Post Syndicated from Matthew Garrett original https://mjg59.dreamwidth.org/69070.html

Libraries are collections of code that are intended to be usable by multiple consumers (if you’re interested in the etymology, watch this video). In the old days we had what we now refer to as “static” libraries, collections of code that existed on disk but which would be copied into newly compiled binaries. We’ve moved beyond that, thankfully, and now make use of what we call “dynamic” or “shared” libraries – instead of the code being copied into the binary, a reference to the library function is incorporated, and at runtime the code is mapped from the on-disk copy of the shared object[1]. This allows libraries to be upgraded without needing to modify the binaries using them, and if multiple applications are using the same library at once it only requires that one copy of the code be kept in RAM.

But for this to work, two things are necessary: when we build a binary, there has to be a way to reference the relevant library functions in the binary; and when we run a binary, the library code needs to be mapped into the process.

(I’m going to somewhat simplify the explanations from here on – things like symbol versioning make this a bit more complicated but aren’t strictly relevant to what I was working on here)

For the first of these, the goal is to replace a call to a function (eg, printf()) with a reference to the actual implementation. This is the job of the linker rather than the compiler (eg, if you use the -c argument to tell gcc to simply compile to an object rather than linking an executable, it’s not going to care about whether or not every function called in your code actually exists or not – that’ll be figured out when you link all the objects together), and the linker needs to know which symbols (which aren’t just functions – libraries can export variables or structures and so on) are available in which libraries. You give the linker a list of libraries, it extracts the symbols available, and resolves the references in your code with references to the library.

But how is that information extracted? Each ELF object has a fixed-size header that contains references to various things, including a reference to a list of “section headers”. Each section has a name and a type, but the ones we’re interested in are .dynstr and .dynsym. .dynstr contains a list of strings, representing the name of each exported symbol. .dynsym is where things get more interesting – it’s a list of structs that contain information about each symbol. This includes a bunch of fairly complicated stuff that you need to care about if you’re actually writing a linker, but the relevant entries for this discussion are an index into .dynstr (which means the .dynsym entry isn’t sufficient to know the name of a symbol, you need to extract that from .dynstr), along with the location of that symbol within the library. The linker can parse this information and obtain a list of symbol names and addresses, and can now replace the call to printf() with a reference to libc instead.

(Note that it’s not possible to simply encode this as “Call this address in this library” – if the library is rebuilt or is a different version, the function could move to a different location)

Experimentally, .dynstr and .dynsym appear to be sufficient for linking a dynamic library at build time – there are other sections related to dynamic linking, but you can link against a library that’s missing them. Runtime is where things get more complicated.

When you run a binary that makes use of dynamic libraries, the code from those libraries needs to be mapped into the resulting process. This is the job of the runtime dynamic linker, or RTLD[2]. The RTLD needs to open every library the process requires, map the relevant code into the process’s address space, and then rewrite the references in the binary into calls to the library code. This requires more information than is present in .dynstr and .dynsym – at the very least, it needs to know the list of required libraries.

There’s a separate section called .dynamic that contains another list of structures, and it’s the data here that’s used for this purpose. For example, .dynamic contains a bunch of entries of type DT_NEEDED – this is the list of libraries that an executable requires. There’s also a bunch of other stuff that’s required to actually make all of this work, but the only thing I’m going to touch on is DT_HASH. Doing all this re-linking at runtime involves resolving the locations of a large number of symbols, and if the only way you can do that is by reading a list from .dynsym and then looking up every name in .dynstr that’s going to take some time. The DT_HASH entry points to a hash table – the RTLD hashes the symbol name it’s trying to resolve, looks it up in that hash table, and gets the symbol entry directly (it still needs to resolve that against .dynstr to make sure it hasn’t hit a hash collision – if it has it needs to look up the next hash entry, but this is still generally faster than walking the entire .dynsym list to find the relevant symbol). There’s also DT_GNU_HASH which fulfills the same purpose as DT_HASH but uses a more complicated algorithm that preforms even better. .dynamic also contains entries pointing at .dynstr and .dynsym, which seems redundant but will become relevant shortly.

So, .dynsym and .dynstr are required at build time, and bother are required along with .dynamic at runtime. This seems simple enough, but obviously there’s a twist and I’m sorry it’s taken so long to get to this point.

I bought a Synology NAS for home backup purposes (my previous solution was a single external USB drive plugged into a small server, which had uncomfortable single point of failure properties). Obviously I decided to poke around at it, and I found something odd – all the libraries Synology ships were entirely lacking any ELF section headers. This meant no .dynstr, .dynsym or .dynamic sections, so how was any of this working? nm asserted that the libraries exported no symbols, and readelf agreed. If I wrote a small app that called a function in one of the libraries and built it, gcc complained that the function was undefined. But executables on the device were clearly resolving the symbols at runtime, and if I loaded them into ghidra the exported functions were visible. If I dlopen()ed them, dlsym() couldn’t resolve the symbols – but if I hardcoded the offset into my code, I could call them directly.

Things finally made sense when I discovered that if I passed the --use-dynamic argument to readelf, I did get a list of exported symbols. It turns out that ELF is weirder than I realised. As well as the aforementioned section headers, ELF objects also include a set of program headers. One of the program header types is PT_DYNAMIC. This typically points to the same data that’s present in the .dynamic section. Remember when I mentioned that .dynamic contained references to .dynsym and .dynstr? This means that simply pointing at .dynamic is sufficient, there’s no need to have separate entries for them.

The same information can be reached from two different locations. The information in the section headers is used at build time, and the information in the program headers at run time[3]. I do not have an explanation for this. But if the information is present in two places, it seems obvious that it should be able to reconstruct the missing section headers in my weird libraries? So that’s what this does. It extracts information from the DYNAMIC entry in the program headers and creates equivalent section headers.

There’s one thing that makes this more difficult than it might seem. The section header for .dynsym has to contain the number of symbols present in the section. And that information doesn’t directly exist in DYNAMIC – to figure out how many symbols exist, you’re expected to walk the hash tables and keep track of the largest number you’ve seen. Since every symbol has to be referenced in the hash table, once you’ve hit every entry the largest number is the number of exported symbols. This seemed annoying to implement, so instead I cheated, added code to simply pass in the number of symbols on the command line, and then just parsed the output of readelf against the original binaries to extract that information and pass it to my tool.

Somehow, this worked. I now have a bunch of library files that I can link into my own binaries to make it easier to figure out how various things on the Synology work. Now, could someone explain (a) why this information is present in two locations, and (b) why the build-time linker and run-time linker disagree on the canonical source of truth?

[1] “Shared object” is the source of the .so filename extension used in various Unix-style operating systems
[2] You’ll note that “RTLD” is not an acryonym for “runtime dynamic linker”, because reasons
[3] For environments using the GNU RTLD, at least – I have no idea whether this is the case in all ELF environments

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Post Syndicated from Matthew Garrett original https://mjg59.dreamwidth.org/68721.html

Earlier this year, after Github accidentally committed their private RSA SSH host key to a public repository, I wrote about how better support for SSH host certificates would allow this sort of situation to be handled in a user-transparent way without any negative impact on security. I was hoping that someone would read this and be inspired to fix the problem but sadly that didn’t happen so I’ve actually written some code myself.

The core part of this is straightforward – if a server presents you with a certificate associated with a host key, then make the trust in that host be whoever signed the certificate rather than just trusting the host key. This means that if someone needs to replace the host key for any reason (such as, for example, them having published the private half), you can replace the host key with a new key and a new certificate, and as long as the new certificate is signed by the same key that the previous certificate was, you’ll trust the new key and key rotation can be carried out without any user errors. Hurrah!

So obviously I wrote that bit and then thought about the failure modes and it turns out there’s an obvious one – if an attacker obtained both the private key and the certificate, what stops them from continuing to use it? The certificate isn’t a secret, so we basically have to assume that anyone who possesses the private key has access to it. We may have silently transitioned to a new host key on the legitimate servers, but a hostile actor able to MITM a user can keep on presenting the old key and the old certificate until it expires.

There’s two ways to deal with this – either have short-lived certificates (ie, issue a new certificate every 24 hours or so even if you haven’t changed the key, and specify that the certificate is invalid after those 24 hours), or have a mechanism to revoke the certificates. The former is viable if you have a very well-engineered certificate issuing operation, but still leaves a window for an attacker to make use of the certificate before it expires. The latter is something SSH has support for, but the spec doesn’t define any mechanism for distributing revocation data.

So, I’ve implemented a new SSH protocol extension that allows a host to send a key revocation list to a client. The idea is that the client authenticates to the server, receives a key revocation list, and will no longer trust any certificates that are contained within that list. This seems simple enough, but a naive implementation opens the client to various DoS attacks. For instance, if you simply revoke any key contained within the received KRL, a hostile server could revoke any certificates that were otherwise trusted by the client. The easy way around this is for the client to ensure that any revoked keys are associated with the same CA that signed the host certificate – that way a compromised host can only revoke certificates associated with that CA, and can’t interfere with anyone else.

Unfortunately that still means that a single compromised host can still trigger revocation of certificates inside that trust domain (ie, a compromised host a.test.com could push a KRL that invalidated the certificate for b.test.com), because there’s no way in the KRL format to indicate that a given revocation is associated with a specific hostname. This means we need a mechanism to verify that the KRL update is legitimate, and the easiest way to handle that is to sign it. The KRL format specifies an in-band signature but this was deprecated earlier this year – instead KRLs are supposed to be signed with the sshsig format. But we control both the server and the client, which means it’s easy enough to send a detached signature as part of the extension data.

Putting this all together: you ssh to a server you’ve never contacted before, and it presents you with a host certificate. Instead of the host key being added to known_hosts, the CA key associated with the certificate is added. From now on, if you ssh to that host and it presents a certificate signed by that CA, it’ll be trusted. Optionally, the host can also send you a KRL and a signature. If the signature is generated by the CA key that you already trust, any certificates in that KRL associated with that CA key will be incorporated into local storage. The expected flow if a key is compromised is that the owner of the host generates a new keypair, obtains a new certificate for the new key, and adds the old certificate to a KRL that is signed with the CA key. The next time the user connects to that host, they receive the new key and new certificate, trust it because it’s signed by the same CA key, and also receive a KRL signed with the same CA that revokes trust in the old certificate.

Obviously this breaks down if a user is MITMed with a compromised key and certificate immediately after the host is compromised – they’ll see a legitimate certificate and won’t receive any revocation list, so will trust the host. But this is the same failure mode that would occur in the absence of keys, where the attacker simply presents the compromised key to the client before trust in the new key has been created. This seems no worse than the status quo, but means that most users will seamlessly transition to a new key and revoke trust in the old key with no effort on their part.

The work in progress tree for this is here – at the point of writing I’ve merely implemented this and made sure it builds, not verified that it actually works or anything. Cleanup should happen over the next few days, and I’ll propose this to upstream if it doesn’t look like there’s any showstopper design issues.

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Post Syndicated from Matthew Garrett original https://mjg59.dreamwidth.org/68537.html

There’s a decent number of laptops with fingerprint readers that are supported by Linux, and Gnome has some nice integration to make use of that for authentication purposes. But if you log in with a fingerprint, the moment you start any app that wants to access stored passwords you’ll get a prompt asking you to type in your password, which feels like it somewhat defeats the point. Mac users don’t have this problem – authenticate with TouchID and all your passwords are available after login. Why the difference?

Fingerprint detection can be done in two primary ways. The first is that a fingerprint reader is effectively just a scanner – it passes a graphical representation of the fingerprint back to the OS and the OS decides whether or not it matches an enrolled finger. The second is for the fingerprint reader to make that determination itself, either storing a set of trusted fingerprints in its own storage or supporting being passed a set of encrypted images to compare against. Fprint supports both of these, but note that in both cases all that we get at the end of the day is a statement of “The fingerprint matched” or “The fingerprint didn’t match” – we can’t associate anything else with that.

Apple’s solution involves wiring the fingerprint reader to a secure enclave, an independently running security chip that can store encrypted secrets or keys and only release them under pre-defined circumstances. Rather than the fingerprint reader providing information directly to the OS, it provides it to the secure enclave. If the fingerprint matches, the secure enclave can then provide some otherwise secret material to the OS. Critically, if the fingerprint doesn’t match, the enclave will never release this material.

And that’s the difference. When you perform TouchID authentication, the secure enclave can decide to release a secret that can be used to decrypt your keyring. We can’t easily do this under Linux because we don’t have an interface to store those secrets. The secret material can’t just be stored on disk – that would allow anyone who had access to the disk to use that material to decrypt the keyring and get access to the passwords, defeating the object. We can’t use the TPM because there’s no secure communications channel between the fingerprint reader and the TPM, so we can’t configure the TPM to release secrets only if an associated fingerprint is provided.

So the simple answer is that fingerprint unlock doesn’t unlock the keyring because there’s currently no secure way to do that. It’s not intransigence on the part of the developers or a conspiracy to make life more annoying. It’d be great to fix it, but I don’t see an easy way to do so at the moment.

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Post Syndicated from Matthew Garrett original https://mjg59.dreamwidth.org/68350.html

“Why does ACPI exist” – – the greatest thread in the history of forums, locked by a moderator after 12,239 pages of heated debate, wait no let me start again.

Why does ACPI exist? In the beforetimes power management on x86 was done by jumping to an opaque BIOS entry point and hoping it would do the right thing. It frequently didn’t. We called this Advanced Power Management (Advanced because before this power management involved custom drivers for every machine and everyone agreed that this was a bad idea), and it involved the firmware having to save and restore the state of every piece of hardware in the system. This meant that assumptions about hardware configuration were baked into the firmware – failed to program your graphics card exactly the way the BIOS expected? Hurrah! It’s only saved and restored a subset of the state that you configured and now potential data corruption for you. The developers of ACPI made the reasonable decision that, well, maybe since the OS was the one setting state in the first place, the OS should restore it.

So far so good. But some state is fundamentally device specific, at a level that the OS generally ignores. How should this state be managed? One way to do that would be to have the OS know about the device specific details. Unfortunately that means you can’t ship the computer without having OS support for it, which means having OS support for every device (exactly what we’d got away from with APM). This, uh, was not an option the PC industry seriously considered. The alternative is that you ship something that abstracts the details of the specific hardware and makes that abstraction available to the OS. This is what ACPI does, and it’s also what things like Device Tree do. Both provide static information about how the platform is configured, which can then be consumed by the OS and avoid needing device-specific drivers or configuration to be built-in.

The main distinction between Device Tree and ACPI is that Device Tree is purely a description of the hardware that exists, and so still requires the OS to know what’s possible – if you add a new type of power controller, for instance, you need to add a driver for that to the OS before you can express that via Device Tree. ACPI decided to include an interpreted language to allow vendors to expose functionality to the OS without the OS needing to know about the underlying hardware. So, for instance, ACPI allows you to associate a device with a function to power down that device. That function may, when executed, trigger a bunch of register accesses to a piece of hardware otherwise not exposed to the OS, and that hardware may then cut the power rail to the device to power it down entirely. And that can be done without the OS having to know anything about the control hardware.

How is this better than just calling into the firmware to do it? Because the fact that ACPI declares that it’s going to access these registers means the OS can figure out that it shouldn’t, because it might otherwise collide with what the firmware is doing. With APM we had no visibility into that – if the OS tried to touch the hardware at the same time APM did, boom, almost impossible to debug failures (This is why various hardware monitoring drivers refuse to load by default on Linux – the firmware declares that it’s going to touch those registers itself, so Linux decides not to in order to avoid race conditions and potential hardware damage. In many cases the firmware offers a collaborative interface to obtain the same data, and a driver can be written to get that. this bug comment discusses this for a specific board)

Unfortunately ACPI doesn’t entirely remove opaque firmware from the equation – ACPI methods can still trigger System Management Mode, which is basically a fancy way to say “Your computer stops running your OS, does something else for a while, and you have no idea what”. This has all the same issues that APM did, in that if the hardware isn’t in exactly the state the firmware expects, bad things can happen. While historically there were a bunch of ACPI-related issues because the spec didn’t define every single possible scenario and also there was no conformance suite (eg, should the interpreter be multi-threaded? Not defined by spec, but influences whether a specific implementation will work or not!), these days overall compatibility is pretty solid and the vast majority of systems work just fine – but we do still have some issues that are largely associated with System Management Mode.

One example is a recent Lenovo one, where the firmware appears to try to poke the NVME drive on resume. There’s some indication that this is intended to deal with transparently unlocking self-encrypting drives on resume, but it seems to do so without taking IOMMU configuration into account and so things explode. It’s kind of understandable why a vendor would implement something like this, but it’s also kind of understandable that doing so without OS cooperation may end badly.

This isn’t something that ACPI enabled – in the absence of ACPI firmware vendors would just be doing this unilaterally with even less OS involvement and we’d probably have even more of these issues. Ideally we’d “simply” have hardware that didn’t support transitioning back to opaque code, but we don’t (ARM has basically the same issue with TrustZone). In the absence of the ideal world, by and large ACPI has been a net improvement in Linux compatibility on x86 systems. It certainly didn’t remove the “Everything is Windows” mentality that many vendors have, but it meant we largely only needed to ensure that Linux behaved the same way as Windows in a finite number of ways (ie, the behaviour of the ACPI interpreter) rather than in every single hardware driver, and so the chances that a new machine will work out of the box are much greater than they were in the pre-ACPI period.

There’s an alternative universe where we decided to teach the kernel about every piece of hardware it should run on. Fortunately (or, well, unfortunately) we’ve seen that in the ARM world. Most device-specific simply never reaches mainline, and most users are stuck running ancient kernels as a result. Imagine every x86 device vendor shipping their own kernel optimised for their hardware, and now imagine how well that works out given the quality of their firmware. Does that really seem better to you?

It’s understandable why ACPI has a poor reputation. But it’s also hard to figure out what would work better in the real world. We could have built something similar on top of Open Firmware instead but the distinction wouldn’t be terribly meaningful – we’d just have Forth instead of the ACPI bytecode language. Longing for a non-ACPI world without presenting something that’s better and actually stands a reasonable chance of adoption doesn’t make the world a better place.

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Post Syndicated from Matthew Garrett original https://mjg59.dreamwidth.org/68004.html

The Free Software Foundation Europe and the Software Freedom Conservancy recently released a statement that they would no longer work with Eben Moglen, chairman of the Software Freedom Law Center. Eben was the general counsel for the Free Software Foundation for over 20 years, and was centrally involved in the development of version 3 of the GNU General Public License. He’s devoted a great deal of his life to furthering free software.

But, as described in the joint statement, he’s also acted abusively towards other members of the free software community. He’s acted poorly towards his own staff. In a professional context, he’s used graphically violent rhetoric to describe people he dislikes. He’s screamed abuse at people attempting to do their job.

And, sadly, none of this comes as a surprise to me. As I wrote in 2017, after it became clear that Eben’s opinions diverged sufficiently from the FSF’s that he could no longer act as general counsel, he responded by threatening an FSF board member at an FSF-run event (various members of the board were willing to tolerate this, which is what led to me quitting the board). There’s over a decade’s evidence of Eben engaging in abusive behaviour towards members of the free software community, be they staff, colleagues, or just volunteers trying to make the world a better place.

When we build communities that tolerate abuse, we exclude anyone unwilling to tolerate being abused[1]. Nobody in the free software community should be expected to deal with being screamed at or threatened. Nobody should be afraid that they’re about to have their sexuality outed by a former boss.

But of course there are some that will defend Eben based on his past contributions. There were people who were willing to defend Hans Reiser on that basis. We need to be clear that what these people are defending is not free software – it’s the right for abusers to abuse. And in the long term, that’s bad for free software.

[1] “Why don’t people just get better at tolerating abuse?” is a terrible response to this. Why don’t abusers stop abusing? There’s fewer of them, and it should be easier.

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Post Syndicated from Matthew Garrett original https://mjg59.dreamwidth.org/67602.html

TPMs contain a set of registers (“Platform Configuration Registers”, or PCRs) that are used to track what a system boots. Each time a new event is measured, a cryptographic hash representing that event is passed to the TPM. The TPM appends that hash to the existing value in the PCR, hashes that, and stores the final result in the PCR. This means that while the PCR’s value depends on the precise sequence and value of the hashes presented to it, the PCR value alone doesn’t tell you what those individual events were. Different PCRs are used to store different event types, but there are still more events than there are PCRs so we can’t avoid this problem by simply storing each event separately.

This is solved using the event log. The event log is simply a record of each event, stored in RAM. The algorithm the TPM uses to calculate the PCR values is known, so we can reproduce that by simply taking the events from the event log and replaying the series of events that were passed to the TPM. If the final calculated value is the same as the value in the PCR, we know that the event log is accurate, which means we now know the value of each individual event and can make an appropriate judgement regarding its security.

If any value in the event log is invalid, we’ll calculate a different PCR value and it won’t match. This isn’t terribly helpful – we know that at least one entry in the event log doesn’t match what was passed to the TPM, but we don’t know which entry. That means we can’t trust any of the events associated with that PCR. If you’re trying to make a security determination based on this, that’s going to be a problem.

PCR 7 is used to track information about the secure boot policy on the system. It contains measurements of whether or not secure boot is enabled, and which keys are trusted and untrusted on the system in question. This is extremely helpful if you want to verify that a system booted with secure boot enabled before allowing it to do something security or safety critical. Unfortunately, if the device gives you an event log that doesn’t replay correctly for PCR 7, you now have no idea what the security state of the system is.

We ran into that this week. Examination of the event log revealed an additional event other than the expected ones – a measurement accompanied by the string “Boot Guard Measured S-CRTM”. Boot Guard is an Intel feature where the CPU verifies the firmware is signed with a trusted key before executing it, and measures information about the firmware in the process. Previously I’d only encountered this as a measurement into PCR 0, which is the PCR used to track information about the firmware itself. But it turns out that at least some versions of Boot Guard also measure information about the Boot Guard policy into PCR 7. The argument for this is that this is effectively part of the secure boot policy – having a measurement of the Boot Guard state tells you whether Boot Guard was enabled, which tells you whether or not the CPU verified a signature on your firmware before running it (as I wrote before, I think Boot Guard has user-hostile default behaviour, and that enforcing this on consumer devices is a bad idea).

But there’s a problem here. The event log is created by the firmware, and the Boot Guard measurements occur before the firmware is executed. So how do we get a log that represents them? That one’s fairly simple – the firmware simply re-calculates the same measurements that Boot Guard did and creates a log entry after the fact[1]. All good.

Except. What if the firmware screws up the calculation and comes up with a different answer? The entry in the event log will now not match what was sent to the TPM, and replaying will fail. And without knowing what the actual value should be, there’s no way to fix this, which means there’s no way to verify the contents of PCR 7 and determine whether or not secure boot was enabled.

But there’s still a fundamental source of truth – the measurement that was sent to the TPM in the first place. Inspired by Henri Nurmi’s work on sniffing Bitlocker encryption keys, I asked a coworker if we could sniff the TPM traffic during boot. The TPM on the board in question uses SPI, a simple bus that can have multiple devices connected to it. In this case the system flash and the TPM are on the same SPI bus, which made things easier. The board had a flash header for external reprogramming of the firmware in the event of failure, and all SPI traffic was visible through that header. Attaching a logic analyser to this header made it simple to generate a record of that. The only problem was that the chip select line on the header was attached to the firmware flash chip, not the TPM. This was worked around by simply telling the analysis software that it should invert the sense of the chip select line, ignoring all traffic that was bound for the flash and paying attention to all other traffic. This worked in this case since the only other device on the bus was the TPM, but would cause problems in the event of multiple devices on the bus all communicating.

With the aid of this analyser plugin, I was able to dump all the TPM traffic and could then search for writes that included the “0182” sequence that corresponds to the command code for a measurement event. This gave me a couple of accesses to the locality 3 registers, which was a strong indication that they were coming from the CPU rather than from the firmware. One was for PCR 0, and one was for PCR 7. This corresponded to the two Boot Guard events that we expected from the event log. The hash in the PCR 0 measurement was the same as the hash in the event log, but the hash in the PCR 7 measurement differed from the hash in the event log. Replacing the event log value with the value actually sent to the TPM resulted in the event log now replaying correctly, supporting the hypothesis that the firmware was failing to correctly reconstruct the event.

What now? The simple thing to do is for us to simply hard code this fixup, but longer term we’d like to figure out how to reconstruct the event so we can calculate the expected value ourselves. Unfortunately there doesn’t seem to be any public documentation on this. Sigh.

[1] What stops firmware on a system with no Boot Guard faking those measurements? TPMs have a concept of “localities”, effectively different privilege levels. When Boot Guard performs its initial measurement into PCR 0, it does so at locality 3, a locality that’s only available to the CPU. This causes PCR 0 to be initialised to a different initial value, affecting the final PCR value. The firmware can’t access locality 3, so can’t perform an equivalent measurement, so can’t fake the value.

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Post Syndicated from Matthew Garrett original https://mjg59.dreamwidth.org/67402.html

Work involves supporting Windows (there’s a lot of specialised hardware design software that’s only supported under Windows, so this isn’t really avoidable), but also involves git, so I’ve been working on extending our support for hardware-backed SSH certificates to Windows and trying to glue that into git. In theory this doesn’t sound like a hard problem, but in practice oh good heavens.

Git for Windows is built on top of msys2, which in turn is built on top of Cygwin. This is an astonishing artifact that allows you to build roughly unmodified POSIXish code on top of Windows, despite the terrible impedance mismatches inherent in this. One is that until 2017, Windows had no native support for Unix sockets. That’s kind of a big deal for compatibility purposes, so Cygwin worked around it. It’s, uh, kind of awful. If you’re not a Cygwin/msys app but you want to implement a socket they can communicate with, you need to implement this undocumented protocol yourself. This isn’t impossible, but ugh.

But going to all this trouble helps you avoid another problem! The Microsoft version of OpenSSH ships an SSH agent that doesn’t use Unix sockets, but uses a named pipe instead. So if you want to communicate between Cygwinish OpenSSH (as is shipped with git for Windows) and the SSH agent shipped with Windows, you need something that bridges between those. The state of the art seems to be to use npiperelay with socat, but if you’re already writing something that implements the Cygwin socket protocol you can just use npipe to talk to the shipped ssh-agent and then export your own socket interface.

And, amazingly, this all works? I’ve managed to hack together an SSH agent (using Go’s SSH agent implementation) that can satisfy hardware backed queries itself, but forward things on to the Windows agent for compatibility with other tooling. Now I just need to figure out how to plumb it through to WSL. Sigh.

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Post Syndicated from Matthew Garrett original https://mjg59.dreamwidth.org/67126.html

I dug out a computer running Fedora 28, which was released 2018-04-01 – over 5 years ago. Backing up the data and re-installing seemed tedious, but the current version of Fedora is 38, and while Fedora supports updates from N to N+2 that was still going to be 5 separate upgrades. That seemed tedious, so I figured I’d just try to do an update from 28 directly to 38. This is, obviously, extremely unsupported, but what could possibly go wrong?

Running sudo dnf system-upgrade download --releasever=38 didn’t successfully resolve dependencies, but sudo dnf system-upgrade download --releasever=38 --allowerasing passed and dnf started downloading 6GB of packages. And then promptly failed, since I didn’t have any of the relevant signing keys. So I downloaded the fedora-gpg-keys package from F38 by hand and tried to install it, and got a signature hdr data: BAD, no. of bytes(88084) out of range error. It turns out that rpm doesn’t handle cases where the signature header is larger than a few K, and RPMs from modern versions of Fedora. The obvious fix would be to install a newer version of rpm, but that wouldn’t be easy without upgrading the rest of the system as well – or, alternatively, downloading a bunch of build depends and building it. Given that I’m already doing all of this in the worst way possible, let’s do something different.

The relevant code in the hdrblobRead function of rpm’s lib/header.c is:

int32_t il_max = HEADER_TAGS_MAX;
int32_t dl_max = HEADER_DATA_MAX;

if (regionTag == RPMTAG_HEADERSIGNATURES) {
il_max = 32;
dl_max = 8192;
}

which indicates that if the header in question is RPMTAG_HEADERSIGNATURES, it sets more restrictive limits on the size (no, I don’t know why). So I installed rpm-libs-debuginfo, ran gdb against librpm.so.8, loaded the symbol file, and then did disassemble hdrblobRead. The relevant chunk ends up being:

0x000000000001bc81 <+81>: cmp $0x3e,%ebx
0x000000000001bc84 <+84>: mov $0xfffffff,%ecx
0x000000000001bc89 <+89>: mov $0x2000,%eax
0x000000000001bc8e <+94>: mov %r12,%rdi
0x000000000001bc91 <+97>: cmovne %ecx,%eax

which is basically “If ebx is not 0x3e, set eax to 0xffffffff – otherwise, set it to 0x2000”. RPMTAG_HEADERSIGNATURES is 62, which is 0x3e, so I just opened librpm.so.8 in hexedit, went to byte 0x1bc81, and replaced 0x3e with 0xfe (an arbitrary invalid value). This has the effect of skipping the if (regionTag == RPMTAG_HEADERSIGNATURES) code and so using the default limits even if the header section in question is the signatures. And with that one byte modification, rpm from F28 would suddenly install the fedora-gpg-keys package from F38. Success!

But short-lived. dnf now believed packages had valid signatures, but sadly there were still issues. A bunch of packages in F38 had files that conflicted with packages in F28. These were largely Python 3 packages that conflicted with Python 2 packages from F28 – jumping this many releases meant that a bunch of explicit replaces and the like no longer existed. The easiest way to solve this was simply to uninstall python 2 before upgrading, and avoiding the entire transition. Another issue was that some data files had moved from libxcrypt-common to libxcrypt, and removing libxcrypt-common would remove libxcrypt and a bunch of important things that depended on it (like, for instance, systemd). So I built a fake empty package that provided libxcrypt-common and removed the actual package. Surely everything would work now?

Ha no. The final obstacle was that several packages depended on rpmlib(CaretInVersions), and building another fake package that provided that didn’t work. I shouted into the void and Bill Nottingham answered – rpmlib dependencies are synthesised by rpm itself, indicating that it has the ability to handle extensions that specific packages are making use of. This made things harder, since the list is hard-coded in the binary. But since I’m already committing crimes against humanity with a hex editor, why not go further? Back to editing librpm.so.8 and finding the list of rpmlib() dependencies it provides. There were a bunch, but I couldn’t really extend the list. What I could do is overwrite existing entries. I tried this a few times but (unsurprisingly) broke other things since packages depended on the feature I’d overwritten. Finally, I rewrote rpmlib(ExplicitPackageProvide) to rpmlib(CaretInVersions) (adding an extra ‘\0’ at the end of it to deal with it being shorter than the original string) and apparently nothing I wanted to install depended on rpmlib(ExplicitPackageProvide) because dnf finished its transaction checks and prompted me to reboot to perform the update. So, I did.

And about an hour later, it rebooted and gave me a whole bunch of errors due to the fact that dbus never got started. A bit of digging revealed that I had no /etc/systemd/system/dbus.service, a symlink that was presumably introduced at some point between F28 and F38 but which didn’t get automatically added in my case because well who knows. That was literally the only thing I needed to fix up after the upgrade, and on the next reboot I was presented with a gdm prompt and had a fully functional F38 machine.

You should not do this. I should not do this. This was a terrible idea. Any situation where you’re binary patching your package manager to get it to let you do something is obviously a bad situation. And with hindsight performing 5 independent upgrades might have been faster. But that would have just involved me typing the same thing 5 times, while this way I learned something. And what I learned is “Terrible ideas sometimes work and so you should definitely act upon them rather than doing the sensible thing”, so like I said, you should not do this in case you learn the same lesson.

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