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Kernel/LWN: Brute/GRKERNSEC_BRUTE, Chroot and Lockless Patterns

Submitted by Roy Schestowitz on Thursday 25th of March 2021 02:08:41 PM Filed under
Linux
  • Handling brute force attacks in the kernel

    A number of different attacks against Linux systems rely on brute-force techniques using the fork() system call, so a new Linux security module (LSM), called "Brute", has been created to detect and thwart such attacks. Repeated fork() calls can be used for various types of attacks, such as exploiting the Stack Clash vulnerability or Heartbleed-style flaws. Version 6 of the Brute patch set was recently posted and looks like it might be heading toward the mainline.

    This patch set has been in the works since it was first posted as an RFC by John Wood in September 2020 (the resend from Kees Cook a few days later may make it easier to see the whole set). It was originally called "fork brute force attack mitigation" or "fbfam", but that name was deemed too cryptic by Jann Horn and Cook. In addition, it was suggested that turning it into an LSM would be desirable. Both of those suggestions were adopted in version 2, which was posted in October.

    But the idea goes back a lot further than that. The grsecurity kernel patches have long had the GRKERNSEC_BRUTE feature to mitigate brute-force exploits of server programs that use fork() as well as exploits of setuid/setgid binaries. A patch from Richard Weinberger in 2014 used a similar technique to delay fork() calls if forked children die due to a fatal error (which may imply it is part of an attack). That effort was not pushed further, so Cook added an issue to the kernel self-protection project (KSPP) GitHub repository, which is where Wood picked up the idea.

  • Unprivileged chroot()

    It is probably fair to say that most Linux developers never end up using chroot() in an application. This system call puts the calling process into a new view of the filesystem, with the passed-in directory as the root directory. It can be used to isolate a process from the bulk of the filesystem, though its security benefits are somewhat limited. Calling chroot() is a privileged operation but, if Mickaël Salaün has his way with this patch set, that will not be true for much longer, in some situations at least.
    Typically, chroot() is used for tasks like "jailing" a network daemon process; should that process be compromised, its ability to access the filesystem will be limited to the directory tree below the new root. The resulting security boundary is not the strongest — there are a number of ways to break out of chroot() jails — but it can still present a barrier to attackers. chroot() can also be used to create a different view of the file system to, for example, run containers within.

    This system call is not available to just anybody; the CAP_SYS_CHROOT capability is required to be able to call chroot(). This restriction is in place to thwart attackers who would otherwise try to confuse (and exploit) setuid programs by running them inside a specially crafted filesystem tree. As a simple example, consider the sort of mayhem that might be possible if setuid programs saw a version of /etc/passwd or /etc/sudoers that was created by an attacker.

    The limitations of chroot() have long limited its applicability; in recent years it has fallen even further out of favor. Mount namespaces are a much more flexible mechanism for creating new views of the filesystem; they can also be harder to break out of. So relatively few developers see a reason to use chroot() for anything new.

  • Lockless patterns: an introduction to compare-and-swap

    In the first part of this series, I showed you the theory behind concurrent memory models and how that theory can be applied to simple loads and stores. However, loads and stores alone are not a practical tool for the building of higher-level synchronization primitives such as spinlocks, mutexes, and condition variables. Even though it is possible to synchronize two threads using the full memory-barrier pattern that was introduced last week (Dekker's algorithm), modern processors provide a way that is easier, more generic, and faster—yes, all three of them—the compare-and-swap operation.

    [...]

    cmpxchg() loads the value pointed to by *ptr and, if it is equal to old, it stores new in its place. Otherwise, no store happens. The value that was loaded is then returned, regardless of whether it matched old or not. The compare and the store are atomic: if the store is performed, you are guaranteed that no thread could sneak in and write a value other than old to *ptr. Because a single operation provides the old version of the value and stores a new one, compare-and-swap is said to be an atomic read-modify-write operation.

    In Linux, the cmpxchg() macro puts strong ordering requirements on the surrounding code. A compare-and-swap operation comprises a load and a store; for the sake of this article, you can consider them to be, respectively, load-acquire and store-release operations. This means that cmpxchg() can synchronize with both load-acquire or store-release operations performed on the same location by other threads.

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