Kernel development

Brief items

Kernel release status

The current development kernel is 3.8-rc2, released on January 2. Linus says: "It's a new year, people are getting back to work, and trying desperately to forget the over-eating that has been going on for the last two weeks. And hey, to celebrate, here's -rc2!" Perhaps as a result of the aforementioned over-eating, this patch is relatively small; things might be expected to pick up a bit by the time -rc3 comes out.

The 3.8-rc1 release happened on December 21. The most significant changes pulled between the last edition's summary and the closing of the merge window include the f2fs filesystem, and a driver for Dialog Semiconductor DA9055 watchdog devices.

Stable updates: no stable updates have been released since December 17. The 3.2.36 update is in the review process as of this writing; its release can be expected at any time.

Comments (none posted)

Quotes of the week

— Frederic Weisbecker

— Thomas Gleixner

— Linus Torvalds

— Russell King

Comments (42 posted)

The Tux3 filesystem returns

The "Tux3" next-generation filesystem project generated a lot of discussion and a fair amount of code before fading into obscurity; LWN last covered this work in 2008. Tux3 developer Daniel Phillips has resurfaced with a new-year posting suggesting that work on this code is resuming. "In brief, the first time Hirofumi ever put together all the kernel pieces in his magical lab over in Tokyo, our Tux3 rocket took off and made it straight to orbit. Or in less metaphorical terms, our first meaningful benchmarks turned in numbers that meet or even slightly beat the illustrious incumbent, Ext4." The code can be found in git://github.com/OGAWAHirofumi/tux3.git for those who want to play with it.

Full Story (comments: 40)

Kernel development news

Improving ticket spinlocks

Spinlocks, being the lowest-level synchronization mechanism in the kernel, are the target of seemingly endless attempts at performance enhancement. The ticket spinlock mechanism used in the mainline has resisted such attempts for a few years. Now, though, some developers have identified a performance bottleneck associated with these locks and are busily trying to come up with an improved version.

A spinlock is so-named because a CPU waiting for a contended lock will "spin" in a tight loop, repeatedly querying the lock until it becomes available. Ticket spinlocks adjust this algorithm by having each waiting CPU take a "ticket" so that each CPU obtains the lock in the order in which it arrived. These locks thus resemble the "take a number" mechanisms found at deli counters or motor vehicle division offices worldwide — though, with luck, the wait is rather shorter than is required to renew a driver's license in your editor's part of the world. Without the ticket mechanism, which was added for the 2.6.25 release, the kernel's spinlocks were unfair; in some situations, some waiters could be starved for an extended period of time.

It has long been understood that lock contention reduces system performance considerably. The simple act of spinning for a lock clearly is not going to be good for performance, but there are also caching issues to take into account. If two CPUs are repeatedly acquiring a spinlock, the memory location representing that lock will bounce back and forth between those CPUs' caches. Even if neither CPU ever has to wait for the lock, the process of moving it between caches will slow things down considerably. For that reason, interest in lockless algorithms has been growing for many years.

In the case of a contended lock, though, cache contention would appear to be less of an issue. A CPU spinning on a lock will cache its contents in a shared mode; no cache bouncing should occur until the CPU owning the lock releases it. Releasing the lock (and its acquisition by another CPU) requires writing to the lock, and that requires exclusive cache access. The cache line movement at that time hurts, but probably not as much as waiting for the lock in the first place. So it would seem that trying to optimize cache behavior in the contended case is not likely to produce much in the way of useful results.

That picture is not complete, though; one must take a couple of other facts into account. Processors do not cache a single value; they cache a "line" of (typically) 128 consecutive bytes as a single unit. In other words, the cache lines in any contemporary processor are almost certainly significantly larger than what is required to hold a spinlock. So when a CPU needs exclusive access to a spinlock's cache line, it also gains exclusive access to a significant chunk of surrounding data. And that is where the other important detail comes into play: spinlocks tend to be embedded within the data structures that they protect, so that surrounding data is typically data of immediate interest to the CPU holding the lock.

Kernel code will acquire a lock to work with (and, usually, modify) a structure's contents. Often, changing a field within the protected structure will require access to the same cache line that holds the structure's spinlock. If the lock is uncontended, that access is not a problem; the CPU owning the lock probably owns the cache line as well. But if the lock is contended, there will be one or more other CPUs constantly querying its value, obtaining shared access to that same cache line and depriving the lock holder of the exclusive access it needs. A subsequent modification of data within the affected cache line will thus incur a cache miss. So CPUs querying a contended lock can slow the lock owner considerably, even though that owner is not accessing the lock directly.

How badly can throughput be impacted? In the description of his patch adding proportional backoff to ticket spinlocks, Rik van Riel describes a microbenchmark that is slowed by a factor of two when there is a single contending CPU, and by as much as a factor of ten with many CPUs in the mix. That is not just a slowdown; that is a catastrophic loss of performance. Needless to say, that is not the sort of behavior that kernel developers like to see.

Rik's solution is simple enough. Rather than spinning tightly and querying a contended lock's status, a waiting CPU should wait a bit more patiently, only querying the lock occasionally. So his patch causes a waiting CPU to loop a number of times doing nothing at all before it gets impatient and checks the lock again. It goes without saying that picking that "number of times" correctly is the key to good performance with this algorithm. While a CPU is looping without querying the lock it cannot be bouncing cache lines around, so the lock holder should be able to make faster progress. But too much looping will cause the lock to sit idle before the owner of the next ticket notices that its turn has come; that, too, will hurt performance.

The first step in Rik's patch series calculates how many CPUs must release the lock before the current CPU can claim it (by subtracting the current CPU's ticket number from the number currently being served) and loops 50 times for every CPU that is ahead in the queue. That is where the "proportional backoff" comes in; the further back in line the CPU is, the longer it will wait between queries of the lock. The result should be a minimizing of idle looping while also minimizing cache traffic.

The number 50 was determined empirically, but it seems unlikely that it will be optimal for all situations. So the final part of Rik's patch set attempts to tune that number dynamically. The dynamic delay factor is increased when the lock is found to be unavailable and decreased when the lock is obtained. The goal is to have a CPU query the lock an average of 2.7 times before obtaining it. The number 2.7, once again, was obtained by running lots of tests and seeing what worked best; subsequent versions of the patch have tweaked this heuristic somewhat. Details aside, the core idea is that the delay factor (a per-CPU value that applies to all contended locks equally) will increase for workloads experiencing more contention, tuning the system appropriately.

That said, the notion of a single delay for all locks is likely to be causing a severe case of raised eyebrows for some readers, and, indeed, it turned out to be inadequate; some locks are rather more contended than others, after all. So the January 3 version of Rik's patch keeps a hashed list (based on the spinlock address) of delay values instead.

Michel Lespinasse ran some experiments of his own to see how well the proportional backoff algorithm worked. In particular, he wanted to figure out whether it was truly necessary to calculate a dynamic delay factor, or whether an optimal static value could be found. His conclusion was that, in fact, a static value is good enough; it might be possible to do a little better with a dynamic value, he said, but the improvement is not enough to justify the added complexity of the tuning mechanism. There is just one little difficulty:

If these results stand, and an appropriate way of picking the static value can be found, then there is probably not a case for adding dynamic backoff to the kernel's spinlock implementation. But the backoff idea in general would appear to be a significant improvement for some workloads. So the chances are good that we will see it added in some form in an upcoming development cycle.

Comments (17 posted)

The mempressure control group proposal

Last November, LWN described the vmpressure_fd() work which implemented a new system call making it possible for user-space applications to be informed when system memory is tight. Those applications could then respond by freeing memory, easing the crunch. That patch set has since evolved considerably.

Based on the feedback that author Anton Vorontsov received, the concept has changed from a new system call to a new, control-group-based subsystem. The controller implementation allows for integration with the memory controller, meaning that applications can receive notifications when their specific control group is running low on memory, even if the system as a whole is not under memory pressure.

As with previous versions of the patch, applications can receive notifications for three levels of memory pressure: "low" (memory reclaim is happening at a low level), "medium" (some swapping is happening), and "oom" (memory pressure is severe). But these notifications may no longer be the primary way in which applications interact with the controller, thanks to the most significant change in comparison to the previous vmpressure_fd() solution: the addition of a user-space "shrinker" interface allowing the kernel to ask user space to free specific amounts of memory when needed. This API was inspired by Andrew Morton's feedback on the first revision of the mempressure control group subsystem patch:

Andrew also worried that application developers may tune their programs against a particular kernel version; subtle behavioral changes in new kernel releases might then cause regressions. In short, Andrew complained, the behavior of the system as a whole was not testable, so there would be no way to know if subsequent kernel changes made performance worse.

Andrew's suggestion was to give more control to the kernel and introduce some kind of interface for user-space memory scanning and freeing (similar in its main concept to the shrink_slab() kernel shrinkers). This interface would control user-space reclaim behavior; if something goes wrong, it will be up to kernel to resolve the issue. It would also give kernel developers the ability to test and tune whole system behavior by writing a compliant user-space test application and running it.

The user-space shrinker implementation by Anton operates on the concept of chunks of an application-defined size. There is an assumption that the application does memory allocations with a specific granularity (the "chunk size," which may be not 100% accurate but the more accurate it is, the better). So if the application caches data in chunks of 1MB, that is the size it will provide to the shrinker interface. That is done through a sequence like this:

1. The application opens the control interface, which is found as the file cgroup.event_control in the controller directory.

2. The shrinker interface (mempressure.shrinker in the controller directory) must also be opened.

3. The eventfd() system call is used to obtain a third file descriptor for notifications.

4. The application then writes a string containing the eventfd() file descriptor number, the mempressure.shrinker file descriptor number, and the chunk size to the control interface.

Occasionally, the application should write a string to the shrinker file indicating how many chunks have been allocated or (using a negative count) freed. The kernel uses this information to maintain an internal count of how many reclaimable chunks the application is currently holding on to.

If the kernel wants the application to free some memory, the notification will come through the eventfd() file descriptor in the form of an integer count of the number of chunks that should be freed. The kernel assumes that the application will free the specified number of chunks before reading from the eventfd() file descriptor again. If the application isn't able to reclaim all chunks for some reason, it should re-add the number of chunks that were not freed by writing to the mempressure.shrinker file.

The patchset also includes an example application (slightly buggy in the current version) for testing the new interface. It creates two threads; the first thread initializes the user-space shrinker mechanism notifications and then tries to allocate memory (more than physically available) in an infinite loop. The second thread listens for user-space shrinker notifications and frees the requested number of chunks (also in an infinite loop). Ideally, during the run of the test application the system shouldn't get into an out-of-memory condition and it also shouldn't use much swap space (if any is available of course).

Various comments were received on the patch set, so at least one more round of changes will be required before this interface can be considered for merging into the mainline. There is also an open question on how this feature interacts with volatile ranges and whether both mechanisms (neither of which has yet been merged) are truly required. So this discussion may continue well into the new year before we end up with reclaimable user-space memory caches in their final form.

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Namespaces in operation, part 1: namespaces overview

The Linux 3.8 merge window saw the acceptance of Eric Biederman's sizeable series of user namespace and related patches. Although there remain some details to finish—for example, a number of Linux filesystems are not yet user-namespace aware—the implementation of user namespaces is now functionally complete.

The completion of the user namespaces work is something of a milestone, for a number of reasons. First, this work represents the completion of one of the most complex namespace implementations to date, as evidenced by the fact that it has been around five years since the first steps in the implementation of user namespaces (in Linux 2.6.23). Second, the namespace work is currently at something of a "stable point", with the implementation of most of the existing namespaces being more or less complete. This does not mean that work on namespaces has finished: other namespaces may be added in the future, and there will probably be further extensions to existing namespaces, such as the addition of namespace isolation for the kernel log. Finally, the recent changes in the implementation of user namespaces are something of a game changer in terms of how namespaces can be used: starting with Linux 3.8, unprivileged processes can create user namespaces in which they have full privileges, which in turn allows any other type of namespace to be created inside a user namespace.

Thus, the present moment seems a good point to take an overview of namespaces and a practical look at the namespace API. This is the first of a series of articles that does so: in this article, we provide an overview of the currently available namespaces; in the follow-on articles, we'll show how the namespace APIs can be used in programs.

The namespaces

Currently, Linux implements six different types of namespaces. The purpose of each namespace is to wrap a particular global system resource in an abstraction that makes it appear to the processes within the namespace that they have their own isolated instance of the global resource. One of the overall goals of namespaces is to support the implementation of containers, a tool for lightweight virtualization (as well as other purposes) that provides a group of processes with the illusion that they are the only processes on the system.

In the discussion below, we present the namespaces in the order that they were implemented (or at least, the order in which the implementations were completed). The CLONE_NEW* identifiers listed in parentheses are the names of the constants used to identify namespace types when employing the namespace-related APIs (clone(), unshare(), and setns()) that we will describe in our follow-on articles.

Mount namespaces (CLONE_NEWNS, Linux 2.4.19) isolate the set of filesystem mount points seen by a group of processes. Thus, processes in different mount namespaces can have different views of the filesystem hierarchy. With the addition of mount namespaces, the mount() and umount() system calls ceased operating on a global set of mount points visible to all processes on the system and instead performed operations that affected just the mount namespace associated with the calling process.

One use of mount namespaces is to create environments that are similar to chroot jails. However, by contrast with the use of the chroot() system call, mount namespaces are a more secure and flexible tool for this task. Other more sophisticated uses of mount namespaces are also possible. For example, separate mount namespaces can be set up in a master-slave relationship, so that the mount events are automatically propagated from one namespace to another; this allows, for example, an optical disk device that is mounted in one namespace to automatically appear in other namespaces.

Mount namespaces were the first type of namespace to be implemented on Linux, appearing in 2002. This fact accounts for the rather generic "NEWNS" moniker (short for "new namespace"): at that time no one seems to have been thinking that other, different types of namespace might be needed in the future.

UTS namespaces (CLONE_NEWUTS, Linux 2.6.19) isolate two system identifiers—nodename and domainname—returned by the uname() system call; the names are set using the sethostname() and setdomainname() system calls. In the context of containers, the UTS namespaces feature allows each container to have its own hostname and NIS domain name. This can be useful for initialization and configuration scripts that tailor their actions based on these names. The term "UTS" derives from the name of the structure passed to the uname() system call: struct utsname. The name of that structure in turn derives from "UNIX Time-sharing System".

IPC namespaces (CLONE_NEWIPC, Linux 2.6.19) isolate certain interprocess communication (IPC) resources, namely, System V IPC objects and (since Linux 2.6.30) POSIX message queues. The common characteristic of these IPC mechanisms is that IPC objects are identified by mechanisms other than filesystem pathnames. Each IPC namespace has its own set of System V IPC identifiers and its own POSIX message queue filesystem.

PID namespaces (CLONE_NEWPID, Linux 2.6.24) isolate the process ID number space. In other words, processes in different PID namespaces can have the same PID. One of the main benefits of PID namespaces is that containers can be migrated between hosts while keeping the same process IDs for the processes inside the container. PID namespaces also allow each container to have its own init (PID 1), the "ancestor of all processes" that manages various system initialization tasks and reaps orphaned child processes when they terminate.

From the point of view of a particular PID namespace instance, a process has two PIDs: the PID inside the namespace, and the PID outside the namespace on the host system. PID namespaces can be nested: a process will have one PID for each of the layers of the hierarchy starting from the PID namespace in which it resides through to the root PID namespace. A process can see (e.g., view via /proc/PID and send signals with kill()) only processes contained in its own PID namespace and the namespaces nested below that PID namespace.

Network namespaces (CLONE_NEWNET, started in Linux 2.4.19 2.6.24 and largely completed by about Linux 2.6.29) provide isolation of the system resources associated with networking. Thus, each network namespace has its own network devices, IP addresses, IP routing tables, /proc/net directory, port numbers, and so on.

Network namespaces make containers useful from a networking perspective: each container can have its own (virtual) network device and its own applications that bind to the per-namespace port number space; suitable routing rules in the host system can direct network packets to the network device associated with a specific container. Thus, for example, it is possible to have multiple containerized web servers on the same host system, with each server bound to port 80 in its (per-container) network namespace.

User namespaces (CLONE_NEWUSER, started in Linux 2.6.23 and completed in Linux 3.8) isolate the user and group ID number spaces. In other words, a process's user and group IDs can be different inside and outside a user namespace. The most interesting case here is that a process can have a normal unprivileged user ID outside a user namespace while at the same time having a user ID of 0 inside the namespace. This means that the process has full root privileges for operations inside the user namespace, but is unprivileged for operations outside the namespace.

Starting in Linux 3.8, unprivileged processes can create user namespaces, which opens up a raft of interesting new possibilities for applications: since an otherwise unprivileged process can hold root privileges inside the user namespace, unprivileged applications now have access to functionality that was formerly limited to root. Eric Biederman has put a lot of effort into making the user namespaces implementation safe and correct. However, the changes wrought by this work are subtle and wide ranging. Thus, it may happen that user namespaces have some as-yet unknown security issues that remain to be found and fixed in the future.

Concluding remarks

It's now around a decade since the implementation of the first Linux namespace. Since that time, the namespace concept has expanded into a more general framework for isolating a range of global resources whose scope was formerly system-wide. As a result, namespaces now provide the basis for a complete lightweight virtualization system, in the form of containers. As the namespace concept has expanded, the associated API has grown—from a single system call (clone()) and one or two /proc files—to include a number of other system calls and many more files under /proc. The details of that API will form the subject of the follow-ups to this article.

Series index

The following list shows later articles in this series, along with their example programs:

• Part 2: the namespaces API
  • demo_uts_s.c: demonstrate the use of UTS namespaces
  • ns_exec.c: join a namespace using setns() and execute a command
  • unshare.c: unshare namespaces and execute a command; similar in concept to unshare(1)
• Part 3: PID namespaces
  • pidns_init_sleep.c: demonstrate PID namespaces
  • multi_pidns.c: create a series of child processes in nested PID namespaces
• Part 4: more on PID namespaces
  • ns_child_exec.c: create a child process that executes a shell command in new namespace(s)
  • simple_init.c: a simple init(1)-style program to be used as the init program in a PID namespace
  • orphan.c: demonstrate that a child becomes orphaned and is adopted by the init process when its parent exits
  • ns_run.c: join one or more namespaces using setns() and execute a command in those namespaces, possibly inside a child process; similar in concept to nsenter(1)
• Part 5: user namespaces
  • demo_userns.c: simple program to create a user namespace and display process credentials and capabilities
  • userns_child_exec.c: create a child process that executes a shell command in new namespace(s); similar to ns_child_exec.c, but with additional options for use with user namespaces
• Part 6: more on user namespaces
  • userns_setns_test.c: test the operation of setns() from two different user namespaces.

Comments (20 posted)

Patches and updates

Kernel trees

• Linus Torvalds: Linux 3.8-rc2 . (January 3, 2013)

• Linus Torvalds: Linux 3.8-rc1 . (December 22, 2012)

• Thomas Gleixner: 3.6.11-rt24 (apocalypse release)c . (December 21, 2012)

• Steven Rostedt: 3.4.24-rt35 . (December 21, 2012)

• Steven Rostedt: 3.4.24-rt36 . (December 27, 2012)

• Steven Rostedt: 3.2.35-rt53 . (December 27, 2012)

• Steven Rostedt: 3.0.57-rt81 . (December 21, 2012)

• Steven Rostedt: 3.0.57-rt82 . (December 27, 2012)

Core kernel code

• Frederic Weisbecker: 3.7-nohz1 . (December 21, 2012)

• Frederic Weisbecker: 3.8-rc1-nohz1 . (December 29, 2012)

• Anton Arapov: uprobes: return probe implementation . (December 21, 2012)

• Andrey Vagin: signalfd: a kernel interface for dumping/restoring pending signals . (December 27, 2012)

• Kent Overstreet: AIO performance improvements/cleanups, v3 . (December 27, 2012)

Development tools

• Stephane Eranian: perf: add memory access sampling support . (December 21, 2012)

• Jovi Zhang: ktap: Another dynamic tracing tool for Linux . (January 1, 2013)

Device drivers

• Terje Bergstrom: Support for Tegra 2D hardware . (December 21, 2012)

• Kim, Milo: rtc: add new lp8788 rtc driver . (December 21, 2012)

• Laxman Dewangan: rtc: add RTC driver for TPS6586x . (December 21, 2012)

• Laxman Dewangan: rtc: add RTC driver for TPS80031/TPS80032 . (December 27, 2012)

• Roland Stigge: gpio: Add block GPIO . (December 27, 2012)

• Solomon Peachy: RFC: Driver ST-E cw1200 driver . (January 1, 2013)

• Guennadi Liakhovetski: media: V4L2: support asynchronous subdevice registration . (January 1, 2013)

Documentation

• Michael Kerrisk (man-pages): man-pages-3.45 is released . (December 22, 2012)

Filesystems and block I/O

• zwu.kernel@gmail.com: vfs: hot data tracking . (December 21, 2012)

• Tejun Heo: block: implement blkcg hierarchy support in cfq, take#2 . (December 28, 2012)

• Zheng Liu: ext4: extent status tree implementation (step2) . (January 1, 2013)

Memory management

• Tang Chen: memory-hotplug: hot-remove physical memory . (December 27, 2012)

• Sha Zhengju: Per-cgroup page stat accounting . (December 27, 2012)

• Bartlomiej Zolnierkiewicz: Timeout Based User-space Low Memory Killer Daemon . (January 1, 2013)

• Minchan Kim: Support volatile for anonymous range . (January 3, 2013)

Networking

• Jiri Pirko: net: introduce upper device lists and remove dev->master . (January 2, 2013)

Architecture-specific

• Hiroshi Doyu: ARM: Initial support for Tegra 114 SoC. . (December 21, 2012)

• Fenghua Yu: x86/microcode: Early load microcode . (December 21, 2012)

• Rik van Riel: [RFC PATCH 0/3] x86,smp: make ticket spinlock proportional backoff w/ auto tuning . (December 22, 2012)

Virtualization and containers

• Daniel Kiper: xen: Initial kexec/kdump implementation . (December 27, 2012)

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