Cleancache

Motivation

Cleancache is a new optional feature provided by the VFS layer that potentially dramatically increases page cache effectiveness for many workloads in many environments at a negligible cost.

Cleancache can be thought of as a page-granularity victim cache for clean pages that the kernel’s pageframe replacement algorithm (PFRA) would like to keep around, but can’t since there isn’t enough memory. So when the PFRA “evicts” a page, it first attempts to use cleancache code to put the data contained in that page into “transcendent memory”, memory that is not directly accessible or addressable by the kernel and is of unknown and possibly time-varying size.

Later, when a cleancache-enabled filesystem wishes to access a page in a file on disk, it first checks cleancache to see if it already contains it; if it does, the page of data is copied into the kernel and a disk access is avoided.

Transcendent memory “drivers” for cleancache are currently implemented in Xen (using hypervisor memory) and zcache (using in-kernel compressed memory) and other implementations are in development.

FAQs are included below.

Implementation Overview

A cleancache “backend” that provides transcendent memory registers itself to the kernel’s cleancache “frontend” by calling cleancache_register_ops, passing a pointer to a cleancache_ops structure with funcs set appropriately. The functions provided must conform to certain semantics as follows:

Most important, cleancache is “ephemeral”. Pages which are copied into cleancache have an indefinite lifetime which is completely unknowable by the kernel and so may or may not still be in cleancache at any later time. Thus, as its name implies, cleancache is not suitable for dirty pages. Cleancache has complete discretion over what pages to preserve and what pages to discard and when.

Mounting a cleancache-enabled filesystem should call “init_fs” to obtain a pool id which, if positive, must be saved in the filesystem’s superblock; a negative return value indicates failure. A “put_page” will copy a (presumably about-to-be-evicted) page into cleancache and associate it with the pool id, a file key, and a page index into the file. (The combination of a pool id, a file key, and an index is sometimes called a “handle”.) A “get_page” will copy the page, if found, from cleancache into kernel memory. An “invalidate_page” will ensure the page no longer is present in cleancache; an “invalidate_inode” will invalidate all pages associated with the specified file; and, when a filesystem is unmounted, an “invalidate_fs” will invalidate all pages in all files specified by the given pool id and also surrender the pool id.

An “init_shared_fs”, like init_fs, obtains a pool id but tells cleancache to treat the pool as shared using a 128-bit UUID as a key. On systems that may run multiple kernels (such as hard partitioned or virtualized systems) that may share a clustered filesystem, and where cleancache may be shared among those kernels, calls to init_shared_fs that specify the same UUID will receive the same pool id, thus allowing the pages to be shared. Note that any security requirements must be imposed outside of the kernel (e.g. by “tools” that control cleancache). Or a cleancache implementation can simply disable shared_init by always returning a negative value.

If a get_page is successful on a non-shared pool, the page is invalidated (thus making cleancache an “exclusive” cache). On a shared pool, the page is NOT invalidated on a successful get_page so that it remains accessible to other sharers. The kernel is responsible for ensuring coherency between cleancache (shared or not), the page cache, and the filesystem, using cleancache invalidate operations as required.

Note that cleancache must enforce put-put-get coherency and get-get coherency. For the former, if two puts are made to the same handle but with different data, say AAA by the first put and BBB by the second, a subsequent get can never return the stale data (AAA). For get-get coherency, if a get for a given handle fails, subsequent gets for that handle will never succeed unless preceded by a successful put with that handle.

Last, cleancache provides no SMP serialization guarantees; if two different Linux threads are simultaneously putting and invalidating a page with the same handle, the results are indeterminate. Callers must lock the page to ensure serial behavior.

Cleancache Performance Metrics

If properly configured, monitoring of cleancache is done via debugfs in the /sys/kernel/debug/cleancache directory. The effectiveness of cleancache can be measured (across all filesystems) with:

succ_gets

number of gets that were successful

failed_gets

number of gets that failed

puts

number of puts attempted (all “succeed”)

invalidates

number of invalidates attempted

A backend implementation may provide additional metrics.

FAQ

• Where’s the value? (Andrew Morton)

Cleancache provides a significant performance benefit to many workloads in many environments with negligible overhead by improving the effectiveness of the pagecache. Clean pagecache pages are saved in transcendent memory (RAM that is otherwise not directly addressable to the kernel); fetching those pages later avoids “refaults” and thus disk reads.

Cleancache (and its sister code “frontswap”) provide interfaces for this transcendent memory (aka “tmem”), which conceptually lies between fast kernel-directly-addressable RAM and slower DMA/asynchronous devices. Disallowing direct kernel or userland reads/writes to tmem is ideal when data is transformed to a different form and size (such as with compression) or secretly moved (as might be useful for write- balancing for some RAM-like devices). Evicted page-cache pages (and swap pages) are a great use for this kind of slower-than-RAM-but-much- faster-than-disk transcendent memory, and the cleancache (and frontswap) “page-object-oriented” specification provides a nice way to read and write – and indirectly “name” – the pages.

In the virtual case, the whole point of virtualization is to statistically multiplex physical resources across the varying demands of multiple virtual machines. This is really hard to do with RAM and efforts to do it well with no kernel change have essentially failed (except in some well-publicized special-case workloads). Cleancache – and frontswap – with a fairly small impact on the kernel, provide a huge amount of flexibility for more dynamic, flexible RAM multiplexing. Specifically, the Xen Transcendent Memory backend allows otherwise “fallow” hypervisor-owned RAM to not only be “time-shared” between multiple virtual machines, but the pages can be compressed and deduplicated to optimize RAM utilization. And when guest OS’s are induced to surrender underutilized RAM (e.g. with “self-ballooning”), page cache pages are the first to go, and cleancache allows those pages to be saved and reclaimed if overall host system memory conditions allow.

And the identical interface used for cleancache can be used in physical systems as well. The zcache driver acts as a memory-hungry device that stores pages of data in a compressed state. And the proposed “RAMster” driver shares RAM across multiple physical systems.

• Why does cleancache have its sticky fingers so deep inside the filesystems and VFS? (Andrew Morton and Christoph Hellwig)

The core hooks for cleancache in VFS are in most cases a single line and the minimum set are placed precisely where needed to maintain coherency (via cleancache_invalidate operations) between cleancache, the page cache, and disk. All hooks compile into nothingness if cleancache is config’ed off and turn into a function-pointer- compare-to-NULL if config’ed on but no backend claims the ops functions, or to a compare-struct-element-to-negative if a backend claims the ops functions but a filesystem doesn’t enable cleancache.

Some filesystems are built entirely on top of VFS and the hooks in VFS are sufficient, so don’t require an “init_fs” hook; the initial implementation of cleancache didn’t provide this hook. But for some filesystems (such as btrfs), the VFS hooks are incomplete and one or more hooks in fs-specific code are required. And for some other filesystems, such as tmpfs, cleancache may be counterproductive. So it seemed prudent to require a filesystem to “opt in” to use cleancache, which requires adding a hook in each filesystem. Not all filesystems are supported by cleancache only because they haven’t been tested. The existing set should be sufficient to validate the concept, the opt-in approach means that untested filesystems are not affected, and the hooks in the existing filesystems should make it very easy to add more filesystems in the future.

The total impact of the hooks to existing fs and mm files is only about 40 lines added (not counting comments and blank lines).

• Why not make cleancache asynchronous and batched so it can more easily interface with real devices with DMA instead of copying each individual page? (Minchan Kim)

The one-page-at-a-time copy semantics simplifies the implementation on both the frontend and backend and also allows the backend to do fancy things on-the-fly like page compression and page deduplication. And since the data is “gone” (copied into/out of the pageframe) before the cleancache get/put call returns, a great deal of race conditions and potential coherency issues are avoided. While the interface seems odd for a “real device” or for real kernel-addressable RAM, it makes perfect sense for transcendent memory.

• Why is non-shared cleancache “exclusive”? And where is the page “invalidated” after a “get”? (Minchan Kim)

The main reason is to free up space in transcendent memory and to avoid unnecessary cleancache_invalidate calls. If you want inclusive, the page can be “put” immediately following the “get”. If put-after-get for inclusive becomes common, the interface could be easily extended to add a “get_no_invalidate” call.

The invalidate is done by the cleancache backend implementation.

• What’s the performance impact?

Performance analysis has been presented at OLS’09 and LCA’10. Briefly, performance gains can be significant on most workloads, especially when memory pressure is high (e.g. when RAM is overcommitted in a virtual workload); and because the hooks are invoked primarily in place of or in addition to a disk read/write, overhead is negligible even in worst case workloads. Basically cleancache replaces I/O with memory-copy-CPU-overhead; on older single-core systems with slow memory-copy speeds, cleancache has little value, but in newer multicore machines, especially consolidated/virtualized machines, it has great value.

• How do I add cleancache support for filesystem X? (Boaz Harrash)

Filesystems that are well-behaved and conform to certain restrictions can utilize cleancache simply by making a call to cleancache_init_fs at mount time. Unusual, misbehaving, or poorly layered filesystems must either add additional hooks and/or undergo extensive additional testing… or should just not enable the optional cleancache.

Some points for a filesystem to consider:

• The FS should be block-device-based (e.g. a ram-based FS such as tmpfs should not enable cleancache)

• To ensure coherency/correctness, the FS must ensure that all file removal or truncation operations either go through VFS or add hooks to do the equivalent cleancache “invalidate” operations

• To ensure coherency/correctness, either inode numbers must be unique across the lifetime of the on-disk file OR the FS must provide an “encode_fh” function.

• The FS must call the VFS superblock alloc and deactivate routines or add hooks to do the equivalent cleancache calls done there.

• To maximize performance, all pages fetched from the FS should go through the do_mpag_readpage routine or the FS should add hooks to do the equivalent (cf. btrfs)

• Currently, the FS blocksize must be the same as PAGESIZE. This is not an architectural restriction, but no backends currently support anything different.

• A clustered FS should invoke the “shared_init_fs” cleancache hook to get best performance for some backends.

• Why not use the KVA of the inode as the key? (Christoph Hellwig)

If cleancache would use the inode virtual address instead of inode/filehandle, the pool id could be eliminated. But, this won’t work because cleancache retains pagecache data pages persistently even when the inode has been pruned from the inode unused list, and only invalidates the data page if the file gets removed/truncated. So if cleancache used the inode kva, there would be potential coherency issues if/when the inode kva is reused for a different file. Alternately, if cleancache invalidated the pages when the inode kva was freed, much of the value of cleancache would be lost because the cache of pages in cleanache is potentially much larger than the kernel pagecache and is most useful if the pages survive inode cache removal.

• Why is a global variable required?

The cleancache_enabled flag is checked in all of the frequently-used cleancache hooks. The alternative is a function call to check a static variable. Since cleancache is enabled dynamically at runtime, systems that don’t enable cleancache would suffer thousands (possibly tens-of-thousands) of unnecessary function calls per second. So the global variable allows cleancache to be enabled by default at compile time, but have insignificant performance impact when cleancache remains disabled at runtime.

• Does cleanache work with KVM?

The memory model of KVM is sufficiently different that a cleancache backend may have less value for KVM. This remains to be tested, especially in an overcommitted system.

• Does cleancache work in userspace? It sounds useful for memory hungry caches like web browsers. (Jamie Lokier)

No plans yet, though we agree it sounds useful, at least for apps that bypass the page cache (e.g. O_DIRECT).

Last updated: Dan Magenheimer, April 13 2011