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I/O Schedulers

Simply sending out requests to the block devices in the order that the kernel issues them, as soon as it issues them, results in awful performance. One of the slowest operations in a modern computer is disk seeks. Each seekpositioning the hard disk's head at the location of a specific blocktakes many milliseconds. Minimizing seeks is absolutely crucial to the system's performance.

Therefore, the kernel does not issue block I/O requests to the disk in the order they are received or as soon as they are received. Instead, it performs operations called merging and sorting to greatly improve the performance of the system as a whole[2]. The subsystem of the kernel that performs these operations is called the I/O scheduler.

[2] This point must be stressed. A system without these features, or wherein these features are poorly implemented, would perform horribly with even a modest number of block I/O operations.

The I/O scheduler divides the resource of disk I/O among the pending block I/O requests in the system. It does this through the merging and sorting of pending requests in the request queue. The I/O scheduler is not to be confused with the process scheduler (see Chapter 4, "Process Scheduling"), which divides the resource of the processor among the processes on the system. The two subsystems are similar in nature but not the same. Both the process scheduler and the I/O scheduler virtualize a resource among multiple objects. In the case of the process scheduler, the processor is virtualized and shared among the processes on the system. This provides the illusion of virtualization inherent in a multitasking and timesharing operating system, such as any Unix. On the other hand, the I/O scheduler virtualizes block devices among multiple outstanding block requests. This is done to minimize disk seeks and ensure optimum disk performance.

The Job of an I/O Scheduler

An I/O scheduler works by managing a block device's request queue. It decides the order of requests in the queue and at what time each request is dispatched to the block device. It manages the request queue with the goal of reducing seeks, which results in greater global throughput. The "global" modifier here is important. An I/O scheduler, very openly, is unfair to some requests at the expense of improving the overall performance of the system.

I/O schedulers perform two primary actions to minimize seeks: merging and sorting. Merging is the coalescing of two or more requests into one. Consider an example request that is submitted to the queue by a filesystemsay, to read a chunk of data from a file. (At this point, of course, everything is occurring in terms of sectors and blocks and not files, but presume that the requested blocks originate from a chunk of a file.) If a request is already in the queue to read from an adjacent sector on the disk (for example, an earlier chunk of the same file), the two requests can be merged into a single request operating on one or more adjacent on-disk sectors. By merging requests, the I/O scheduler reduces the overhead of multiple requests down to a single request. More importantly, only a single command needs to be issued to the disk and servicing the multiple requests can be done without seeking. Consequently, merging requests reduces overhead and minimizes seeks.

Now, assume your fictional read request is submitted to the request queue, but there is no read request to an adjacent sector. You are therefore unable to merge this request with any other request. Now, you could simply stick this request onto the tail of the queue. But, what if there are other requests to a similar location on the disk? Would it not make sense to insert this new request into the queue at a spot near other requests operating on physically near sectors? In fact, I/O schedulers do exactly this. The entire request queue is kept sorted, sector-wise, so that all seeking activity along the queue moves (as much as possible) sequentially over the sectors of the hard disk. The goal is not just to minimize each individual seek, but to minimize all seeking by keeping the disk head moving in a straight line. This is similar to the algorithm employed in elevatorselevators do not jump all over, wildly, from floor to floor. Instead, they try to move gracefully in a single direction. When the final floor is reached in one direction, the elevator can reverse course and move in the other direction. Because of this similarity, I/O schedulers (or sometimes just their sorting algorithm) are called elevators.

The Linus Elevator

Now let's look at some real-life I/O schedulers. The first I/O scheduler is called the Linus Elevator (yes, Linus has an elevator named after him!). It was the default I/O scheduler in 2.4. In 2.6, it was replaced by the following I/O schedulers that we will look athowever, because this elevator is simpler than the subsequent ones, while performing many of the same functions, it still deserves discussion.

The Linus Elevator performs both merging and sorting. When a request is added to the queue, it is first checked against every other pending request to see whether it is a possible candidate for merging. The Linus Elevator I/O scheduler performs both front and back merging. The type of merging performed depends on the location of the existing adjacent request. If the new request immediately proceeds an existing request, it is front merged. Conversely, if the new request immediately precedes an existing request, it is back merged. Because of the way files are laid out (usually by increasing sector number) and the I/O operations performed in a typical workload (data is normally read from start to finish and not in reverse), front merging is very rare compared to back merging. Nonetheless, the Linus Elevator checks for and performs both types of merge.

If the merge attempt fails, a possible insertion point in the queue (a location in the queue where the new request fits sector-wise between the existing requests) is then sought. If one is found, the new request is inserted there. If a suitable location is not found, the request is added to the tail of the queue. Additionally, if an existing request is found in the queue that is older than a predefined threshold, the new request is added to the tail of the queue even if it can be insertion sorted elsewhere. This prevents many requests to nearby on-disk locations from indefinitely starving requests to other locations on the disk. Unfortunately, this "age" check is not very efficient. It does not provide any real attempt to service requests in a given timeframeit merely stops insertion-sorting requests after a suitable delay. This improves latency but can still lead to request starvation, which was the big must-fix of the 2.4 I/O scheduler.

In summary, when a request is added to the queue, four operations are possible. In order, they are

1.
First, if a request to an adjacent on-disk sector is in the queue, the existing request and the new request are merged into a single request.

2.
Second, if a request in the queue is sufficiently old, the new request is inserted at the tail of the queue to prevent starvation of the other, older, requests.

3.
Next, if there is a suitable location sector-wise in the queue, the new request is inserted there. This keeps the queue sorted by physical location on disk.

4.
Finally, if no such suitable insertion point exists, the request is inserted at the tail of the queue

The Deadline I/O Scheduler

The Deadline I/O scheduler sought to prevent the starvation caused by the Linus Elevator. In the interest of minimizing seeks, heavy disk I/O operations to one area of the disk can indefinitely starve request operations to another part of the disk. Indeed, a stream of requests to the same area of the disk can result in other far-off requests never being serviced. This starvation is unfair.

Worse, the general issue of request starvation introduces a specific instance of the problem known as writes starving reads. Write operations can usually be committed to disk whenever the kernel gets around to them, entirely asynchronous with respect to the submitting application. Read operations are quite different. Normally, when an application submits a read request, the application blocks until the request is fulfilled. That is, read requests occur synchronously with respect to the submitting application. Although system response is largely unaffected by write latency (the time required to commit a write request), read latency (the time required to commit a read request) is very important. Write latency has little bearing on application performance[3], but an application must wait, twiddling its thumbs, for the completion of each read request. Consequently, read latency is very important to the performance of the system.

[3] We still do not want to delay write requests indefinitely, however, because the kernel wants to ensure that data is eventually written to disk to prevent in-memory buffers from growing too large or too old.

Compounding the problem, read requests tend to be dependent on each other. For example, consider the reading of a large number of files. Each read occurs in small buffered chunks. The application does not start reading the next chunk (or the next file, for that matter) until the previous chunk is read from disk and returned to the application. Worse, both read and write operations require the reading of various metadata, such as inodes. Reading these blocks off the disk further serializes I/O. Consequently, if each read request is individually starved, the total delay to such applications compounds and can grow enormous. Recognizing that the asynchrony and interdependency of read requests results in a much stronger bearing of read latency on the performance of the system, the Deadline I/O scheduler implements several features to ensure that request starvation in general, and read starvation in specific, is minimized.

Note that reducing request starvation comes at a cost to global throughput. Even the Linus Elevator makes this compromise, albeit in a much milder mannerthe Linus Elevator could provide better overall throughput (via a greater minimization of seeks) if it always inserted requests into the queue sector-wise and never checked for old requests and reverted to insertion at the tail of the queue. Although minimizing seeks is very important, indefinite starvation is not good either. The Deadline I/O scheduler, therefore, works harder to limit starvation while still providing good global throughput. Make no mistake: It is a tough act to provide request fairness, yet maximize global throughput.

In the Deadline I/O scheduler, each request is associated with an expiration time. By default, the expiration time is 500 milliseconds in the future for read requests and 5 seconds in the future for write requests. The Deadline I/O scheduler operates similarly to the Linus Elevator in that it maintains a request queue sorted by physical location on disk. It calls this queue the sorted queue. When a new request is submitted to the sorted queue, the Deadline I/O scheduler performs merging and insertion like the Linus Elevator[4]. The Deadline I/O scheduler also, however, inserts the request into a second queue that depends on the type of request. Read requests are sorted into a special read FIFO queue and write requests are inserted into a special write FIFO queue. Although the normal queue is sorted by on-disk sector, these queues are kept FIFO (effectively, they are sorted by time). Consequently, new requests are always added to the tail of the queue. Under normal operation, the Deadline I/O scheduler pulls requests from the head of the sorted queue into the dispatch queue. The dispatch queue is then fed to the disk drive. This results in minimal seeks.

[4] Performing front merging is optional in the Deadline I/O scheduler, however. It is not always worth the trouble because many workloads have very few requests that can be front merged.

If the request at the head of either the write FIFO queue or the read FIFO queue expires (that is, if the current time becomes greater than the expiration time associated with the request), the Deadline I/O scheduler then begins servicing requests from the FIFO queue. In this manner, the Deadline I/O scheduler attempts to ensure that no request is outstanding longer than its expiration time. See Figure 13.3.

Figure 13.3. The three queues of the Deadline I/O scheduler.


Note that the Deadline I/O scheduler does not make any strict guarantees over request latency. It is capable, however, of generally committing requests on or before their expiration. This prevents request starvation. Because read requests are given a substantially smaller expiration value than write requests, the Deadline I/O scheduler also works to ensure that write requests do not starve read requests. This preference toward read requests provides minimized read latency.

The Deadline I/O scheduler lives in drivers/block/deadline-iosched.c.

The Anticipatory I/O Scheduler

Although the Deadline I/O scheduler does a great job minimizing read latency, it does so at the expense of global throughput. Consider a system undergoing heavy write activity. Every time a read request is submitted, the I/O scheduler quickly rushes to handle the read request. This results in the disk seeking over to where the read is, performing the read operation, and then seeking back to continue the ongoing write operation, repeating this little charade for each read request. The preference toward read requests is a good thing, but the resulting pair of seeks (one to the location of the read request and another back to the ongoing write) is detrimental to global disk throughput. The Anticipatory I/O scheduler aims to continue to provide excellent read latency, but also provide excellent global throughput.

First, the Anticipatory I/O scheduler starts with the Deadline I/O scheduler as its base. Therefore, it is not entirely different. The Anticipatory I/O scheduler implements three queues (plus the dispatch queue) and expirations for each request, just like the Deadline I/O scheduler. The major change is the addition of an anticipation heuristic.

The Anticipatory I/O scheduler attempts to minimize the seek storm that accompanies read requests issued during other disk I/O activity. When a read request is issued, it is handled as usual, within its usual expiration period. After the request is submitted, however, the Anticipatory I/O scheduler does not immediately seek back and return to handling other requests. Instead, it does absolutely nothing for a few milliseconds (the actual value is configurable; by default it is six milliseconds). In those few milliseconds, there is a good chance that the application will submit another read request. Any requests issued to an adjacent area of the disk are immediately handled. After the waiting period elapses, the Anticipatory I/O scheduler seeks back to where it left off and continues handling the previous requests.

It is important to note that the few milliseconds that are spent in anticipation for more requests are well worth it if they minimize even a modest percentage of the back-and-forth seeking that results from the servicing of read requests during other heavy requests. If an adjacent I/O request is issued within the waiting period, the I/O scheduler just saved a pair of seeks. As more and more reads are issued to the same area of disk, many more seeks are prevented.

Of course, if no activity occurs within the waiting period, the Anticipatory I/O scheduler loses and a few milliseconds are wasted. The key to reaping maximum benefit from the Anticipatory I/O scheduler is correctly anticipating the actions of applications and filesystems. This is done via a set of statistics and associated heuristics. The Anticipatory I/O scheduler keeps track of per-process statistics pertaining to block I/O habits in hopes of correctly anticipating the actions of applications. With a sufficiently high percentage of correct anticipations, the Anticipatory I/O scheduler can greatly reduce the penalty of seeking to service read requests, while still providing the attention to such requests that system response requires. This allows the Anticipatory I/O scheduler to minimize read latency, while also minimizing the number and duration of seeks. This results in low system latency and high system throughput.

The Anticipatory I/O scheduler lives in the file drivers/block/as-iosched.c in the kernel source tree. It is the default I/O scheduler in the Linux kernel. It performs well across most workloads. It is ideal for servers, although it performs very poorly on certain uncommon but critical workloads involving seek-happy databases.

The Complete Fair Queuing I/O Scheduler

The Complete Fair Queuing (CFQ) I/O scheduler is an I/O scheduler designed for specialized workloads, but that in practice actually provides good performance across multiple workloads. It is fundamentally different from the previous I/O schedulers that have been covered, however.

The CFQ I/O scheduler assigns incoming I/O requests to specific queues based on the process originating the I/O request. For example, I/O requests from process foo go in foo's queues, and I/O requests from process bar go in bar's queue. Within each queue, requests are coalesced with adjacent requests and insertion sorted. The queues are thus kept sorted sector-wise, as with the other I/O scheduler's queues. The difference with the CFQ I/O scheduler is that there is one queue for each process submitting I/O.

The CFQ I/O scheduler then services the queues round robin, plucking a configurable number of requests (by default, four) from each queue before continuing on to the next. This provides fairness at a per-process level, assuring that each process receives a fair slice of the disk's bandwidth. The intended workload is multimedia, in which such a fair algorithm can guarantee that, for example, an audio player is able to always refill its audio buffers from disk in time. In practice, however, the CFQ I/O scheduler performs well in many scenarios.

The Complete Fair Queuing I/O scheduler lives in drivers/block/cfq-iosched.c. It is recommended for desktop workloads, although it performs reasonably well in nearly all workloads without any pathological corner cases.

The Noop I/O Scheduler

A fourth and final I/O scheduler is the Noop I/O scheduler, so named because it is basically a noopit does not do much. The Noop I/O scheduler does not perform sorting or any other form of seek-prevention whatsoever. In turn, it has no need to implement anything akin to the slick algorithms to minimize request latency that you saw in the previous three I/O schedulers.

The Noop I/O scheduler does perform merging, however, as its lone chore. When a new request is submitted to the queue, it is coalesced with any adjacent requests. Other than this operation, the Noop I/O Scheduler truly is a noop, merely maintaining the request queue in near-FIFO order, from which the block device driver can pluck requests.

The Noop I/O scheduler is not the lazy and worthless I/O scheduler of the bunch; its lack of hard work is with reason. It is intended for block devices that are truly random-access, such as flash memory cards. If a block device has little or no overhead associated with "seeking," then there is no need for insertion sorting of incoming requests, and the Noop I/O scheduler is the ideal candidate.

The Noop I/O scheduler lives in drivers/block/noop-iosched.c. It is intended only for random-access devices.

I/O Scheduler Selection

You have now seen four different I/O schedulers in the 2.6 kernel. Each of these I/O schedulers can be enabled and built into the kernel. By default, block devices use the Anticipatory I/O scheduler. This can be overridden via the boot-time option elevator=foo on the kernel command line, where foo is a valid and enabled I/O Scheduler. See Table 13.2.

Table 13.2. Parameters Given to the elevator Option

Parameter

I/O Scheduler

as

Anticipatory

cfq

Complete Fair Queuing

deadline

Deadline

noop

Noop


For example, the kernel command line option elevator=cfq would enable use of the Complete Fair Queuing I/O scheduler for all block devices.

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