Root/
1 | ============================ |
2 | LINUX KERNEL MEMORY BARRIERS |
3 | ============================ |
4 | |
5 | By: David Howells <dhowells@redhat.com> |
6 | Paul E. McKenney <paulmck@linux.vnet.ibm.com> |
7 | |
8 | Contents: |
9 | |
10 | (*) Abstract memory access model. |
11 | |
12 | - Device operations. |
13 | - Guarantees. |
14 | |
15 | (*) What are memory barriers? |
16 | |
17 | - Varieties of memory barrier. |
18 | - What may not be assumed about memory barriers? |
19 | - Data dependency barriers. |
20 | - Control dependencies. |
21 | - SMP barrier pairing. |
22 | - Examples of memory barrier sequences. |
23 | - Read memory barriers vs load speculation. |
24 | - Transitivity |
25 | |
26 | (*) Explicit kernel barriers. |
27 | |
28 | - Compiler barrier. |
29 | - CPU memory barriers. |
30 | - MMIO write barrier. |
31 | |
32 | (*) Implicit kernel memory barriers. |
33 | |
34 | - Locking functions. |
35 | - Interrupt disabling functions. |
36 | - Sleep and wake-up functions. |
37 | - Miscellaneous functions. |
38 | |
39 | (*) Inter-CPU locking barrier effects. |
40 | |
41 | - Locks vs memory accesses. |
42 | - Locks vs I/O accesses. |
43 | |
44 | (*) Where are memory barriers needed? |
45 | |
46 | - Interprocessor interaction. |
47 | - Atomic operations. |
48 | - Accessing devices. |
49 | - Interrupts. |
50 | |
51 | (*) Kernel I/O barrier effects. |
52 | |
53 | (*) Assumed minimum execution ordering model. |
54 | |
55 | (*) The effects of the cpu cache. |
56 | |
57 | - Cache coherency. |
58 | - Cache coherency vs DMA. |
59 | - Cache coherency vs MMIO. |
60 | |
61 | (*) The things CPUs get up to. |
62 | |
63 | - And then there's the Alpha. |
64 | |
65 | (*) Example uses. |
66 | |
67 | - Circular buffers. |
68 | |
69 | (*) References. |
70 | |
71 | |
72 | ============================ |
73 | ABSTRACT MEMORY ACCESS MODEL |
74 | ============================ |
75 | |
76 | Consider the following abstract model of the system: |
77 | |
78 | : : |
79 | : : |
80 | : : |
81 | +-------+ : +--------+ : +-------+ |
82 | | | : | | : | | |
83 | | | : | | : | | |
84 | | CPU 1 |<----->| Memory |<----->| CPU 2 | |
85 | | | : | | : | | |
86 | | | : | | : | | |
87 | +-------+ : +--------+ : +-------+ |
88 | ^ : ^ : ^ |
89 | | : | : | |
90 | | : | : | |
91 | | : v : | |
92 | | : +--------+ : | |
93 | | : | | : | |
94 | | : | | : | |
95 | +---------->| Device |<----------+ |
96 | : | | : |
97 | : | | : |
98 | : +--------+ : |
99 | : : |
100 | |
101 | Each CPU executes a program that generates memory access operations. In the |
102 | abstract CPU, memory operation ordering is very relaxed, and a CPU may actually |
103 | perform the memory operations in any order it likes, provided program causality |
104 | appears to be maintained. Similarly, the compiler may also arrange the |
105 | instructions it emits in any order it likes, provided it doesn't affect the |
106 | apparent operation of the program. |
107 | |
108 | So in the above diagram, the effects of the memory operations performed by a |
109 | CPU are perceived by the rest of the system as the operations cross the |
110 | interface between the CPU and rest of the system (the dotted lines). |
111 | |
112 | |
113 | For example, consider the following sequence of events: |
114 | |
115 | CPU 1 CPU 2 |
116 | =============== =============== |
117 | { A == 1; B == 2 } |
118 | A = 3; x = A; |
119 | B = 4; y = B; |
120 | |
121 | The set of accesses as seen by the memory system in the middle can be arranged |
122 | in 24 different combinations: |
123 | |
124 | STORE A=3, STORE B=4, x=LOAD A->3, y=LOAD B->4 |
125 | STORE A=3, STORE B=4, y=LOAD B->4, x=LOAD A->3 |
126 | STORE A=3, x=LOAD A->3, STORE B=4, y=LOAD B->4 |
127 | STORE A=3, x=LOAD A->3, y=LOAD B->2, STORE B=4 |
128 | STORE A=3, y=LOAD B->2, STORE B=4, x=LOAD A->3 |
129 | STORE A=3, y=LOAD B->2, x=LOAD A->3, STORE B=4 |
130 | STORE B=4, STORE A=3, x=LOAD A->3, y=LOAD B->4 |
131 | STORE B=4, ... |
132 | ... |
133 | |
134 | and can thus result in four different combinations of values: |
135 | |
136 | x == 1, y == 2 |
137 | x == 1, y == 4 |
138 | x == 3, y == 2 |
139 | x == 3, y == 4 |
140 | |
141 | |
142 | Furthermore, the stores committed by a CPU to the memory system may not be |
143 | perceived by the loads made by another CPU in the same order as the stores were |
144 | committed. |
145 | |
146 | |
147 | As a further example, consider this sequence of events: |
148 | |
149 | CPU 1 CPU 2 |
150 | =============== =============== |
151 | { A == 1, B == 2, C = 3, P == &A, Q == &C } |
152 | B = 4; Q = P; |
153 | P = &B D = *Q; |
154 | |
155 | There is an obvious data dependency here, as the value loaded into D depends on |
156 | the address retrieved from P by CPU 2. At the end of the sequence, any of the |
157 | following results are possible: |
158 | |
159 | (Q == &A) and (D == 1) |
160 | (Q == &B) and (D == 2) |
161 | (Q == &B) and (D == 4) |
162 | |
163 | Note that CPU 2 will never try and load C into D because the CPU will load P |
164 | into Q before issuing the load of *Q. |
165 | |
166 | |
167 | DEVICE OPERATIONS |
168 | ----------------- |
169 | |
170 | Some devices present their control interfaces as collections of memory |
171 | locations, but the order in which the control registers are accessed is very |
172 | important. For instance, imagine an ethernet card with a set of internal |
173 | registers that are accessed through an address port register (A) and a data |
174 | port register (D). To read internal register 5, the following code might then |
175 | be used: |
176 | |
177 | *A = 5; |
178 | x = *D; |
179 | |
180 | but this might show up as either of the following two sequences: |
181 | |
182 | STORE *A = 5, x = LOAD *D |
183 | x = LOAD *D, STORE *A = 5 |
184 | |
185 | the second of which will almost certainly result in a malfunction, since it set |
186 | the address _after_ attempting to read the register. |
187 | |
188 | |
189 | GUARANTEES |
190 | ---------- |
191 | |
192 | There are some minimal guarantees that may be expected of a CPU: |
193 | |
194 | (*) On any given CPU, dependent memory accesses will be issued in order, with |
195 | respect to itself. This means that for: |
196 | |
197 | Q = P; D = *Q; |
198 | |
199 | the CPU will issue the following memory operations: |
200 | |
201 | Q = LOAD P, D = LOAD *Q |
202 | |
203 | and always in that order. |
204 | |
205 | (*) Overlapping loads and stores within a particular CPU will appear to be |
206 | ordered within that CPU. This means that for: |
207 | |
208 | a = *X; *X = b; |
209 | |
210 | the CPU will only issue the following sequence of memory operations: |
211 | |
212 | a = LOAD *X, STORE *X = b |
213 | |
214 | And for: |
215 | |
216 | *X = c; d = *X; |
217 | |
218 | the CPU will only issue: |
219 | |
220 | STORE *X = c, d = LOAD *X |
221 | |
222 | (Loads and stores overlap if they are targeted at overlapping pieces of |
223 | memory). |
224 | |
225 | And there are a number of things that _must_ or _must_not_ be assumed: |
226 | |
227 | (*) It _must_not_ be assumed that independent loads and stores will be issued |
228 | in the order given. This means that for: |
229 | |
230 | X = *A; Y = *B; *D = Z; |
231 | |
232 | we may get any of the following sequences: |
233 | |
234 | X = LOAD *A, Y = LOAD *B, STORE *D = Z |
235 | X = LOAD *A, STORE *D = Z, Y = LOAD *B |
236 | Y = LOAD *B, X = LOAD *A, STORE *D = Z |
237 | Y = LOAD *B, STORE *D = Z, X = LOAD *A |
238 | STORE *D = Z, X = LOAD *A, Y = LOAD *B |
239 | STORE *D = Z, Y = LOAD *B, X = LOAD *A |
240 | |
241 | (*) It _must_ be assumed that overlapping memory accesses may be merged or |
242 | discarded. This means that for: |
243 | |
244 | X = *A; Y = *(A + 4); |
245 | |
246 | we may get any one of the following sequences: |
247 | |
248 | X = LOAD *A; Y = LOAD *(A + 4); |
249 | Y = LOAD *(A + 4); X = LOAD *A; |
250 | {X, Y} = LOAD {*A, *(A + 4) }; |
251 | |
252 | And for: |
253 | |
254 | *A = X; Y = *A; |
255 | |
256 | we may get either of: |
257 | |
258 | STORE *A = X; Y = LOAD *A; |
259 | STORE *A = Y = X; |
260 | |
261 | |
262 | ========================= |
263 | WHAT ARE MEMORY BARRIERS? |
264 | ========================= |
265 | |
266 | As can be seen above, independent memory operations are effectively performed |
267 | in random order, but this can be a problem for CPU-CPU interaction and for I/O. |
268 | What is required is some way of intervening to instruct the compiler and the |
269 | CPU to restrict the order. |
270 | |
271 | Memory barriers are such interventions. They impose a perceived partial |
272 | ordering over the memory operations on either side of the barrier. |
273 | |
274 | Such enforcement is important because the CPUs and other devices in a system |
275 | can use a variety of tricks to improve performance, including reordering, |
276 | deferral and combination of memory operations; speculative loads; speculative |
277 | branch prediction and various types of caching. Memory barriers are used to |
278 | override or suppress these tricks, allowing the code to sanely control the |
279 | interaction of multiple CPUs and/or devices. |
280 | |
281 | |
282 | VARIETIES OF MEMORY BARRIER |
283 | --------------------------- |
284 | |
285 | Memory barriers come in four basic varieties: |
286 | |
287 | (1) Write (or store) memory barriers. |
288 | |
289 | A write memory barrier gives a guarantee that all the STORE operations |
290 | specified before the barrier will appear to happen before all the STORE |
291 | operations specified after the barrier with respect to the other |
292 | components of the system. |
293 | |
294 | A write barrier is a partial ordering on stores only; it is not required |
295 | to have any effect on loads. |
296 | |
297 | A CPU can be viewed as committing a sequence of store operations to the |
298 | memory system as time progresses. All stores before a write barrier will |
299 | occur in the sequence _before_ all the stores after the write barrier. |
300 | |
301 | [!] Note that write barriers should normally be paired with read or data |
302 | dependency barriers; see the "SMP barrier pairing" subsection. |
303 | |
304 | |
305 | (2) Data dependency barriers. |
306 | |
307 | A data dependency barrier is a weaker form of read barrier. In the case |
308 | where two loads are performed such that the second depends on the result |
309 | of the first (eg: the first load retrieves the address to which the second |
310 | load will be directed), a data dependency barrier would be required to |
311 | make sure that the target of the second load is updated before the address |
312 | obtained by the first load is accessed. |
313 | |
314 | A data dependency barrier is a partial ordering on interdependent loads |
315 | only; it is not required to have any effect on stores, independent loads |
316 | or overlapping loads. |
317 | |
318 | As mentioned in (1), the other CPUs in the system can be viewed as |
319 | committing sequences of stores to the memory system that the CPU being |
320 | considered can then perceive. A data dependency barrier issued by the CPU |
321 | under consideration guarantees that for any load preceding it, if that |
322 | load touches one of a sequence of stores from another CPU, then by the |
323 | time the barrier completes, the effects of all the stores prior to that |
324 | touched by the load will be perceptible to any loads issued after the data |
325 | dependency barrier. |
326 | |
327 | See the "Examples of memory barrier sequences" subsection for diagrams |
328 | showing the ordering constraints. |
329 | |
330 | [!] Note that the first load really has to have a _data_ dependency and |
331 | not a control dependency. If the address for the second load is dependent |
332 | on the first load, but the dependency is through a conditional rather than |
333 | actually loading the address itself, then it's a _control_ dependency and |
334 | a full read barrier or better is required. See the "Control dependencies" |
335 | subsection for more information. |
336 | |
337 | [!] Note that data dependency barriers should normally be paired with |
338 | write barriers; see the "SMP barrier pairing" subsection. |
339 | |
340 | |
341 | (3) Read (or load) memory barriers. |
342 | |
343 | A read barrier is a data dependency barrier plus a guarantee that all the |
344 | LOAD operations specified before the barrier will appear to happen before |
345 | all the LOAD operations specified after the barrier with respect to the |
346 | other components of the system. |
347 | |
348 | A read barrier is a partial ordering on loads only; it is not required to |
349 | have any effect on stores. |
350 | |
351 | Read memory barriers imply data dependency barriers, and so can substitute |
352 | for them. |
353 | |
354 | [!] Note that read barriers should normally be paired with write barriers; |
355 | see the "SMP barrier pairing" subsection. |
356 | |
357 | |
358 | (4) General memory barriers. |
359 | |
360 | A general memory barrier gives a guarantee that all the LOAD and STORE |
361 | operations specified before the barrier will appear to happen before all |
362 | the LOAD and STORE operations specified after the barrier with respect to |
363 | the other components of the system. |
364 | |
365 | A general memory barrier is a partial ordering over both loads and stores. |
366 | |
367 | General memory barriers imply both read and write memory barriers, and so |
368 | can substitute for either. |
369 | |
370 | |
371 | And a couple of implicit varieties: |
372 | |
373 | (5) LOCK operations. |
374 | |
375 | This acts as a one-way permeable barrier. It guarantees that all memory |
376 | operations after the LOCK operation will appear to happen after the LOCK |
377 | operation with respect to the other components of the system. |
378 | |
379 | Memory operations that occur before a LOCK operation may appear to happen |
380 | after it completes. |
381 | |
382 | A LOCK operation should almost always be paired with an UNLOCK operation. |
383 | |
384 | |
385 | (6) UNLOCK operations. |
386 | |
387 | This also acts as a one-way permeable barrier. It guarantees that all |
388 | memory operations before the UNLOCK operation will appear to happen before |
389 | the UNLOCK operation with respect to the other components of the system. |
390 | |
391 | Memory operations that occur after an UNLOCK operation may appear to |
392 | happen before it completes. |
393 | |
394 | LOCK and UNLOCK operations are guaranteed to appear with respect to each |
395 | other strictly in the order specified. |
396 | |
397 | The use of LOCK and UNLOCK operations generally precludes the need for |
398 | other sorts of memory barrier (but note the exceptions mentioned in the |
399 | subsection "MMIO write barrier"). |
400 | |
401 | |
402 | Memory barriers are only required where there's a possibility of interaction |
403 | between two CPUs or between a CPU and a device. If it can be guaranteed that |
404 | there won't be any such interaction in any particular piece of code, then |
405 | memory barriers are unnecessary in that piece of code. |
406 | |
407 | |
408 | Note that these are the _minimum_ guarantees. Different architectures may give |
409 | more substantial guarantees, but they may _not_ be relied upon outside of arch |
410 | specific code. |
411 | |
412 | |
413 | WHAT MAY NOT BE ASSUMED ABOUT MEMORY BARRIERS? |
414 | ---------------------------------------------- |
415 | |
416 | There are certain things that the Linux kernel memory barriers do not guarantee: |
417 | |
418 | (*) There is no guarantee that any of the memory accesses specified before a |
419 | memory barrier will be _complete_ by the completion of a memory barrier |
420 | instruction; the barrier can be considered to draw a line in that CPU's |
421 | access queue that accesses of the appropriate type may not cross. |
422 | |
423 | (*) There is no guarantee that issuing a memory barrier on one CPU will have |
424 | any direct effect on another CPU or any other hardware in the system. The |
425 | indirect effect will be the order in which the second CPU sees the effects |
426 | of the first CPU's accesses occur, but see the next point: |
427 | |
428 | (*) There is no guarantee that a CPU will see the correct order of effects |
429 | from a second CPU's accesses, even _if_ the second CPU uses a memory |
430 | barrier, unless the first CPU _also_ uses a matching memory barrier (see |
431 | the subsection on "SMP Barrier Pairing"). |
432 | |
433 | (*) There is no guarantee that some intervening piece of off-the-CPU |
434 | hardware[*] will not reorder the memory accesses. CPU cache coherency |
435 | mechanisms should propagate the indirect effects of a memory barrier |
436 | between CPUs, but might not do so in order. |
437 | |
438 | [*] For information on bus mastering DMA and coherency please read: |
439 | |
440 | Documentation/PCI/pci.txt |
441 | Documentation/PCI/PCI-DMA-mapping.txt |
442 | Documentation/DMA-API.txt |
443 | |
444 | |
445 | DATA DEPENDENCY BARRIERS |
446 | ------------------------ |
447 | |
448 | The usage requirements of data dependency barriers are a little subtle, and |
449 | it's not always obvious that they're needed. To illustrate, consider the |
450 | following sequence of events: |
451 | |
452 | CPU 1 CPU 2 |
453 | =============== =============== |
454 | { A == 1, B == 2, C = 3, P == &A, Q == &C } |
455 | B = 4; |
456 | <write barrier> |
457 | P = &B |
458 | Q = P; |
459 | D = *Q; |
460 | |
461 | There's a clear data dependency here, and it would seem that by the end of the |
462 | sequence, Q must be either &A or &B, and that: |
463 | |
464 | (Q == &A) implies (D == 1) |
465 | (Q == &B) implies (D == 4) |
466 | |
467 | But! CPU 2's perception of P may be updated _before_ its perception of B, thus |
468 | leading to the following situation: |
469 | |
470 | (Q == &B) and (D == 2) ???? |
471 | |
472 | Whilst this may seem like a failure of coherency or causality maintenance, it |
473 | isn't, and this behaviour can be observed on certain real CPUs (such as the DEC |
474 | Alpha). |
475 | |
476 | To deal with this, a data dependency barrier or better must be inserted |
477 | between the address load and the data load: |
478 | |
479 | CPU 1 CPU 2 |
480 | =============== =============== |
481 | { A == 1, B == 2, C = 3, P == &A, Q == &C } |
482 | B = 4; |
483 | <write barrier> |
484 | P = &B |
485 | Q = P; |
486 | <data dependency barrier> |
487 | D = *Q; |
488 | |
489 | This enforces the occurrence of one of the two implications, and prevents the |
490 | third possibility from arising. |
491 | |
492 | [!] Note that this extremely counterintuitive situation arises most easily on |
493 | machines with split caches, so that, for example, one cache bank processes |
494 | even-numbered cache lines and the other bank processes odd-numbered cache |
495 | lines. The pointer P might be stored in an odd-numbered cache line, and the |
496 | variable B might be stored in an even-numbered cache line. Then, if the |
497 | even-numbered bank of the reading CPU's cache is extremely busy while the |
498 | odd-numbered bank is idle, one can see the new value of the pointer P (&B), |
499 | but the old value of the variable B (2). |
500 | |
501 | |
502 | Another example of where data dependency barriers might by required is where a |
503 | number is read from memory and then used to calculate the index for an array |
504 | access: |
505 | |
506 | CPU 1 CPU 2 |
507 | =============== =============== |
508 | { M[0] == 1, M[1] == 2, M[3] = 3, P == 0, Q == 3 } |
509 | M[1] = 4; |
510 | <write barrier> |
511 | P = 1 |
512 | Q = P; |
513 | <data dependency barrier> |
514 | D = M[Q]; |
515 | |
516 | |
517 | The data dependency barrier is very important to the RCU system, for example. |
518 | See rcu_dereference() in include/linux/rcupdate.h. This permits the current |
519 | target of an RCU'd pointer to be replaced with a new modified target, without |
520 | the replacement target appearing to be incompletely initialised. |
521 | |
522 | See also the subsection on "Cache Coherency" for a more thorough example. |
523 | |
524 | |
525 | CONTROL DEPENDENCIES |
526 | -------------------- |
527 | |
528 | A control dependency requires a full read memory barrier, not simply a data |
529 | dependency barrier to make it work correctly. Consider the following bit of |
530 | code: |
531 | |
532 | q = &a; |
533 | if (p) |
534 | q = &b; |
535 | <data dependency barrier> |
536 | x = *q; |
537 | |
538 | This will not have the desired effect because there is no actual data |
539 | dependency, but rather a control dependency that the CPU may short-circuit by |
540 | attempting to predict the outcome in advance. In such a case what's actually |
541 | required is: |
542 | |
543 | q = &a; |
544 | if (p) |
545 | q = &b; |
546 | <read barrier> |
547 | x = *q; |
548 | |
549 | |
550 | SMP BARRIER PAIRING |
551 | ------------------- |
552 | |
553 | When dealing with CPU-CPU interactions, certain types of memory barrier should |
554 | always be paired. A lack of appropriate pairing is almost certainly an error. |
555 | |
556 | A write barrier should always be paired with a data dependency barrier or read |
557 | barrier, though a general barrier would also be viable. Similarly a read |
558 | barrier or a data dependency barrier should always be paired with at least an |
559 | write barrier, though, again, a general barrier is viable: |
560 | |
561 | CPU 1 CPU 2 |
562 | =============== =============== |
563 | a = 1; |
564 | <write barrier> |
565 | b = 2; x = b; |
566 | <read barrier> |
567 | y = a; |
568 | |
569 | Or: |
570 | |
571 | CPU 1 CPU 2 |
572 | =============== =============================== |
573 | a = 1; |
574 | <write barrier> |
575 | b = &a; x = b; |
576 | <data dependency barrier> |
577 | y = *x; |
578 | |
579 | Basically, the read barrier always has to be there, even though it can be of |
580 | the "weaker" type. |
581 | |
582 | [!] Note that the stores before the write barrier would normally be expected to |
583 | match the loads after the read barrier or the data dependency barrier, and vice |
584 | versa: |
585 | |
586 | CPU 1 CPU 2 |
587 | =============== =============== |
588 | a = 1; }---- --->{ v = c |
589 | b = 2; } \ / { w = d |
590 | <write barrier> \ <read barrier> |
591 | c = 3; } / \ { x = a; |
592 | d = 4; }---- --->{ y = b; |
593 | |
594 | |
595 | EXAMPLES OF MEMORY BARRIER SEQUENCES |
596 | ------------------------------------ |
597 | |
598 | Firstly, write barriers act as partial orderings on store operations. |
599 | Consider the following sequence of events: |
600 | |
601 | CPU 1 |
602 | ======================= |
603 | STORE A = 1 |
604 | STORE B = 2 |
605 | STORE C = 3 |
606 | <write barrier> |
607 | STORE D = 4 |
608 | STORE E = 5 |
609 | |
610 | This sequence of events is committed to the memory coherence system in an order |
611 | that the rest of the system might perceive as the unordered set of { STORE A, |
612 | STORE B, STORE C } all occurring before the unordered set of { STORE D, STORE E |
613 | }: |
614 | |
615 | +-------+ : : |
616 | | | +------+ |
617 | | |------>| C=3 | } /\ |
618 | | | : +------+ }----- \ -----> Events perceptible to |
619 | | | : | A=1 | } \/ the rest of the system |
620 | | | : +------+ } |
621 | | CPU 1 | : | B=2 | } |
622 | | | +------+ } |
623 | | | wwwwwwwwwwwwwwww } <--- At this point the write barrier |
624 | | | +------+ } requires all stores prior to the |
625 | | | : | E=5 | } barrier to be committed before |
626 | | | : +------+ } further stores may take place |
627 | | |------>| D=4 | } |
628 | | | +------+ |
629 | +-------+ : : |
630 | | |
631 | | Sequence in which stores are committed to the |
632 | | memory system by CPU 1 |
633 | V |
634 | |
635 | |
636 | Secondly, data dependency barriers act as partial orderings on data-dependent |
637 | loads. Consider the following sequence of events: |
638 | |
639 | CPU 1 CPU 2 |
640 | ======================= ======================= |
641 | { B = 7; X = 9; Y = 8; C = &Y } |
642 | STORE A = 1 |
643 | STORE B = 2 |
644 | <write barrier> |
645 | STORE C = &B LOAD X |
646 | STORE D = 4 LOAD C (gets &B) |
647 | LOAD *C (reads B) |
648 | |
649 | Without intervention, CPU 2 may perceive the events on CPU 1 in some |
650 | effectively random order, despite the write barrier issued by CPU 1: |
651 | |
652 | +-------+ : : : : |
653 | | | +------+ +-------+ | Sequence of update |
654 | | |------>| B=2 |----- --->| Y->8 | | of perception on |
655 | | | : +------+ \ +-------+ | CPU 2 |
656 | | CPU 1 | : | A=1 | \ --->| C->&Y | V |
657 | | | +------+ | +-------+ |
658 | | | wwwwwwwwwwwwwwww | : : |
659 | | | +------+ | : : |
660 | | | : | C=&B |--- | : : +-------+ |
661 | | | : +------+ \ | +-------+ | | |
662 | | |------>| D=4 | ----------->| C->&B |------>| | |
663 | | | +------+ | +-------+ | | |
664 | +-------+ : : | : : | | |
665 | | : : | | |
666 | | : : | CPU 2 | |
667 | | +-------+ | | |
668 | Apparently incorrect ---> | | B->7 |------>| | |
669 | perception of B (!) | +-------+ | | |
670 | | : : | | |
671 | | +-------+ | | |
672 | The load of X holds ---> \ | X->9 |------>| | |
673 | up the maintenance \ +-------+ | | |
674 | of coherence of B ----->| B->2 | +-------+ |
675 | +-------+ |
676 | : : |
677 | |
678 | |
679 | In the above example, CPU 2 perceives that B is 7, despite the load of *C |
680 | (which would be B) coming after the LOAD of C. |
681 | |
682 | If, however, a data dependency barrier were to be placed between the load of C |
683 | and the load of *C (ie: B) on CPU 2: |
684 | |
685 | CPU 1 CPU 2 |
686 | ======================= ======================= |
687 | { B = 7; X = 9; Y = 8; C = &Y } |
688 | STORE A = 1 |
689 | STORE B = 2 |
690 | <write barrier> |
691 | STORE C = &B LOAD X |
692 | STORE D = 4 LOAD C (gets &B) |
693 | <data dependency barrier> |
694 | LOAD *C (reads B) |
695 | |
696 | then the following will occur: |
697 | |
698 | +-------+ : : : : |
699 | | | +------+ +-------+ |
700 | | |------>| B=2 |----- --->| Y->8 | |
701 | | | : +------+ \ +-------+ |
702 | | CPU 1 | : | A=1 | \ --->| C->&Y | |
703 | | | +------+ | +-------+ |
704 | | | wwwwwwwwwwwwwwww | : : |
705 | | | +------+ | : : |
706 | | | : | C=&B |--- | : : +-------+ |
707 | | | : +------+ \ | +-------+ | | |
708 | | |------>| D=4 | ----------->| C->&B |------>| | |
709 | | | +------+ | +-------+ | | |
710 | +-------+ : : | : : | | |
711 | | : : | | |
712 | | : : | CPU 2 | |
713 | | +-------+ | | |
714 | | | X->9 |------>| | |
715 | | +-------+ | | |
716 | Makes sure all effects ---> \ ddddddddddddddddd | | |
717 | prior to the store of C \ +-------+ | | |
718 | are perceptible to ----->| B->2 |------>| | |
719 | subsequent loads +-------+ | | |
720 | : : +-------+ |
721 | |
722 | |
723 | And thirdly, a read barrier acts as a partial order on loads. Consider the |
724 | following sequence of events: |
725 | |
726 | CPU 1 CPU 2 |
727 | ======================= ======================= |
728 | { A = 0, B = 9 } |
729 | STORE A=1 |
730 | <write barrier> |
731 | STORE B=2 |
732 | LOAD B |
733 | LOAD A |
734 | |
735 | Without intervention, CPU 2 may then choose to perceive the events on CPU 1 in |
736 | some effectively random order, despite the write barrier issued by CPU 1: |
737 | |
738 | +-------+ : : : : |
739 | | | +------+ +-------+ |
740 | | |------>| A=1 |------ --->| A->0 | |
741 | | | +------+ \ +-------+ |
742 | | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 | |
743 | | | +------+ | +-------+ |
744 | | |------>| B=2 |--- | : : |
745 | | | +------+ \ | : : +-------+ |
746 | +-------+ : : \ | +-------+ | | |
747 | ---------->| B->2 |------>| | |
748 | | +-------+ | CPU 2 | |
749 | | | A->0 |------>| | |
750 | | +-------+ | | |
751 | | : : +-------+ |
752 | \ : : |
753 | \ +-------+ |
754 | ---->| A->1 | |
755 | +-------+ |
756 | : : |
757 | |
758 | |
759 | If, however, a read barrier were to be placed between the load of B and the |
760 | load of A on CPU 2: |
761 | |
762 | CPU 1 CPU 2 |
763 | ======================= ======================= |
764 | { A = 0, B = 9 } |
765 | STORE A=1 |
766 | <write barrier> |
767 | STORE B=2 |
768 | LOAD B |
769 | <read barrier> |
770 | LOAD A |
771 | |
772 | then the partial ordering imposed by CPU 1 will be perceived correctly by CPU |
773 | 2: |
774 | |
775 | +-------+ : : : : |
776 | | | +------+ +-------+ |
777 | | |------>| A=1 |------ --->| A->0 | |
778 | | | +------+ \ +-------+ |
779 | | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 | |
780 | | | +------+ | +-------+ |
781 | | |------>| B=2 |--- | : : |
782 | | | +------+ \ | : : +-------+ |
783 | +-------+ : : \ | +-------+ | | |
784 | ---------->| B->2 |------>| | |
785 | | +-------+ | CPU 2 | |
786 | | : : | | |
787 | | : : | | |
788 | At this point the read ----> \ rrrrrrrrrrrrrrrrr | | |
789 | barrier causes all effects \ +-------+ | | |
790 | prior to the storage of B ---->| A->1 |------>| | |
791 | to be perceptible to CPU 2 +-------+ | | |
792 | : : +-------+ |
793 | |
794 | |
795 | To illustrate this more completely, consider what could happen if the code |
796 | contained a load of A either side of the read barrier: |
797 | |
798 | CPU 1 CPU 2 |
799 | ======================= ======================= |
800 | { A = 0, B = 9 } |
801 | STORE A=1 |
802 | <write barrier> |
803 | STORE B=2 |
804 | LOAD B |
805 | LOAD A [first load of A] |
806 | <read barrier> |
807 | LOAD A [second load of A] |
808 | |
809 | Even though the two loads of A both occur after the load of B, they may both |
810 | come up with different values: |
811 | |
812 | +-------+ : : : : |
813 | | | +------+ +-------+ |
814 | | |------>| A=1 |------ --->| A->0 | |
815 | | | +------+ \ +-------+ |
816 | | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 | |
817 | | | +------+ | +-------+ |
818 | | |------>| B=2 |--- | : : |
819 | | | +------+ \ | : : +-------+ |
820 | +-------+ : : \ | +-------+ | | |
821 | ---------->| B->2 |------>| | |
822 | | +-------+ | CPU 2 | |
823 | | : : | | |
824 | | : : | | |
825 | | +-------+ | | |
826 | | | A->0 |------>| 1st | |
827 | | +-------+ | | |
828 | At this point the read ----> \ rrrrrrrrrrrrrrrrr | | |
829 | barrier causes all effects \ +-------+ | | |
830 | prior to the storage of B ---->| A->1 |------>| 2nd | |
831 | to be perceptible to CPU 2 +-------+ | | |
832 | : : +-------+ |
833 | |
834 | |
835 | But it may be that the update to A from CPU 1 becomes perceptible to CPU 2 |
836 | before the read barrier completes anyway: |
837 | |
838 | +-------+ : : : : |
839 | | | +------+ +-------+ |
840 | | |------>| A=1 |------ --->| A->0 | |
841 | | | +------+ \ +-------+ |
842 | | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 | |
843 | | | +------+ | +-------+ |
844 | | |------>| B=2 |--- | : : |
845 | | | +------+ \ | : : +-------+ |
846 | +-------+ : : \ | +-------+ | | |
847 | ---------->| B->2 |------>| | |
848 | | +-------+ | CPU 2 | |
849 | | : : | | |
850 | \ : : | | |
851 | \ +-------+ | | |
852 | ---->| A->1 |------>| 1st | |
853 | +-------+ | | |
854 | rrrrrrrrrrrrrrrrr | | |
855 | +-------+ | | |
856 | | A->1 |------>| 2nd | |
857 | +-------+ | | |
858 | : : +-------+ |
859 | |
860 | |
861 | The guarantee is that the second load will always come up with A == 1 if the |
862 | load of B came up with B == 2. No such guarantee exists for the first load of |
863 | A; that may come up with either A == 0 or A == 1. |
864 | |
865 | |
866 | READ MEMORY BARRIERS VS LOAD SPECULATION |
867 | ---------------------------------------- |
868 | |
869 | Many CPUs speculate with loads: that is they see that they will need to load an |
870 | item from memory, and they find a time where they're not using the bus for any |
871 | other loads, and so do the load in advance - even though they haven't actually |
872 | got to that point in the instruction execution flow yet. This permits the |
873 | actual load instruction to potentially complete immediately because the CPU |
874 | already has the value to hand. |
875 | |
876 | It may turn out that the CPU didn't actually need the value - perhaps because a |
877 | branch circumvented the load - in which case it can discard the value or just |
878 | cache it for later use. |
879 | |
880 | Consider: |
881 | |
882 | CPU 1 CPU 2 |
883 | ======================= ======================= |
884 | LOAD B |
885 | DIVIDE } Divide instructions generally |
886 | DIVIDE } take a long time to perform |
887 | LOAD A |
888 | |
889 | Which might appear as this: |
890 | |
891 | : : +-------+ |
892 | +-------+ | | |
893 | --->| B->2 |------>| | |
894 | +-------+ | CPU 2 | |
895 | : :DIVIDE | | |
896 | +-------+ | | |
897 | The CPU being busy doing a ---> --->| A->0 |~~~~ | | |
898 | division speculates on the +-------+ ~ | | |
899 | LOAD of A : : ~ | | |
900 | : :DIVIDE | | |
901 | : : ~ | | |
902 | Once the divisions are complete --> : : ~-->| | |
903 | the CPU can then perform the : : | | |
904 | LOAD with immediate effect : : +-------+ |
905 | |
906 | |
907 | Placing a read barrier or a data dependency barrier just before the second |
908 | load: |
909 | |
910 | CPU 1 CPU 2 |
911 | ======================= ======================= |
912 | LOAD B |
913 | DIVIDE |
914 | DIVIDE |
915 | <read barrier> |
916 | LOAD A |
917 | |
918 | will force any value speculatively obtained to be reconsidered to an extent |
919 | dependent on the type of barrier used. If there was no change made to the |
920 | speculated memory location, then the speculated value will just be used: |
921 | |
922 | : : +-------+ |
923 | +-------+ | | |
924 | --->| B->2 |------>| | |
925 | +-------+ | CPU 2 | |
926 | : :DIVIDE | | |
927 | +-------+ | | |
928 | The CPU being busy doing a ---> --->| A->0 |~~~~ | | |
929 | division speculates on the +-------+ ~ | | |
930 | LOAD of A : : ~ | | |
931 | : :DIVIDE | | |
932 | : : ~ | | |
933 | : : ~ | | |
934 | rrrrrrrrrrrrrrrr~ | | |
935 | : : ~ | | |
936 | : : ~-->| | |
937 | : : | | |
938 | : : +-------+ |
939 | |
940 | |
941 | but if there was an update or an invalidation from another CPU pending, then |
942 | the speculation will be cancelled and the value reloaded: |
943 | |
944 | : : +-------+ |
945 | +-------+ | | |
946 | --->| B->2 |------>| | |
947 | +-------+ | CPU 2 | |
948 | : :DIVIDE | | |
949 | +-------+ | | |
950 | The CPU being busy doing a ---> --->| A->0 |~~~~ | | |
951 | division speculates on the +-------+ ~ | | |
952 | LOAD of A : : ~ | | |
953 | : :DIVIDE | | |
954 | : : ~ | | |
955 | : : ~ | | |
956 | rrrrrrrrrrrrrrrrr | | |
957 | +-------+ | | |
958 | The speculation is discarded ---> --->| A->1 |------>| | |
959 | and an updated value is +-------+ | | |
960 | retrieved : : +-------+ |
961 | |
962 | |
963 | TRANSITIVITY |
964 | ------------ |
965 | |
966 | Transitivity is a deeply intuitive notion about ordering that is not |
967 | always provided by real computer systems. The following example |
968 | demonstrates transitivity (also called "cumulativity"): |
969 | |
970 | CPU 1 CPU 2 CPU 3 |
971 | ======================= ======================= ======================= |
972 | { X = 0, Y = 0 } |
973 | STORE X=1 LOAD X STORE Y=1 |
974 | <general barrier> <general barrier> |
975 | LOAD Y LOAD X |
976 | |
977 | Suppose that CPU 2's load from X returns 1 and its load from Y returns 0. |
978 | This indicates that CPU 2's load from X in some sense follows CPU 1's |
979 | store to X and that CPU 2's load from Y in some sense preceded CPU 3's |
980 | store to Y. The question is then "Can CPU 3's load from X return 0?" |
981 | |
982 | Because CPU 2's load from X in some sense came after CPU 1's store, it |
983 | is natural to expect that CPU 3's load from X must therefore return 1. |
984 | This expectation is an example of transitivity: if a load executing on |
985 | CPU A follows a load from the same variable executing on CPU B, then |
986 | CPU A's load must either return the same value that CPU B's load did, |
987 | or must return some later value. |
988 | |
989 | In the Linux kernel, use of general memory barriers guarantees |
990 | transitivity. Therefore, in the above example, if CPU 2's load from X |
991 | returns 1 and its load from Y returns 0, then CPU 3's load from X must |
992 | also return 1. |
993 | |
994 | However, transitivity is -not- guaranteed for read or write barriers. |
995 | For example, suppose that CPU 2's general barrier in the above example |
996 | is changed to a read barrier as shown below: |
997 | |
998 | CPU 1 CPU 2 CPU 3 |
999 | ======================= ======================= ======================= |
1000 | { X = 0, Y = 0 } |
1001 | STORE X=1 LOAD X STORE Y=1 |
1002 | <read barrier> <general barrier> |
1003 | LOAD Y LOAD X |
1004 | |
1005 | This substitution destroys transitivity: in this example, it is perfectly |
1006 | legal for CPU 2's load from X to return 1, its load from Y to return 0, |
1007 | and CPU 3's load from X to return 0. |
1008 | |
1009 | The key point is that although CPU 2's read barrier orders its pair |
1010 | of loads, it does not guarantee to order CPU 1's store. Therefore, if |
1011 | this example runs on a system where CPUs 1 and 2 share a store buffer |
1012 | or a level of cache, CPU 2 might have early access to CPU 1's writes. |
1013 | General barriers are therefore required to ensure that all CPUs agree |
1014 | on the combined order of CPU 1's and CPU 2's accesses. |
1015 | |
1016 | To reiterate, if your code requires transitivity, use general barriers |
1017 | throughout. |
1018 | |
1019 | |
1020 | ======================== |
1021 | EXPLICIT KERNEL BARRIERS |
1022 | ======================== |
1023 | |
1024 | The Linux kernel has a variety of different barriers that act at different |
1025 | levels: |
1026 | |
1027 | (*) Compiler barrier. |
1028 | |
1029 | (*) CPU memory barriers. |
1030 | |
1031 | (*) MMIO write barrier. |
1032 | |
1033 | |
1034 | COMPILER BARRIER |
1035 | ---------------- |
1036 | |
1037 | The Linux kernel has an explicit compiler barrier function that prevents the |
1038 | compiler from moving the memory accesses either side of it to the other side: |
1039 | |
1040 | barrier(); |
1041 | |
1042 | This is a general barrier - lesser varieties of compiler barrier do not exist. |
1043 | |
1044 | The compiler barrier has no direct effect on the CPU, which may then reorder |
1045 | things however it wishes. |
1046 | |
1047 | |
1048 | CPU MEMORY BARRIERS |
1049 | ------------------- |
1050 | |
1051 | The Linux kernel has eight basic CPU memory barriers: |
1052 | |
1053 | TYPE MANDATORY SMP CONDITIONAL |
1054 | =============== ======================= =========================== |
1055 | GENERAL mb() smp_mb() |
1056 | WRITE wmb() smp_wmb() |
1057 | READ rmb() smp_rmb() |
1058 | DATA DEPENDENCY read_barrier_depends() smp_read_barrier_depends() |
1059 | |
1060 | |
1061 | All memory barriers except the data dependency barriers imply a compiler |
1062 | barrier. Data dependencies do not impose any additional compiler ordering. |
1063 | |
1064 | Aside: In the case of data dependencies, the compiler would be expected to |
1065 | issue the loads in the correct order (eg. `a[b]` would have to load the value |
1066 | of b before loading a[b]), however there is no guarantee in the C specification |
1067 | that the compiler may not speculate the value of b (eg. is equal to 1) and load |
1068 | a before b (eg. tmp = a[1]; if (b != 1) tmp = a[b]; ). There is also the |
1069 | problem of a compiler reloading b after having loaded a[b], thus having a newer |
1070 | copy of b than a[b]. A consensus has not yet been reached about these problems, |
1071 | however the ACCESS_ONCE macro is a good place to start looking. |
1072 | |
1073 | SMP memory barriers are reduced to compiler barriers on uniprocessor compiled |
1074 | systems because it is assumed that a CPU will appear to be self-consistent, |
1075 | and will order overlapping accesses correctly with respect to itself. |
1076 | |
1077 | [!] Note that SMP memory barriers _must_ be used to control the ordering of |
1078 | references to shared memory on SMP systems, though the use of locking instead |
1079 | is sufficient. |
1080 | |
1081 | Mandatory barriers should not be used to control SMP effects, since mandatory |
1082 | barriers unnecessarily impose overhead on UP systems. They may, however, be |
1083 | used to control MMIO effects on accesses through relaxed memory I/O windows. |
1084 | These are required even on non-SMP systems as they affect the order in which |
1085 | memory operations appear to a device by prohibiting both the compiler and the |
1086 | CPU from reordering them. |
1087 | |
1088 | |
1089 | There are some more advanced barrier functions: |
1090 | |
1091 | (*) set_mb(var, value) |
1092 | |
1093 | This assigns the value to the variable and then inserts a full memory |
1094 | barrier after it, depending on the function. It isn't guaranteed to |
1095 | insert anything more than a compiler barrier in a UP compilation. |
1096 | |
1097 | |
1098 | (*) smp_mb__before_atomic_dec(); |
1099 | (*) smp_mb__after_atomic_dec(); |
1100 | (*) smp_mb__before_atomic_inc(); |
1101 | (*) smp_mb__after_atomic_inc(); |
1102 | |
1103 | These are for use with atomic add, subtract, increment and decrement |
1104 | functions that don't return a value, especially when used for reference |
1105 | counting. These functions do not imply memory barriers. |
1106 | |
1107 | As an example, consider a piece of code that marks an object as being dead |
1108 | and then decrements the object's reference count: |
1109 | |
1110 | obj->dead = 1; |
1111 | smp_mb__before_atomic_dec(); |
1112 | atomic_dec(&obj->ref_count); |
1113 | |
1114 | This makes sure that the death mark on the object is perceived to be set |
1115 | *before* the reference counter is decremented. |
1116 | |
1117 | See Documentation/atomic_ops.txt for more information. See the "Atomic |
1118 | operations" subsection for information on where to use these. |
1119 | |
1120 | |
1121 | (*) smp_mb__before_clear_bit(void); |
1122 | (*) smp_mb__after_clear_bit(void); |
1123 | |
1124 | These are for use similar to the atomic inc/dec barriers. These are |
1125 | typically used for bitwise unlocking operations, so care must be taken as |
1126 | there are no implicit memory barriers here either. |
1127 | |
1128 | Consider implementing an unlock operation of some nature by clearing a |
1129 | locking bit. The clear_bit() would then need to be barriered like this: |
1130 | |
1131 | smp_mb__before_clear_bit(); |
1132 | clear_bit( ... ); |
1133 | |
1134 | This prevents memory operations before the clear leaking to after it. See |
1135 | the subsection on "Locking Functions" with reference to UNLOCK operation |
1136 | implications. |
1137 | |
1138 | See Documentation/atomic_ops.txt for more information. See the "Atomic |
1139 | operations" subsection for information on where to use these. |
1140 | |
1141 | |
1142 | MMIO WRITE BARRIER |
1143 | ------------------ |
1144 | |
1145 | The Linux kernel also has a special barrier for use with memory-mapped I/O |
1146 | writes: |
1147 | |
1148 | mmiowb(); |
1149 | |
1150 | This is a variation on the mandatory write barrier that causes writes to weakly |
1151 | ordered I/O regions to be partially ordered. Its effects may go beyond the |
1152 | CPU->Hardware interface and actually affect the hardware at some level. |
1153 | |
1154 | See the subsection "Locks vs I/O accesses" for more information. |
1155 | |
1156 | |
1157 | =============================== |
1158 | IMPLICIT KERNEL MEMORY BARRIERS |
1159 | =============================== |
1160 | |
1161 | Some of the other functions in the linux kernel imply memory barriers, amongst |
1162 | which are locking and scheduling functions. |
1163 | |
1164 | This specification is a _minimum_ guarantee; any particular architecture may |
1165 | provide more substantial guarantees, but these may not be relied upon outside |
1166 | of arch specific code. |
1167 | |
1168 | |
1169 | LOCKING FUNCTIONS |
1170 | ----------------- |
1171 | |
1172 | The Linux kernel has a number of locking constructs: |
1173 | |
1174 | (*) spin locks |
1175 | (*) R/W spin locks |
1176 | (*) mutexes |
1177 | (*) semaphores |
1178 | (*) R/W semaphores |
1179 | (*) RCU |
1180 | |
1181 | In all cases there are variants on "LOCK" operations and "UNLOCK" operations |
1182 | for each construct. These operations all imply certain barriers: |
1183 | |
1184 | (1) LOCK operation implication: |
1185 | |
1186 | Memory operations issued after the LOCK will be completed after the LOCK |
1187 | operation has completed. |
1188 | |
1189 | Memory operations issued before the LOCK may be completed after the LOCK |
1190 | operation has completed. |
1191 | |
1192 | (2) UNLOCK operation implication: |
1193 | |
1194 | Memory operations issued before the UNLOCK will be completed before the |
1195 | UNLOCK operation has completed. |
1196 | |
1197 | Memory operations issued after the UNLOCK may be completed before the |
1198 | UNLOCK operation has completed. |
1199 | |
1200 | (3) LOCK vs LOCK implication: |
1201 | |
1202 | All LOCK operations issued before another LOCK operation will be completed |
1203 | before that LOCK operation. |
1204 | |
1205 | (4) LOCK vs UNLOCK implication: |
1206 | |
1207 | All LOCK operations issued before an UNLOCK operation will be completed |
1208 | before the UNLOCK operation. |
1209 | |
1210 | All UNLOCK operations issued before a LOCK operation will be completed |
1211 | before the LOCK operation. |
1212 | |
1213 | (5) Failed conditional LOCK implication: |
1214 | |
1215 | Certain variants of the LOCK operation may fail, either due to being |
1216 | unable to get the lock immediately, or due to receiving an unblocked |
1217 | signal whilst asleep waiting for the lock to become available. Failed |
1218 | locks do not imply any sort of barrier. |
1219 | |
1220 | Therefore, from (1), (2) and (4) an UNLOCK followed by an unconditional LOCK is |
1221 | equivalent to a full barrier, but a LOCK followed by an UNLOCK is not. |
1222 | |
1223 | [!] Note: one of the consequences of LOCKs and UNLOCKs being only one-way |
1224 | barriers is that the effects of instructions outside of a critical section |
1225 | may seep into the inside of the critical section. |
1226 | |
1227 | A LOCK followed by an UNLOCK may not be assumed to be full memory barrier |
1228 | because it is possible for an access preceding the LOCK to happen after the |
1229 | LOCK, and an access following the UNLOCK to happen before the UNLOCK, and the |
1230 | two accesses can themselves then cross: |
1231 | |
1232 | *A = a; |
1233 | LOCK |
1234 | UNLOCK |
1235 | *B = b; |
1236 | |
1237 | may occur as: |
1238 | |
1239 | LOCK, STORE *B, STORE *A, UNLOCK |
1240 | |
1241 | Locks and semaphores may not provide any guarantee of ordering on UP compiled |
1242 | systems, and so cannot be counted on in such a situation to actually achieve |
1243 | anything at all - especially with respect to I/O accesses - unless combined |
1244 | with interrupt disabling operations. |
1245 | |
1246 | See also the section on "Inter-CPU locking barrier effects". |
1247 | |
1248 | |
1249 | As an example, consider the following: |
1250 | |
1251 | *A = a; |
1252 | *B = b; |
1253 | LOCK |
1254 | *C = c; |
1255 | *D = d; |
1256 | UNLOCK |
1257 | *E = e; |
1258 | *F = f; |
1259 | |
1260 | The following sequence of events is acceptable: |
1261 | |
1262 | LOCK, {*F,*A}, *E, {*C,*D}, *B, UNLOCK |
1263 | |
1264 | [+] Note that {*F,*A} indicates a combined access. |
1265 | |
1266 | But none of the following are: |
1267 | |
1268 | {*F,*A}, *B, LOCK, *C, *D, UNLOCK, *E |
1269 | *A, *B, *C, LOCK, *D, UNLOCK, *E, *F |
1270 | *A, *B, LOCK, *C, UNLOCK, *D, *E, *F |
1271 | *B, LOCK, *C, *D, UNLOCK, {*F,*A}, *E |
1272 | |
1273 | |
1274 | |
1275 | INTERRUPT DISABLING FUNCTIONS |
1276 | ----------------------------- |
1277 | |
1278 | Functions that disable interrupts (LOCK equivalent) and enable interrupts |
1279 | (UNLOCK equivalent) will act as compiler barriers only. So if memory or I/O |
1280 | barriers are required in such a situation, they must be provided from some |
1281 | other means. |
1282 | |
1283 | |
1284 | SLEEP AND WAKE-UP FUNCTIONS |
1285 | --------------------------- |
1286 | |
1287 | Sleeping and waking on an event flagged in global data can be viewed as an |
1288 | interaction between two pieces of data: the task state of the task waiting for |
1289 | the event and the global data used to indicate the event. To make sure that |
1290 | these appear to happen in the right order, the primitives to begin the process |
1291 | of going to sleep, and the primitives to initiate a wake up imply certain |
1292 | barriers. |
1293 | |
1294 | Firstly, the sleeper normally follows something like this sequence of events: |
1295 | |
1296 | for (;;) { |
1297 | set_current_state(TASK_UNINTERRUPTIBLE); |
1298 | if (event_indicated) |
1299 | break; |
1300 | schedule(); |
1301 | } |
1302 | |
1303 | A general memory barrier is interpolated automatically by set_current_state() |
1304 | after it has altered the task state: |
1305 | |
1306 | CPU 1 |
1307 | =============================== |
1308 | set_current_state(); |
1309 | set_mb(); |
1310 | STORE current->state |
1311 | <general barrier> |
1312 | LOAD event_indicated |
1313 | |
1314 | set_current_state() may be wrapped by: |
1315 | |
1316 | prepare_to_wait(); |
1317 | prepare_to_wait_exclusive(); |
1318 | |
1319 | which therefore also imply a general memory barrier after setting the state. |
1320 | The whole sequence above is available in various canned forms, all of which |
1321 | interpolate the memory barrier in the right place: |
1322 | |
1323 | wait_event(); |
1324 | wait_event_interruptible(); |
1325 | wait_event_interruptible_exclusive(); |
1326 | wait_event_interruptible_timeout(); |
1327 | wait_event_killable(); |
1328 | wait_event_timeout(); |
1329 | wait_on_bit(); |
1330 | wait_on_bit_lock(); |
1331 | |
1332 | |
1333 | Secondly, code that performs a wake up normally follows something like this: |
1334 | |
1335 | event_indicated = 1; |
1336 | wake_up(&event_wait_queue); |
1337 | |
1338 | or: |
1339 | |
1340 | event_indicated = 1; |
1341 | wake_up_process(event_daemon); |
1342 | |
1343 | A write memory barrier is implied by wake_up() and co. if and only if they wake |
1344 | something up. The barrier occurs before the task state is cleared, and so sits |
1345 | between the STORE to indicate the event and the STORE to set TASK_RUNNING: |
1346 | |
1347 | CPU 1 CPU 2 |
1348 | =============================== =============================== |
1349 | set_current_state(); STORE event_indicated |
1350 | set_mb(); wake_up(); |
1351 | STORE current->state <write barrier> |
1352 | <general barrier> STORE current->state |
1353 | LOAD event_indicated |
1354 | |
1355 | The available waker functions include: |
1356 | |
1357 | complete(); |
1358 | wake_up(); |
1359 | wake_up_all(); |
1360 | wake_up_bit(); |
1361 | wake_up_interruptible(); |
1362 | wake_up_interruptible_all(); |
1363 | wake_up_interruptible_nr(); |
1364 | wake_up_interruptible_poll(); |
1365 | wake_up_interruptible_sync(); |
1366 | wake_up_interruptible_sync_poll(); |
1367 | wake_up_locked(); |
1368 | wake_up_locked_poll(); |
1369 | wake_up_nr(); |
1370 | wake_up_poll(); |
1371 | wake_up_process(); |
1372 | |
1373 | |
1374 | [!] Note that the memory barriers implied by the sleeper and the waker do _not_ |
1375 | order multiple stores before the wake-up with respect to loads of those stored |
1376 | values after the sleeper has called set_current_state(). For instance, if the |
1377 | sleeper does: |
1378 | |
1379 | set_current_state(TASK_INTERRUPTIBLE); |
1380 | if (event_indicated) |
1381 | break; |
1382 | __set_current_state(TASK_RUNNING); |
1383 | do_something(my_data); |
1384 | |
1385 | and the waker does: |
1386 | |
1387 | my_data = value; |
1388 | event_indicated = 1; |
1389 | wake_up(&event_wait_queue); |
1390 | |
1391 | there's no guarantee that the change to event_indicated will be perceived by |
1392 | the sleeper as coming after the change to my_data. In such a circumstance, the |
1393 | code on both sides must interpolate its own memory barriers between the |
1394 | separate data accesses. Thus the above sleeper ought to do: |
1395 | |
1396 | set_current_state(TASK_INTERRUPTIBLE); |
1397 | if (event_indicated) { |
1398 | smp_rmb(); |
1399 | do_something(my_data); |
1400 | } |
1401 | |
1402 | and the waker should do: |
1403 | |
1404 | my_data = value; |
1405 | smp_wmb(); |
1406 | event_indicated = 1; |
1407 | wake_up(&event_wait_queue); |
1408 | |
1409 | |
1410 | MISCELLANEOUS FUNCTIONS |
1411 | ----------------------- |
1412 | |
1413 | Other functions that imply barriers: |
1414 | |
1415 | (*) schedule() and similar imply full memory barriers. |
1416 | |
1417 | |
1418 | ================================= |
1419 | INTER-CPU LOCKING BARRIER EFFECTS |
1420 | ================================= |
1421 | |
1422 | On SMP systems locking primitives give a more substantial form of barrier: one |
1423 | that does affect memory access ordering on other CPUs, within the context of |
1424 | conflict on any particular lock. |
1425 | |
1426 | |
1427 | LOCKS VS MEMORY ACCESSES |
1428 | ------------------------ |
1429 | |
1430 | Consider the following: the system has a pair of spinlocks (M) and (Q), and |
1431 | three CPUs; then should the following sequence of events occur: |
1432 | |
1433 | CPU 1 CPU 2 |
1434 | =============================== =============================== |
1435 | *A = a; *E = e; |
1436 | LOCK M LOCK Q |
1437 | *B = b; *F = f; |
1438 | *C = c; *G = g; |
1439 | UNLOCK M UNLOCK Q |
1440 | *D = d; *H = h; |
1441 | |
1442 | Then there is no guarantee as to what order CPU 3 will see the accesses to *A |
1443 | through *H occur in, other than the constraints imposed by the separate locks |
1444 | on the separate CPUs. It might, for example, see: |
1445 | |
1446 | *E, LOCK M, LOCK Q, *G, *C, *F, *A, *B, UNLOCK Q, *D, *H, UNLOCK M |
1447 | |
1448 | But it won't see any of: |
1449 | |
1450 | *B, *C or *D preceding LOCK M |
1451 | *A, *B or *C following UNLOCK M |
1452 | *F, *G or *H preceding LOCK Q |
1453 | *E, *F or *G following UNLOCK Q |
1454 | |
1455 | |
1456 | However, if the following occurs: |
1457 | |
1458 | CPU 1 CPU 2 |
1459 | =============================== =============================== |
1460 | *A = a; |
1461 | LOCK M [1] |
1462 | *B = b; |
1463 | *C = c; |
1464 | UNLOCK M [1] |
1465 | *D = d; *E = e; |
1466 | LOCK M [2] |
1467 | *F = f; |
1468 | *G = g; |
1469 | UNLOCK M [2] |
1470 | *H = h; |
1471 | |
1472 | CPU 3 might see: |
1473 | |
1474 | *E, LOCK M [1], *C, *B, *A, UNLOCK M [1], |
1475 | LOCK M [2], *H, *F, *G, UNLOCK M [2], *D |
1476 | |
1477 | But assuming CPU 1 gets the lock first, CPU 3 won't see any of: |
1478 | |
1479 | *B, *C, *D, *F, *G or *H preceding LOCK M [1] |
1480 | *A, *B or *C following UNLOCK M [1] |
1481 | *F, *G or *H preceding LOCK M [2] |
1482 | *A, *B, *C, *E, *F or *G following UNLOCK M [2] |
1483 | |
1484 | |
1485 | LOCKS VS I/O ACCESSES |
1486 | --------------------- |
1487 | |
1488 | Under certain circumstances (especially involving NUMA), I/O accesses within |
1489 | two spinlocked sections on two different CPUs may be seen as interleaved by the |
1490 | PCI bridge, because the PCI bridge does not necessarily participate in the |
1491 | cache-coherence protocol, and is therefore incapable of issuing the required |
1492 | read memory barriers. |
1493 | |
1494 | For example: |
1495 | |
1496 | CPU 1 CPU 2 |
1497 | =============================== =============================== |
1498 | spin_lock(Q) |
1499 | writel(0, ADDR) |
1500 | writel(1, DATA); |
1501 | spin_unlock(Q); |
1502 | spin_lock(Q); |
1503 | writel(4, ADDR); |
1504 | writel(5, DATA); |
1505 | spin_unlock(Q); |
1506 | |
1507 | may be seen by the PCI bridge as follows: |
1508 | |
1509 | STORE *ADDR = 0, STORE *ADDR = 4, STORE *DATA = 1, STORE *DATA = 5 |
1510 | |
1511 | which would probably cause the hardware to malfunction. |
1512 | |
1513 | |
1514 | What is necessary here is to intervene with an mmiowb() before dropping the |
1515 | spinlock, for example: |
1516 | |
1517 | CPU 1 CPU 2 |
1518 | =============================== =============================== |
1519 | spin_lock(Q) |
1520 | writel(0, ADDR) |
1521 | writel(1, DATA); |
1522 | mmiowb(); |
1523 | spin_unlock(Q); |
1524 | spin_lock(Q); |
1525 | writel(4, ADDR); |
1526 | writel(5, DATA); |
1527 | mmiowb(); |
1528 | spin_unlock(Q); |
1529 | |
1530 | this will ensure that the two stores issued on CPU 1 appear at the PCI bridge |
1531 | before either of the stores issued on CPU 2. |
1532 | |
1533 | |
1534 | Furthermore, following a store by a load from the same device obviates the need |
1535 | for the mmiowb(), because the load forces the store to complete before the load |
1536 | is performed: |
1537 | |
1538 | CPU 1 CPU 2 |
1539 | =============================== =============================== |
1540 | spin_lock(Q) |
1541 | writel(0, ADDR) |
1542 | a = readl(DATA); |
1543 | spin_unlock(Q); |
1544 | spin_lock(Q); |
1545 | writel(4, ADDR); |
1546 | b = readl(DATA); |
1547 | spin_unlock(Q); |
1548 | |
1549 | |
1550 | See Documentation/DocBook/deviceiobook.tmpl for more information. |
1551 | |
1552 | |
1553 | ================================= |
1554 | WHERE ARE MEMORY BARRIERS NEEDED? |
1555 | ================================= |
1556 | |
1557 | Under normal operation, memory operation reordering is generally not going to |
1558 | be a problem as a single-threaded linear piece of code will still appear to |
1559 | work correctly, even if it's in an SMP kernel. There are, however, four |
1560 | circumstances in which reordering definitely _could_ be a problem: |
1561 | |
1562 | (*) Interprocessor interaction. |
1563 | |
1564 | (*) Atomic operations. |
1565 | |
1566 | (*) Accessing devices. |
1567 | |
1568 | (*) Interrupts. |
1569 | |
1570 | |
1571 | INTERPROCESSOR INTERACTION |
1572 | -------------------------- |
1573 | |
1574 | When there's a system with more than one processor, more than one CPU in the |
1575 | system may be working on the same data set at the same time. This can cause |
1576 | synchronisation problems, and the usual way of dealing with them is to use |
1577 | locks. Locks, however, are quite expensive, and so it may be preferable to |
1578 | operate without the use of a lock if at all possible. In such a case |
1579 | operations that affect both CPUs may have to be carefully ordered to prevent |
1580 | a malfunction. |
1581 | |
1582 | Consider, for example, the R/W semaphore slow path. Here a waiting process is |
1583 | queued on the semaphore, by virtue of it having a piece of its stack linked to |
1584 | the semaphore's list of waiting processes: |
1585 | |
1586 | struct rw_semaphore { |
1587 | ... |
1588 | spinlock_t lock; |
1589 | struct list_head waiters; |
1590 | }; |
1591 | |
1592 | struct rwsem_waiter { |
1593 | struct list_head list; |
1594 | struct task_struct *task; |
1595 | }; |
1596 | |
1597 | To wake up a particular waiter, the up_read() or up_write() functions have to: |
1598 | |
1599 | (1) read the next pointer from this waiter's record to know as to where the |
1600 | next waiter record is; |
1601 | |
1602 | (2) read the pointer to the waiter's task structure; |
1603 | |
1604 | (3) clear the task pointer to tell the waiter it has been given the semaphore; |
1605 | |
1606 | (4) call wake_up_process() on the task; and |
1607 | |
1608 | (5) release the reference held on the waiter's task struct. |
1609 | |
1610 | In other words, it has to perform this sequence of events: |
1611 | |
1612 | LOAD waiter->list.next; |
1613 | LOAD waiter->task; |
1614 | STORE waiter->task; |
1615 | CALL wakeup |
1616 | RELEASE task |
1617 | |
1618 | and if any of these steps occur out of order, then the whole thing may |
1619 | malfunction. |
1620 | |
1621 | Once it has queued itself and dropped the semaphore lock, the waiter does not |
1622 | get the lock again; it instead just waits for its task pointer to be cleared |
1623 | before proceeding. Since the record is on the waiter's stack, this means that |
1624 | if the task pointer is cleared _before_ the next pointer in the list is read, |
1625 | another CPU might start processing the waiter and might clobber the waiter's |
1626 | stack before the up*() function has a chance to read the next pointer. |
1627 | |
1628 | Consider then what might happen to the above sequence of events: |
1629 | |
1630 | CPU 1 CPU 2 |
1631 | =============================== =============================== |
1632 | down_xxx() |
1633 | Queue waiter |
1634 | Sleep |
1635 | up_yyy() |
1636 | LOAD waiter->task; |
1637 | STORE waiter->task; |
1638 | Woken up by other event |
1639 | <preempt> |
1640 | Resume processing |
1641 | down_xxx() returns |
1642 | call foo() |
1643 | foo() clobbers *waiter |
1644 | </preempt> |
1645 | LOAD waiter->list.next; |
1646 | --- OOPS --- |
1647 | |
1648 | This could be dealt with using the semaphore lock, but then the down_xxx() |
1649 | function has to needlessly get the spinlock again after being woken up. |
1650 | |
1651 | The way to deal with this is to insert a general SMP memory barrier: |
1652 | |
1653 | LOAD waiter->list.next; |
1654 | LOAD waiter->task; |
1655 | smp_mb(); |
1656 | STORE waiter->task; |
1657 | CALL wakeup |
1658 | RELEASE task |
1659 | |
1660 | In this case, the barrier makes a guarantee that all memory accesses before the |
1661 | barrier will appear to happen before all the memory accesses after the barrier |
1662 | with respect to the other CPUs on the system. It does _not_ guarantee that all |
1663 | the memory accesses before the barrier will be complete by the time the barrier |
1664 | instruction itself is complete. |
1665 | |
1666 | On a UP system - where this wouldn't be a problem - the smp_mb() is just a |
1667 | compiler barrier, thus making sure the compiler emits the instructions in the |
1668 | right order without actually intervening in the CPU. Since there's only one |
1669 | CPU, that CPU's dependency ordering logic will take care of everything else. |
1670 | |
1671 | |
1672 | ATOMIC OPERATIONS |
1673 | ----------------- |
1674 | |
1675 | Whilst they are technically interprocessor interaction considerations, atomic |
1676 | operations are noted specially as some of them imply full memory barriers and |
1677 | some don't, but they're very heavily relied on as a group throughout the |
1678 | kernel. |
1679 | |
1680 | Any atomic operation that modifies some state in memory and returns information |
1681 | about the state (old or new) implies an SMP-conditional general memory barrier |
1682 | (smp_mb()) on each side of the actual operation (with the exception of |
1683 | explicit lock operations, described later). These include: |
1684 | |
1685 | xchg(); |
1686 | cmpxchg(); |
1687 | atomic_cmpxchg(); |
1688 | atomic_inc_return(); |
1689 | atomic_dec_return(); |
1690 | atomic_add_return(); |
1691 | atomic_sub_return(); |
1692 | atomic_inc_and_test(); |
1693 | atomic_dec_and_test(); |
1694 | atomic_sub_and_test(); |
1695 | atomic_add_negative(); |
1696 | atomic_add_unless(); /* when succeeds (returns 1) */ |
1697 | test_and_set_bit(); |
1698 | test_and_clear_bit(); |
1699 | test_and_change_bit(); |
1700 | |
1701 | These are used for such things as implementing LOCK-class and UNLOCK-class |
1702 | operations and adjusting reference counters towards object destruction, and as |
1703 | such the implicit memory barrier effects are necessary. |
1704 | |
1705 | |
1706 | The following operations are potential problems as they do _not_ imply memory |
1707 | barriers, but might be used for implementing such things as UNLOCK-class |
1708 | operations: |
1709 | |
1710 | atomic_set(); |
1711 | set_bit(); |
1712 | clear_bit(); |
1713 | change_bit(); |
1714 | |
1715 | With these the appropriate explicit memory barrier should be used if necessary |
1716 | (smp_mb__before_clear_bit() for instance). |
1717 | |
1718 | |
1719 | The following also do _not_ imply memory barriers, and so may require explicit |
1720 | memory barriers under some circumstances (smp_mb__before_atomic_dec() for |
1721 | instance): |
1722 | |
1723 | atomic_add(); |
1724 | atomic_sub(); |
1725 | atomic_inc(); |
1726 | atomic_dec(); |
1727 | |
1728 | If they're used for statistics generation, then they probably don't need memory |
1729 | barriers, unless there's a coupling between statistical data. |
1730 | |
1731 | If they're used for reference counting on an object to control its lifetime, |
1732 | they probably don't need memory barriers because either the reference count |
1733 | will be adjusted inside a locked section, or the caller will already hold |
1734 | sufficient references to make the lock, and thus a memory barrier unnecessary. |
1735 | |
1736 | If they're used for constructing a lock of some description, then they probably |
1737 | do need memory barriers as a lock primitive generally has to do things in a |
1738 | specific order. |
1739 | |
1740 | Basically, each usage case has to be carefully considered as to whether memory |
1741 | barriers are needed or not. |
1742 | |
1743 | The following operations are special locking primitives: |
1744 | |
1745 | test_and_set_bit_lock(); |
1746 | clear_bit_unlock(); |
1747 | __clear_bit_unlock(); |
1748 | |
1749 | These implement LOCK-class and UNLOCK-class operations. These should be used in |
1750 | preference to other operations when implementing locking primitives, because |
1751 | their implementations can be optimised on many architectures. |
1752 | |
1753 | [!] Note that special memory barrier primitives are available for these |
1754 | situations because on some CPUs the atomic instructions used imply full memory |
1755 | barriers, and so barrier instructions are superfluous in conjunction with them, |
1756 | and in such cases the special barrier primitives will be no-ops. |
1757 | |
1758 | See Documentation/atomic_ops.txt for more information. |
1759 | |
1760 | |
1761 | ACCESSING DEVICES |
1762 | ----------------- |
1763 | |
1764 | Many devices can be memory mapped, and so appear to the CPU as if they're just |
1765 | a set of memory locations. To control such a device, the driver usually has to |
1766 | make the right memory accesses in exactly the right order. |
1767 | |
1768 | However, having a clever CPU or a clever compiler creates a potential problem |
1769 | in that the carefully sequenced accesses in the driver code won't reach the |
1770 | device in the requisite order if the CPU or the compiler thinks it is more |
1771 | efficient to reorder, combine or merge accesses - something that would cause |
1772 | the device to malfunction. |
1773 | |
1774 | Inside of the Linux kernel, I/O should be done through the appropriate accessor |
1775 | routines - such as inb() or writel() - which know how to make such accesses |
1776 | appropriately sequential. Whilst this, for the most part, renders the explicit |
1777 | use of memory barriers unnecessary, there are a couple of situations where they |
1778 | might be needed: |
1779 | |
1780 | (1) On some systems, I/O stores are not strongly ordered across all CPUs, and |
1781 | so for _all_ general drivers locks should be used and mmiowb() must be |
1782 | issued prior to unlocking the critical section. |
1783 | |
1784 | (2) If the accessor functions are used to refer to an I/O memory window with |
1785 | relaxed memory access properties, then _mandatory_ memory barriers are |
1786 | required to enforce ordering. |
1787 | |
1788 | See Documentation/DocBook/deviceiobook.tmpl for more information. |
1789 | |
1790 | |
1791 | INTERRUPTS |
1792 | ---------- |
1793 | |
1794 | A driver may be interrupted by its own interrupt service routine, and thus the |
1795 | two parts of the driver may interfere with each other's attempts to control or |
1796 | access the device. |
1797 | |
1798 | This may be alleviated - at least in part - by disabling local interrupts (a |
1799 | form of locking), such that the critical operations are all contained within |
1800 | the interrupt-disabled section in the driver. Whilst the driver's interrupt |
1801 | routine is executing, the driver's core may not run on the same CPU, and its |
1802 | interrupt is not permitted to happen again until the current interrupt has been |
1803 | handled, thus the interrupt handler does not need to lock against that. |
1804 | |
1805 | However, consider a driver that was talking to an ethernet card that sports an |
1806 | address register and a data register. If that driver's core talks to the card |
1807 | under interrupt-disablement and then the driver's interrupt handler is invoked: |
1808 | |
1809 | LOCAL IRQ DISABLE |
1810 | writew(ADDR, 3); |
1811 | writew(DATA, y); |
1812 | LOCAL IRQ ENABLE |
1813 | <interrupt> |
1814 | writew(ADDR, 4); |
1815 | q = readw(DATA); |
1816 | </interrupt> |
1817 | |
1818 | The store to the data register might happen after the second store to the |
1819 | address register if ordering rules are sufficiently relaxed: |
1820 | |
1821 | STORE *ADDR = 3, STORE *ADDR = 4, STORE *DATA = y, q = LOAD *DATA |
1822 | |
1823 | |
1824 | If ordering rules are relaxed, it must be assumed that accesses done inside an |
1825 | interrupt disabled section may leak outside of it and may interleave with |
1826 | accesses performed in an interrupt - and vice versa - unless implicit or |
1827 | explicit barriers are used. |
1828 | |
1829 | Normally this won't be a problem because the I/O accesses done inside such |
1830 | sections will include synchronous load operations on strictly ordered I/O |
1831 | registers that form implicit I/O barriers. If this isn't sufficient then an |
1832 | mmiowb() may need to be used explicitly. |
1833 | |
1834 | |
1835 | A similar situation may occur between an interrupt routine and two routines |
1836 | running on separate CPUs that communicate with each other. If such a case is |
1837 | likely, then interrupt-disabling locks should be used to guarantee ordering. |
1838 | |
1839 | |
1840 | ========================== |
1841 | KERNEL I/O BARRIER EFFECTS |
1842 | ========================== |
1843 | |
1844 | When accessing I/O memory, drivers should use the appropriate accessor |
1845 | functions: |
1846 | |
1847 | (*) inX(), outX(): |
1848 | |
1849 | These are intended to talk to I/O space rather than memory space, but |
1850 | that's primarily a CPU-specific concept. The i386 and x86_64 processors do |
1851 | indeed have special I/O space access cycles and instructions, but many |
1852 | CPUs don't have such a concept. |
1853 | |
1854 | The PCI bus, amongst others, defines an I/O space concept which - on such |
1855 | CPUs as i386 and x86_64 - readily maps to the CPU's concept of I/O |
1856 | space. However, it may also be mapped as a virtual I/O space in the CPU's |
1857 | memory map, particularly on those CPUs that don't support alternate I/O |
1858 | spaces. |
1859 | |
1860 | Accesses to this space may be fully synchronous (as on i386), but |
1861 | intermediary bridges (such as the PCI host bridge) may not fully honour |
1862 | that. |
1863 | |
1864 | They are guaranteed to be fully ordered with respect to each other. |
1865 | |
1866 | They are not guaranteed to be fully ordered with respect to other types of |
1867 | memory and I/O operation. |
1868 | |
1869 | (*) readX(), writeX(): |
1870 | |
1871 | Whether these are guaranteed to be fully ordered and uncombined with |
1872 | respect to each other on the issuing CPU depends on the characteristics |
1873 | defined for the memory window through which they're accessing. On later |
1874 | i386 architecture machines, for example, this is controlled by way of the |
1875 | MTRR registers. |
1876 | |
1877 | Ordinarily, these will be guaranteed to be fully ordered and uncombined, |
1878 | provided they're not accessing a prefetchable device. |
1879 | |
1880 | However, intermediary hardware (such as a PCI bridge) may indulge in |
1881 | deferral if it so wishes; to flush a store, a load from the same location |
1882 | is preferred[*], but a load from the same device or from configuration |
1883 | space should suffice for PCI. |
1884 | |
1885 | [*] NOTE! attempting to load from the same location as was written to may |
1886 | cause a malfunction - consider the 16550 Rx/Tx serial registers for |
1887 | example. |
1888 | |
1889 | Used with prefetchable I/O memory, an mmiowb() barrier may be required to |
1890 | force stores to be ordered. |
1891 | |
1892 | Please refer to the PCI specification for more information on interactions |
1893 | between PCI transactions. |
1894 | |
1895 | (*) readX_relaxed() |
1896 | |
1897 | These are similar to readX(), but are not guaranteed to be ordered in any |
1898 | way. Be aware that there is no I/O read barrier available. |
1899 | |
1900 | (*) ioreadX(), iowriteX() |
1901 | |
1902 | These will perform appropriately for the type of access they're actually |
1903 | doing, be it inX()/outX() or readX()/writeX(). |
1904 | |
1905 | |
1906 | ======================================== |
1907 | ASSUMED MINIMUM EXECUTION ORDERING MODEL |
1908 | ======================================== |
1909 | |
1910 | It has to be assumed that the conceptual CPU is weakly-ordered but that it will |
1911 | maintain the appearance of program causality with respect to itself. Some CPUs |
1912 | (such as i386 or x86_64) are more constrained than others (such as powerpc or |
1913 | frv), and so the most relaxed case (namely DEC Alpha) must be assumed outside |
1914 | of arch-specific code. |
1915 | |
1916 | This means that it must be considered that the CPU will execute its instruction |
1917 | stream in any order it feels like - or even in parallel - provided that if an |
1918 | instruction in the stream depends on an earlier instruction, then that |
1919 | earlier instruction must be sufficiently complete[*] before the later |
1920 | instruction may proceed; in other words: provided that the appearance of |
1921 | causality is maintained. |
1922 | |
1923 | [*] Some instructions have more than one effect - such as changing the |
1924 | condition codes, changing registers or changing memory - and different |
1925 | instructions may depend on different effects. |
1926 | |
1927 | A CPU may also discard any instruction sequence that winds up having no |
1928 | ultimate effect. For example, if two adjacent instructions both load an |
1929 | immediate value into the same register, the first may be discarded. |
1930 | |
1931 | |
1932 | Similarly, it has to be assumed that compiler might reorder the instruction |
1933 | stream in any way it sees fit, again provided the appearance of causality is |
1934 | maintained. |
1935 | |
1936 | |
1937 | ============================ |
1938 | THE EFFECTS OF THE CPU CACHE |
1939 | ============================ |
1940 | |
1941 | The way cached memory operations are perceived across the system is affected to |
1942 | a certain extent by the caches that lie between CPUs and memory, and by the |
1943 | memory coherence system that maintains the consistency of state in the system. |
1944 | |
1945 | As far as the way a CPU interacts with another part of the system through the |
1946 | caches goes, the memory system has to include the CPU's caches, and memory |
1947 | barriers for the most part act at the interface between the CPU and its cache |
1948 | (memory barriers logically act on the dotted line in the following diagram): |
1949 | |
1950 | <--- CPU ---> : <----------- Memory -----------> |
1951 | : |
1952 | +--------+ +--------+ : +--------+ +-----------+ |
1953 | | | | | : | | | | +--------+ |
1954 | | CPU | | Memory | : | CPU | | | | | |
1955 | | Core |--->| Access |----->| Cache |<-->| | | | |
1956 | | | | Queue | : | | | |--->| Memory | |
1957 | | | | | : | | | | | | |
1958 | +--------+ +--------+ : +--------+ | | | | |
1959 | : | Cache | +--------+ |
1960 | : | Coherency | |
1961 | : | Mechanism | +--------+ |
1962 | +--------+ +--------+ : +--------+ | | | | |
1963 | | | | | : | | | | | | |
1964 | | CPU | | Memory | : | CPU | | |--->| Device | |
1965 | | Core |--->| Access |----->| Cache |<-->| | | | |
1966 | | | | Queue | : | | | | | | |
1967 | | | | | : | | | | +--------+ |
1968 | +--------+ +--------+ : +--------+ +-----------+ |
1969 | : |
1970 | : |
1971 | |
1972 | Although any particular load or store may not actually appear outside of the |
1973 | CPU that issued it since it may have been satisfied within the CPU's own cache, |
1974 | it will still appear as if the full memory access had taken place as far as the |
1975 | other CPUs are concerned since the cache coherency mechanisms will migrate the |
1976 | cacheline over to the accessing CPU and propagate the effects upon conflict. |
1977 | |
1978 | The CPU core may execute instructions in any order it deems fit, provided the |
1979 | expected program causality appears to be maintained. Some of the instructions |
1980 | generate load and store operations which then go into the queue of memory |
1981 | accesses to be performed. The core may place these in the queue in any order |
1982 | it wishes, and continue execution until it is forced to wait for an instruction |
1983 | to complete. |
1984 | |
1985 | What memory barriers are concerned with is controlling the order in which |
1986 | accesses cross from the CPU side of things to the memory side of things, and |
1987 | the order in which the effects are perceived to happen by the other observers |
1988 | in the system. |
1989 | |
1990 | [!] Memory barriers are _not_ needed within a given CPU, as CPUs always see |
1991 | their own loads and stores as if they had happened in program order. |
1992 | |
1993 | [!] MMIO or other device accesses may bypass the cache system. This depends on |
1994 | the properties of the memory window through which devices are accessed and/or |
1995 | the use of any special device communication instructions the CPU may have. |
1996 | |
1997 | |
1998 | CACHE COHERENCY |
1999 | --------------- |
2000 | |
2001 | Life isn't quite as simple as it may appear above, however: for while the |
2002 | caches are expected to be coherent, there's no guarantee that that coherency |
2003 | will be ordered. This means that whilst changes made on one CPU will |
2004 | eventually become visible on all CPUs, there's no guarantee that they will |
2005 | become apparent in the same order on those other CPUs. |
2006 | |
2007 | |
2008 | Consider dealing with a system that has a pair of CPUs (1 & 2), each of which |
2009 | has a pair of parallel data caches (CPU 1 has A/B, and CPU 2 has C/D): |
2010 | |
2011 | : |
2012 | : +--------+ |
2013 | : +---------+ | | |
2014 | +--------+ : +--->| Cache A |<------->| | |
2015 | | | : | +---------+ | | |
2016 | | CPU 1 |<---+ | | |
2017 | | | : | +---------+ | | |
2018 | +--------+ : +--->| Cache B |<------->| | |
2019 | : +---------+ | | |
2020 | : | Memory | |
2021 | : +---------+ | System | |
2022 | +--------+ : +--->| Cache C |<------->| | |
2023 | | | : | +---------+ | | |
2024 | | CPU 2 |<---+ | | |
2025 | | | : | +---------+ | | |
2026 | +--------+ : +--->| Cache D |<------->| | |
2027 | : +---------+ | | |
2028 | : +--------+ |
2029 | : |
2030 | |
2031 | Imagine the system has the following properties: |
2032 | |
2033 | (*) an odd-numbered cache line may be in cache A, cache C or it may still be |
2034 | resident in memory; |
2035 | |
2036 | (*) an even-numbered cache line may be in cache B, cache D or it may still be |
2037 | resident in memory; |
2038 | |
2039 | (*) whilst the CPU core is interrogating one cache, the other cache may be |
2040 | making use of the bus to access the rest of the system - perhaps to |
2041 | displace a dirty cacheline or to do a speculative load; |
2042 | |
2043 | (*) each cache has a queue of operations that need to be applied to that cache |
2044 | to maintain coherency with the rest of the system; |
2045 | |
2046 | (*) the coherency queue is not flushed by normal loads to lines already |
2047 | present in the cache, even though the contents of the queue may |
2048 | potentially affect those loads. |
2049 | |
2050 | Imagine, then, that two writes are made on the first CPU, with a write barrier |
2051 | between them to guarantee that they will appear to reach that CPU's caches in |
2052 | the requisite order: |
2053 | |
2054 | CPU 1 CPU 2 COMMENT |
2055 | =============== =============== ======================================= |
2056 | u == 0, v == 1 and p == &u, q == &u |
2057 | v = 2; |
2058 | smp_wmb(); Make sure change to v is visible before |
2059 | change to p |
2060 | <A:modify v=2> v is now in cache A exclusively |
2061 | p = &v; |
2062 | <B:modify p=&v> p is now in cache B exclusively |
2063 | |
2064 | The write memory barrier forces the other CPUs in the system to perceive that |
2065 | the local CPU's caches have apparently been updated in the correct order. But |
2066 | now imagine that the second CPU wants to read those values: |
2067 | |
2068 | CPU 1 CPU 2 COMMENT |
2069 | =============== =============== ======================================= |
2070 | ... |
2071 | q = p; |
2072 | x = *q; |
2073 | |
2074 | The above pair of reads may then fail to happen in the expected order, as the |
2075 | cacheline holding p may get updated in one of the second CPU's caches whilst |
2076 | the update to the cacheline holding v is delayed in the other of the second |
2077 | CPU's caches by some other cache event: |
2078 | |
2079 | CPU 1 CPU 2 COMMENT |
2080 | =============== =============== ======================================= |
2081 | u == 0, v == 1 and p == &u, q == &u |
2082 | v = 2; |
2083 | smp_wmb(); |
2084 | <A:modify v=2> <C:busy> |
2085 | <C:queue v=2> |
2086 | p = &v; q = p; |
2087 | <D:request p> |
2088 | <B:modify p=&v> <D:commit p=&v> |
2089 | <D:read p> |
2090 | x = *q; |
2091 | <C:read *q> Reads from v before v updated in cache |
2092 | <C:unbusy> |
2093 | <C:commit v=2> |
2094 | |
2095 | Basically, whilst both cachelines will be updated on CPU 2 eventually, there's |
2096 | no guarantee that, without intervention, the order of update will be the same |
2097 | as that committed on CPU 1. |
2098 | |
2099 | |
2100 | To intervene, we need to interpolate a data dependency barrier or a read |
2101 | barrier between the loads. This will force the cache to commit its coherency |
2102 | queue before processing any further requests: |
2103 | |
2104 | CPU 1 CPU 2 COMMENT |
2105 | =============== =============== ======================================= |
2106 | u == 0, v == 1 and p == &u, q == &u |
2107 | v = 2; |
2108 | smp_wmb(); |
2109 | <A:modify v=2> <C:busy> |
2110 | <C:queue v=2> |
2111 | p = &v; q = p; |
2112 | <D:request p> |
2113 | <B:modify p=&v> <D:commit p=&v> |
2114 | <D:read p> |
2115 | smp_read_barrier_depends() |
2116 | <C:unbusy> |
2117 | <C:commit v=2> |
2118 | x = *q; |
2119 | <C:read *q> Reads from v after v updated in cache |
2120 | |
2121 | |
2122 | This sort of problem can be encountered on DEC Alpha processors as they have a |
2123 | split cache that improves performance by making better use of the data bus. |
2124 | Whilst most CPUs do imply a data dependency barrier on the read when a memory |
2125 | access depends on a read, not all do, so it may not be relied on. |
2126 | |
2127 | Other CPUs may also have split caches, but must coordinate between the various |
2128 | cachelets for normal memory accesses. The semantics of the Alpha removes the |
2129 | need for coordination in the absence of memory barriers. |
2130 | |
2131 | |
2132 | CACHE COHERENCY VS DMA |
2133 | ---------------------- |
2134 | |
2135 | Not all systems maintain cache coherency with respect to devices doing DMA. In |
2136 | such cases, a device attempting DMA may obtain stale data from RAM because |
2137 | dirty cache lines may be resident in the caches of various CPUs, and may not |
2138 | have been written back to RAM yet. To deal with this, the appropriate part of |
2139 | the kernel must flush the overlapping bits of cache on each CPU (and maybe |
2140 | invalidate them as well). |
2141 | |
2142 | In addition, the data DMA'd to RAM by a device may be overwritten by dirty |
2143 | cache lines being written back to RAM from a CPU's cache after the device has |
2144 | installed its own data, or cache lines present in the CPU's cache may simply |
2145 | obscure the fact that RAM has been updated, until at such time as the cacheline |
2146 | is discarded from the CPU's cache and reloaded. To deal with this, the |
2147 | appropriate part of the kernel must invalidate the overlapping bits of the |
2148 | cache on each CPU. |
2149 | |
2150 | See Documentation/cachetlb.txt for more information on cache management. |
2151 | |
2152 | |
2153 | CACHE COHERENCY VS MMIO |
2154 | ----------------------- |
2155 | |
2156 | Memory mapped I/O usually takes place through memory locations that are part of |
2157 | a window in the CPU's memory space that has different properties assigned than |
2158 | the usual RAM directed window. |
2159 | |
2160 | Amongst these properties is usually the fact that such accesses bypass the |
2161 | caching entirely and go directly to the device buses. This means MMIO accesses |
2162 | may, in effect, overtake accesses to cached memory that were emitted earlier. |
2163 | A memory barrier isn't sufficient in such a case, but rather the cache must be |
2164 | flushed between the cached memory write and the MMIO access if the two are in |
2165 | any way dependent. |
2166 | |
2167 | |
2168 | ========================= |
2169 | THE THINGS CPUS GET UP TO |
2170 | ========================= |
2171 | |
2172 | A programmer might take it for granted that the CPU will perform memory |
2173 | operations in exactly the order specified, so that if the CPU is, for example, |
2174 | given the following piece of code to execute: |
2175 | |
2176 | a = *A; |
2177 | *B = b; |
2178 | c = *C; |
2179 | d = *D; |
2180 | *E = e; |
2181 | |
2182 | they would then expect that the CPU will complete the memory operation for each |
2183 | instruction before moving on to the next one, leading to a definite sequence of |
2184 | operations as seen by external observers in the system: |
2185 | |
2186 | LOAD *A, STORE *B, LOAD *C, LOAD *D, STORE *E. |
2187 | |
2188 | |
2189 | Reality is, of course, much messier. With many CPUs and compilers, the above |
2190 | assumption doesn't hold because: |
2191 | |
2192 | (*) loads are more likely to need to be completed immediately to permit |
2193 | execution progress, whereas stores can often be deferred without a |
2194 | problem; |
2195 | |
2196 | (*) loads may be done speculatively, and the result discarded should it prove |
2197 | to have been unnecessary; |
2198 | |
2199 | (*) loads may be done speculatively, leading to the result having been fetched |
2200 | at the wrong time in the expected sequence of events; |
2201 | |
2202 | (*) the order of the memory accesses may be rearranged to promote better use |
2203 | of the CPU buses and caches; |
2204 | |
2205 | (*) loads and stores may be combined to improve performance when talking to |
2206 | memory or I/O hardware that can do batched accesses of adjacent locations, |
2207 | thus cutting down on transaction setup costs (memory and PCI devices may |
2208 | both be able to do this); and |
2209 | |
2210 | (*) the CPU's data cache may affect the ordering, and whilst cache-coherency |
2211 | mechanisms may alleviate this - once the store has actually hit the cache |
2212 | - there's no guarantee that the coherency management will be propagated in |
2213 | order to other CPUs. |
2214 | |
2215 | So what another CPU, say, might actually observe from the above piece of code |
2216 | is: |
2217 | |
2218 | LOAD *A, ..., LOAD {*C,*D}, STORE *E, STORE *B |
2219 | |
2220 | (Where "LOAD {*C,*D}" is a combined load) |
2221 | |
2222 | |
2223 | However, it is guaranteed that a CPU will be self-consistent: it will see its |
2224 | _own_ accesses appear to be correctly ordered, without the need for a memory |
2225 | barrier. For instance with the following code: |
2226 | |
2227 | U = *A; |
2228 | *A = V; |
2229 | *A = W; |
2230 | X = *A; |
2231 | *A = Y; |
2232 | Z = *A; |
2233 | |
2234 | and assuming no intervention by an external influence, it can be assumed that |
2235 | the final result will appear to be: |
2236 | |
2237 | U == the original value of *A |
2238 | X == W |
2239 | Z == Y |
2240 | *A == Y |
2241 | |
2242 | The code above may cause the CPU to generate the full sequence of memory |
2243 | accesses: |
2244 | |
2245 | U=LOAD *A, STORE *A=V, STORE *A=W, X=LOAD *A, STORE *A=Y, Z=LOAD *A |
2246 | |
2247 | in that order, but, without intervention, the sequence may have almost any |
2248 | combination of elements combined or discarded, provided the program's view of |
2249 | the world remains consistent. |
2250 | |
2251 | The compiler may also combine, discard or defer elements of the sequence before |
2252 | the CPU even sees them. |
2253 | |
2254 | For instance: |
2255 | |
2256 | *A = V; |
2257 | *A = W; |
2258 | |
2259 | may be reduced to: |
2260 | |
2261 | *A = W; |
2262 | |
2263 | since, without a write barrier, it can be assumed that the effect of the |
2264 | storage of V to *A is lost. Similarly: |
2265 | |
2266 | *A = Y; |
2267 | Z = *A; |
2268 | |
2269 | may, without a memory barrier, be reduced to: |
2270 | |
2271 | *A = Y; |
2272 | Z = Y; |
2273 | |
2274 | and the LOAD operation never appear outside of the CPU. |
2275 | |
2276 | |
2277 | AND THEN THERE'S THE ALPHA |
2278 | -------------------------- |
2279 | |
2280 | The DEC Alpha CPU is one of the most relaxed CPUs there is. Not only that, |
2281 | some versions of the Alpha CPU have a split data cache, permitting them to have |
2282 | two semantically-related cache lines updated at separate times. This is where |
2283 | the data dependency barrier really becomes necessary as this synchronises both |
2284 | caches with the memory coherence system, thus making it seem like pointer |
2285 | changes vs new data occur in the right order. |
2286 | |
2287 | The Alpha defines the Linux kernel's memory barrier model. |
2288 | |
2289 | See the subsection on "Cache Coherency" above. |
2290 | |
2291 | |
2292 | ============ |
2293 | EXAMPLE USES |
2294 | ============ |
2295 | |
2296 | CIRCULAR BUFFERS |
2297 | ---------------- |
2298 | |
2299 | Memory barriers can be used to implement circular buffering without the need |
2300 | of a lock to serialise the producer with the consumer. See: |
2301 | |
2302 | Documentation/circular-buffers.txt |
2303 | |
2304 | for details. |
2305 | |
2306 | |
2307 | ========== |
2308 | REFERENCES |
2309 | ========== |
2310 | |
2311 | Alpha AXP Architecture Reference Manual, Second Edition (Sites & Witek, |
2312 | Digital Press) |
2313 | Chapter 5.2: Physical Address Space Characteristics |
2314 | Chapter 5.4: Caches and Write Buffers |
2315 | Chapter 5.5: Data Sharing |
2316 | Chapter 5.6: Read/Write Ordering |
2317 | |
2318 | AMD64 Architecture Programmer's Manual Volume 2: System Programming |
2319 | Chapter 7.1: Memory-Access Ordering |
2320 | Chapter 7.4: Buffering and Combining Memory Writes |
2321 | |
2322 | IA-32 Intel Architecture Software Developer's Manual, Volume 3: |
2323 | System Programming Guide |
2324 | Chapter 7.1: Locked Atomic Operations |
2325 | Chapter 7.2: Memory Ordering |
2326 | Chapter 7.4: Serializing Instructions |
2327 | |
2328 | The SPARC Architecture Manual, Version 9 |
2329 | Chapter 8: Memory Models |
2330 | Appendix D: Formal Specification of the Memory Models |
2331 | Appendix J: Programming with the Memory Models |
2332 | |
2333 | UltraSPARC Programmer Reference Manual |
2334 | Chapter 5: Memory Accesses and Cacheability |
2335 | Chapter 15: Sparc-V9 Memory Models |
2336 | |
2337 | UltraSPARC III Cu User's Manual |
2338 | Chapter 9: Memory Models |
2339 | |
2340 | UltraSPARC IIIi Processor User's Manual |
2341 | Chapter 8: Memory Models |
2342 | |
2343 | UltraSPARC Architecture 2005 |
2344 | Chapter 9: Memory |
2345 | Appendix D: Formal Specifications of the Memory Models |
2346 | |
2347 | UltraSPARC T1 Supplement to the UltraSPARC Architecture 2005 |
2348 | Chapter 8: Memory Models |
2349 | Appendix F: Caches and Cache Coherency |
2350 | |
2351 | Solaris Internals, Core Kernel Architecture, p63-68: |
2352 | Chapter 3.3: Hardware Considerations for Locks and |
2353 | Synchronization |
2354 | |
2355 | Unix Systems for Modern Architectures, Symmetric Multiprocessing and Caching |
2356 | for Kernel Programmers: |
2357 | Chapter 13: Other Memory Models |
2358 | |
2359 | Intel Itanium Architecture Software Developer's Manual: Volume 1: |
2360 | Section 2.6: Speculation |
2361 | Section 4.4: Memory Access |
2362 |
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