Commit | Line | Data |
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108b42b4 DH |
1 | ============================ |
2 | LINUX KERNEL MEMORY BARRIERS | |
3 | ============================ | |
4 | ||
5 | By: David Howells <dhowells@redhat.com> | |
90fddabf | 6 | Paul E. McKenney <paulmck@linux.vnet.ibm.com> |
108b42b4 DH |
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. | |
670bd95e | 23 | - Read memory barriers vs load speculation. |
241e6663 | 24 | - Transitivity |
108b42b4 DH |
25 | |
26 | (*) Explicit kernel barriers. | |
27 | ||
28 | - Compiler barrier. | |
81fc6323 | 29 | - CPU memory barriers. |
108b42b4 DH |
30 | - MMIO write barrier. |
31 | ||
32 | (*) Implicit kernel memory barriers. | |
33 | ||
166bda71 | 34 | - Lock acquisition functions. |
108b42b4 | 35 | - Interrupt disabling functions. |
50fa610a | 36 | - Sleep and wake-up functions. |
108b42b4 DH |
37 | - Miscellaneous functions. |
38 | ||
166bda71 | 39 | (*) Inter-CPU acquiring barrier effects. |
108b42b4 | 40 | |
166bda71 SP |
41 | - Acquires vs memory accesses. |
42 | - Acquires vs I/O accesses. | |
108b42b4 DH |
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. | |
01e1cd6d | 64 | - Virtual Machine Guests. |
108b42b4 | 65 | |
90fddabf DH |
66 | (*) Example uses. |
67 | ||
68 | - Circular buffers. | |
69 | ||
108b42b4 DH |
70 | (*) References. |
71 | ||
72 | ||
73 | ============================ | |
74 | ABSTRACT MEMORY ACCESS MODEL | |
75 | ============================ | |
76 | ||
77 | Consider the following abstract model of the system: | |
78 | ||
79 | : : | |
80 | : : | |
81 | : : | |
82 | +-------+ : +--------+ : +-------+ | |
83 | | | : | | : | | | |
84 | | | : | | : | | | |
85 | | CPU 1 |<----->| Memory |<----->| CPU 2 | | |
86 | | | : | | : | | | |
87 | | | : | | : | | | |
88 | +-------+ : +--------+ : +-------+ | |
89 | ^ : ^ : ^ | |
90 | | : | : | | |
91 | | : | : | | |
92 | | : v : | | |
93 | | : +--------+ : | | |
94 | | : | | : | | |
95 | | : | | : | | |
96 | +---------->| Device |<----------+ | |
97 | : | | : | |
98 | : | | : | |
99 | : +--------+ : | |
100 | : : | |
101 | ||
102 | Each CPU executes a program that generates memory access operations. In the | |
103 | abstract CPU, memory operation ordering is very relaxed, and a CPU may actually | |
104 | perform the memory operations in any order it likes, provided program causality | |
105 | appears to be maintained. Similarly, the compiler may also arrange the | |
106 | instructions it emits in any order it likes, provided it doesn't affect the | |
107 | apparent operation of the program. | |
108 | ||
109 | So in the above diagram, the effects of the memory operations performed by a | |
110 | CPU are perceived by the rest of the system as the operations cross the | |
111 | interface between the CPU and rest of the system (the dotted lines). | |
112 | ||
113 | ||
114 | For example, consider the following sequence of events: | |
115 | ||
116 | CPU 1 CPU 2 | |
117 | =============== =============== | |
118 | { A == 1; B == 2 } | |
615cc2c9 AD |
119 | A = 3; x = B; |
120 | B = 4; y = A; | |
108b42b4 DH |
121 | |
122 | The set of accesses as seen by the memory system in the middle can be arranged | |
123 | in 24 different combinations: | |
124 | ||
8ab8b3e1 PK |
125 | STORE A=3, STORE B=4, y=LOAD A->3, x=LOAD B->4 |
126 | STORE A=3, STORE B=4, x=LOAD B->4, y=LOAD A->3 | |
127 | STORE A=3, y=LOAD A->3, STORE B=4, x=LOAD B->4 | |
128 | STORE A=3, y=LOAD A->3, x=LOAD B->2, STORE B=4 | |
129 | STORE A=3, x=LOAD B->2, STORE B=4, y=LOAD A->3 | |
130 | STORE A=3, x=LOAD B->2, y=LOAD A->3, STORE B=4 | |
131 | STORE B=4, STORE A=3, y=LOAD A->3, x=LOAD B->4 | |
108b42b4 DH |
132 | STORE B=4, ... |
133 | ... | |
134 | ||
135 | and can thus result in four different combinations of values: | |
136 | ||
8ab8b3e1 PK |
137 | x == 2, y == 1 |
138 | x == 2, y == 3 | |
139 | x == 4, y == 1 | |
140 | x == 4, y == 3 | |
108b42b4 DH |
141 | |
142 | ||
143 | Furthermore, the stores committed by a CPU to the memory system may not be | |
144 | perceived by the loads made by another CPU in the same order as the stores were | |
145 | committed. | |
146 | ||
147 | ||
148 | As a further example, consider this sequence of events: | |
149 | ||
150 | CPU 1 CPU 2 | |
151 | =============== =============== | |
3dbf0913 | 152 | { A == 1, B == 2, C == 3, P == &A, Q == &C } |
108b42b4 DH |
153 | B = 4; Q = P; |
154 | P = &B D = *Q; | |
155 | ||
156 | There is an obvious data dependency here, as the value loaded into D depends on | |
157 | the address retrieved from P by CPU 2. At the end of the sequence, any of the | |
158 | following results are possible: | |
159 | ||
160 | (Q == &A) and (D == 1) | |
161 | (Q == &B) and (D == 2) | |
162 | (Q == &B) and (D == 4) | |
163 | ||
164 | Note that CPU 2 will never try and load C into D because the CPU will load P | |
165 | into Q before issuing the load of *Q. | |
166 | ||
167 | ||
168 | DEVICE OPERATIONS | |
169 | ----------------- | |
170 | ||
171 | Some devices present their control interfaces as collections of memory | |
172 | locations, but the order in which the control registers are accessed is very | |
173 | important. For instance, imagine an ethernet card with a set of internal | |
174 | registers that are accessed through an address port register (A) and a data | |
175 | port register (D). To read internal register 5, the following code might then | |
176 | be used: | |
177 | ||
178 | *A = 5; | |
179 | x = *D; | |
180 | ||
181 | but this might show up as either of the following two sequences: | |
182 | ||
183 | STORE *A = 5, x = LOAD *D | |
184 | x = LOAD *D, STORE *A = 5 | |
185 | ||
186 | the second of which will almost certainly result in a malfunction, since it set | |
187 | the address _after_ attempting to read the register. | |
188 | ||
189 | ||
190 | GUARANTEES | |
191 | ---------- | |
192 | ||
193 | There are some minimal guarantees that may be expected of a CPU: | |
194 | ||
195 | (*) On any given CPU, dependent memory accesses will be issued in order, with | |
196 | respect to itself. This means that for: | |
197 | ||
f84cfbb0 | 198 | Q = READ_ONCE(P); smp_read_barrier_depends(); D = READ_ONCE(*Q); |
108b42b4 DH |
199 | |
200 | the CPU will issue the following memory operations: | |
201 | ||
202 | Q = LOAD P, D = LOAD *Q | |
203 | ||
2ecf8101 | 204 | and always in that order. On most systems, smp_read_barrier_depends() |
9af194ce | 205 | does nothing, but it is required for DEC Alpha. The READ_ONCE() |
f84cfbb0 CM |
206 | is required to prevent compiler mischief. Please note that you |
207 | should normally use something like rcu_dereference() instead of | |
208 | open-coding smp_read_barrier_depends(). | |
108b42b4 DH |
209 | |
210 | (*) Overlapping loads and stores within a particular CPU will appear to be | |
211 | ordered within that CPU. This means that for: | |
212 | ||
9af194ce | 213 | a = READ_ONCE(*X); WRITE_ONCE(*X, b); |
108b42b4 DH |
214 | |
215 | the CPU will only issue the following sequence of memory operations: | |
216 | ||
217 | a = LOAD *X, STORE *X = b | |
218 | ||
219 | And for: | |
220 | ||
9af194ce | 221 | WRITE_ONCE(*X, c); d = READ_ONCE(*X); |
108b42b4 DH |
222 | |
223 | the CPU will only issue: | |
224 | ||
225 | STORE *X = c, d = LOAD *X | |
226 | ||
fa00e7e1 | 227 | (Loads and stores overlap if they are targeted at overlapping pieces of |
108b42b4 DH |
228 | memory). |
229 | ||
230 | And there are a number of things that _must_ or _must_not_ be assumed: | |
231 | ||
9af194ce PM |
232 | (*) It _must_not_ be assumed that the compiler will do what you want |
233 | with memory references that are not protected by READ_ONCE() and | |
234 | WRITE_ONCE(). Without them, the compiler is within its rights to | |
235 | do all sorts of "creative" transformations, which are covered in | |
895f5542 | 236 | the COMPILER BARRIER section. |
2ecf8101 | 237 | |
108b42b4 DH |
238 | (*) It _must_not_ be assumed that independent loads and stores will be issued |
239 | in the order given. This means that for: | |
240 | ||
241 | X = *A; Y = *B; *D = Z; | |
242 | ||
243 | we may get any of the following sequences: | |
244 | ||
245 | X = LOAD *A, Y = LOAD *B, STORE *D = Z | |
246 | X = LOAD *A, STORE *D = Z, Y = LOAD *B | |
247 | Y = LOAD *B, X = LOAD *A, STORE *D = Z | |
248 | Y = LOAD *B, STORE *D = Z, X = LOAD *A | |
249 | STORE *D = Z, X = LOAD *A, Y = LOAD *B | |
250 | STORE *D = Z, Y = LOAD *B, X = LOAD *A | |
251 | ||
252 | (*) It _must_ be assumed that overlapping memory accesses may be merged or | |
253 | discarded. This means that for: | |
254 | ||
255 | X = *A; Y = *(A + 4); | |
256 | ||
257 | we may get any one of the following sequences: | |
258 | ||
259 | X = LOAD *A; Y = LOAD *(A + 4); | |
260 | Y = LOAD *(A + 4); X = LOAD *A; | |
261 | {X, Y} = LOAD {*A, *(A + 4) }; | |
262 | ||
263 | And for: | |
264 | ||
f191eec5 | 265 | *A = X; *(A + 4) = Y; |
108b42b4 | 266 | |
f191eec5 | 267 | we may get any of: |
108b42b4 | 268 | |
f191eec5 PM |
269 | STORE *A = X; STORE *(A + 4) = Y; |
270 | STORE *(A + 4) = Y; STORE *A = X; | |
271 | STORE {*A, *(A + 4) } = {X, Y}; | |
108b42b4 | 272 | |
432fbf3c PM |
273 | And there are anti-guarantees: |
274 | ||
275 | (*) These guarantees do not apply to bitfields, because compilers often | |
276 | generate code to modify these using non-atomic read-modify-write | |
277 | sequences. Do not attempt to use bitfields to synchronize parallel | |
278 | algorithms. | |
279 | ||
280 | (*) Even in cases where bitfields are protected by locks, all fields | |
281 | in a given bitfield must be protected by one lock. If two fields | |
282 | in a given bitfield are protected by different locks, the compiler's | |
283 | non-atomic read-modify-write sequences can cause an update to one | |
284 | field to corrupt the value of an adjacent field. | |
285 | ||
286 | (*) These guarantees apply only to properly aligned and sized scalar | |
287 | variables. "Properly sized" currently means variables that are | |
288 | the same size as "char", "short", "int" and "long". "Properly | |
289 | aligned" means the natural alignment, thus no constraints for | |
290 | "char", two-byte alignment for "short", four-byte alignment for | |
291 | "int", and either four-byte or eight-byte alignment for "long", | |
292 | on 32-bit and 64-bit systems, respectively. Note that these | |
293 | guarantees were introduced into the C11 standard, so beware when | |
294 | using older pre-C11 compilers (for example, gcc 4.6). The portion | |
295 | of the standard containing this guarantee is Section 3.14, which | |
296 | defines "memory location" as follows: | |
297 | ||
298 | memory location | |
299 | either an object of scalar type, or a maximal sequence | |
300 | of adjacent bit-fields all having nonzero width | |
301 | ||
302 | NOTE 1: Two threads of execution can update and access | |
303 | separate memory locations without interfering with | |
304 | each other. | |
305 | ||
306 | NOTE 2: A bit-field and an adjacent non-bit-field member | |
307 | are in separate memory locations. The same applies | |
308 | to two bit-fields, if one is declared inside a nested | |
309 | structure declaration and the other is not, or if the two | |
310 | are separated by a zero-length bit-field declaration, | |
311 | or if they are separated by a non-bit-field member | |
312 | declaration. It is not safe to concurrently update two | |
313 | bit-fields in the same structure if all members declared | |
314 | between them are also bit-fields, no matter what the | |
315 | sizes of those intervening bit-fields happen to be. | |
316 | ||
108b42b4 DH |
317 | |
318 | ========================= | |
319 | WHAT ARE MEMORY BARRIERS? | |
320 | ========================= | |
321 | ||
322 | As can be seen above, independent memory operations are effectively performed | |
323 | in random order, but this can be a problem for CPU-CPU interaction and for I/O. | |
324 | What is required is some way of intervening to instruct the compiler and the | |
325 | CPU to restrict the order. | |
326 | ||
327 | Memory barriers are such interventions. They impose a perceived partial | |
2b94895b DH |
328 | ordering over the memory operations on either side of the barrier. |
329 | ||
330 | Such enforcement is important because the CPUs and other devices in a system | |
81fc6323 | 331 | can use a variety of tricks to improve performance, including reordering, |
2b94895b DH |
332 | deferral and combination of memory operations; speculative loads; speculative |
333 | branch prediction and various types of caching. Memory barriers are used to | |
334 | override or suppress these tricks, allowing the code to sanely control the | |
335 | interaction of multiple CPUs and/or devices. | |
108b42b4 DH |
336 | |
337 | ||
338 | VARIETIES OF MEMORY BARRIER | |
339 | --------------------------- | |
340 | ||
341 | Memory barriers come in four basic varieties: | |
342 | ||
343 | (1) Write (or store) memory barriers. | |
344 | ||
345 | A write memory barrier gives a guarantee that all the STORE operations | |
346 | specified before the barrier will appear to happen before all the STORE | |
347 | operations specified after the barrier with respect to the other | |
348 | components of the system. | |
349 | ||
350 | A write barrier is a partial ordering on stores only; it is not required | |
351 | to have any effect on loads. | |
352 | ||
6bc39274 | 353 | A CPU can be viewed as committing a sequence of store operations to the |
108b42b4 DH |
354 | memory system as time progresses. All stores before a write barrier will |
355 | occur in the sequence _before_ all the stores after the write barrier. | |
356 | ||
357 | [!] Note that write barriers should normally be paired with read or data | |
358 | dependency barriers; see the "SMP barrier pairing" subsection. | |
359 | ||
360 | ||
361 | (2) Data dependency barriers. | |
362 | ||
363 | A data dependency barrier is a weaker form of read barrier. In the case | |
364 | where two loads are performed such that the second depends on the result | |
365 | of the first (eg: the first load retrieves the address to which the second | |
366 | load will be directed), a data dependency barrier would be required to | |
367 | make sure that the target of the second load is updated before the address | |
368 | obtained by the first load is accessed. | |
369 | ||
370 | A data dependency barrier is a partial ordering on interdependent loads | |
371 | only; it is not required to have any effect on stores, independent loads | |
372 | or overlapping loads. | |
373 | ||
374 | As mentioned in (1), the other CPUs in the system can be viewed as | |
375 | committing sequences of stores to the memory system that the CPU being | |
376 | considered can then perceive. A data dependency barrier issued by the CPU | |
377 | under consideration guarantees that for any load preceding it, if that | |
378 | load touches one of a sequence of stores from another CPU, then by the | |
379 | time the barrier completes, the effects of all the stores prior to that | |
380 | touched by the load will be perceptible to any loads issued after the data | |
381 | dependency barrier. | |
382 | ||
383 | See the "Examples of memory barrier sequences" subsection for diagrams | |
384 | showing the ordering constraints. | |
385 | ||
386 | [!] Note that the first load really has to have a _data_ dependency and | |
387 | not a control dependency. If the address for the second load is dependent | |
388 | on the first load, but the dependency is through a conditional rather than | |
389 | actually loading the address itself, then it's a _control_ dependency and | |
390 | a full read barrier or better is required. See the "Control dependencies" | |
391 | subsection for more information. | |
392 | ||
393 | [!] Note that data dependency barriers should normally be paired with | |
394 | write barriers; see the "SMP barrier pairing" subsection. | |
395 | ||
396 | ||
397 | (3) Read (or load) memory barriers. | |
398 | ||
399 | A read barrier is a data dependency barrier plus a guarantee that all the | |
400 | LOAD operations specified before the barrier will appear to happen before | |
401 | all the LOAD operations specified after the barrier with respect to the | |
402 | other components of the system. | |
403 | ||
404 | A read barrier is a partial ordering on loads only; it is not required to | |
405 | have any effect on stores. | |
406 | ||
407 | Read memory barriers imply data dependency barriers, and so can substitute | |
408 | for them. | |
409 | ||
410 | [!] Note that read barriers should normally be paired with write barriers; | |
411 | see the "SMP barrier pairing" subsection. | |
412 | ||
413 | ||
414 | (4) General memory barriers. | |
415 | ||
670bd95e DH |
416 | A general memory barrier gives a guarantee that all the LOAD and STORE |
417 | operations specified before the barrier will appear to happen before all | |
418 | the LOAD and STORE operations specified after the barrier with respect to | |
419 | the other components of the system. | |
420 | ||
421 | A general memory barrier is a partial ordering over both loads and stores. | |
108b42b4 DH |
422 | |
423 | General memory barriers imply both read and write memory barriers, and so | |
424 | can substitute for either. | |
425 | ||
426 | ||
427 | And a couple of implicit varieties: | |
428 | ||
2e4f5382 | 429 | (5) ACQUIRE operations. |
108b42b4 DH |
430 | |
431 | This acts as a one-way permeable barrier. It guarantees that all memory | |
2e4f5382 PZ |
432 | operations after the ACQUIRE operation will appear to happen after the |
433 | ACQUIRE operation with respect to the other components of the system. | |
787df638 DB |
434 | ACQUIRE operations include LOCK operations and both smp_load_acquire() |
435 | and smp_cond_acquire() operations. The later builds the necessary ACQUIRE | |
436 | semantics from relying on a control dependency and smp_rmb(). | |
108b42b4 | 437 | |
2e4f5382 PZ |
438 | Memory operations that occur before an ACQUIRE operation may appear to |
439 | happen after it completes. | |
108b42b4 | 440 | |
2e4f5382 PZ |
441 | An ACQUIRE operation should almost always be paired with a RELEASE |
442 | operation. | |
108b42b4 DH |
443 | |
444 | ||
2e4f5382 | 445 | (6) RELEASE operations. |
108b42b4 DH |
446 | |
447 | This also acts as a one-way permeable barrier. It guarantees that all | |
2e4f5382 PZ |
448 | memory operations before the RELEASE operation will appear to happen |
449 | before the RELEASE operation with respect to the other components of the | |
450 | system. RELEASE operations include UNLOCK operations and | |
451 | smp_store_release() operations. | |
108b42b4 | 452 | |
2e4f5382 | 453 | Memory operations that occur after a RELEASE operation may appear to |
108b42b4 DH |
454 | happen before it completes. |
455 | ||
2e4f5382 PZ |
456 | The use of ACQUIRE and RELEASE operations generally precludes the need |
457 | for other sorts of memory barrier (but note the exceptions mentioned in | |
458 | the subsection "MMIO write barrier"). In addition, a RELEASE+ACQUIRE | |
459 | pair is -not- guaranteed to act as a full memory barrier. However, after | |
460 | an ACQUIRE on a given variable, all memory accesses preceding any prior | |
461 | RELEASE on that same variable are guaranteed to be visible. In other | |
462 | words, within a given variable's critical section, all accesses of all | |
463 | previous critical sections for that variable are guaranteed to have | |
464 | completed. | |
17eb88e0 | 465 | |
2e4f5382 PZ |
466 | This means that ACQUIRE acts as a minimal "acquire" operation and |
467 | RELEASE acts as a minimal "release" operation. | |
108b42b4 DH |
468 | |
469 | ||
470 | Memory barriers are only required where there's a possibility of interaction | |
471 | between two CPUs or between a CPU and a device. If it can be guaranteed that | |
472 | there won't be any such interaction in any particular piece of code, then | |
473 | memory barriers are unnecessary in that piece of code. | |
474 | ||
475 | ||
476 | Note that these are the _minimum_ guarantees. Different architectures may give | |
477 | more substantial guarantees, but they may _not_ be relied upon outside of arch | |
478 | specific code. | |
479 | ||
480 | ||
481 | WHAT MAY NOT BE ASSUMED ABOUT MEMORY BARRIERS? | |
482 | ---------------------------------------------- | |
483 | ||
484 | There are certain things that the Linux kernel memory barriers do not guarantee: | |
485 | ||
486 | (*) There is no guarantee that any of the memory accesses specified before a | |
487 | memory barrier will be _complete_ by the completion of a memory barrier | |
488 | instruction; the barrier can be considered to draw a line in that CPU's | |
489 | access queue that accesses of the appropriate type may not cross. | |
490 | ||
491 | (*) There is no guarantee that issuing a memory barrier on one CPU will have | |
492 | any direct effect on another CPU or any other hardware in the system. The | |
493 | indirect effect will be the order in which the second CPU sees the effects | |
494 | of the first CPU's accesses occur, but see the next point: | |
495 | ||
6bc39274 | 496 | (*) There is no guarantee that a CPU will see the correct order of effects |
108b42b4 DH |
497 | from a second CPU's accesses, even _if_ the second CPU uses a memory |
498 | barrier, unless the first CPU _also_ uses a matching memory barrier (see | |
499 | the subsection on "SMP Barrier Pairing"). | |
500 | ||
501 | (*) There is no guarantee that some intervening piece of off-the-CPU | |
502 | hardware[*] will not reorder the memory accesses. CPU cache coherency | |
503 | mechanisms should propagate the indirect effects of a memory barrier | |
504 | between CPUs, but might not do so in order. | |
505 | ||
506 | [*] For information on bus mastering DMA and coherency please read: | |
507 | ||
4b5ff469 | 508 | Documentation/PCI/pci.txt |
395cf969 | 509 | Documentation/DMA-API-HOWTO.txt |
108b42b4 DH |
510 | Documentation/DMA-API.txt |
511 | ||
512 | ||
513 | DATA DEPENDENCY BARRIERS | |
514 | ------------------------ | |
515 | ||
516 | The usage requirements of data dependency barriers are a little subtle, and | |
517 | it's not always obvious that they're needed. To illustrate, consider the | |
518 | following sequence of events: | |
519 | ||
2ecf8101 PM |
520 | CPU 1 CPU 2 |
521 | =============== =============== | |
3dbf0913 | 522 | { A == 1, B == 2, C == 3, P == &A, Q == &C } |
108b42b4 DH |
523 | B = 4; |
524 | <write barrier> | |
9af194ce PM |
525 | WRITE_ONCE(P, &B) |
526 | Q = READ_ONCE(P); | |
2ecf8101 | 527 | D = *Q; |
108b42b4 DH |
528 | |
529 | There's a clear data dependency here, and it would seem that by the end of the | |
530 | sequence, Q must be either &A or &B, and that: | |
531 | ||
532 | (Q == &A) implies (D == 1) | |
533 | (Q == &B) implies (D == 4) | |
534 | ||
81fc6323 | 535 | But! CPU 2's perception of P may be updated _before_ its perception of B, thus |
108b42b4 DH |
536 | leading to the following situation: |
537 | ||
538 | (Q == &B) and (D == 2) ???? | |
539 | ||
540 | Whilst this may seem like a failure of coherency or causality maintenance, it | |
541 | isn't, and this behaviour can be observed on certain real CPUs (such as the DEC | |
542 | Alpha). | |
543 | ||
2b94895b DH |
544 | To deal with this, a data dependency barrier or better must be inserted |
545 | between the address load and the data load: | |
108b42b4 | 546 | |
2ecf8101 PM |
547 | CPU 1 CPU 2 |
548 | =============== =============== | |
3dbf0913 | 549 | { A == 1, B == 2, C == 3, P == &A, Q == &C } |
108b42b4 DH |
550 | B = 4; |
551 | <write barrier> | |
9af194ce PM |
552 | WRITE_ONCE(P, &B); |
553 | Q = READ_ONCE(P); | |
2ecf8101 PM |
554 | <data dependency barrier> |
555 | D = *Q; | |
108b42b4 DH |
556 | |
557 | This enforces the occurrence of one of the two implications, and prevents the | |
558 | third possibility from arising. | |
559 | ||
92a84dd2 PM |
560 | A data-dependency barrier must also order against dependent writes: |
561 | ||
562 | CPU 1 CPU 2 | |
563 | =============== =============== | |
564 | { A == 1, B == 2, C = 3, P == &A, Q == &C } | |
565 | B = 4; | |
566 | <write barrier> | |
567 | WRITE_ONCE(P, &B); | |
568 | Q = READ_ONCE(P); | |
569 | <data dependency barrier> | |
570 | *Q = 5; | |
571 | ||
572 | The data-dependency barrier must order the read into Q with the store | |
573 | into *Q. This prohibits this outcome: | |
574 | ||
575 | (Q == B) && (B == 4) | |
576 | ||
577 | Please note that this pattern should be rare. After all, the whole point | |
578 | of dependency ordering is to -prevent- writes to the data structure, along | |
579 | with the expensive cache misses associated with those writes. This pattern | |
580 | can be used to record rare error conditions and the like, and the ordering | |
581 | prevents such records from being lost. | |
582 | ||
583 | ||
108b42b4 DH |
584 | [!] Note that this extremely counterintuitive situation arises most easily on |
585 | machines with split caches, so that, for example, one cache bank processes | |
586 | even-numbered cache lines and the other bank processes odd-numbered cache | |
587 | lines. The pointer P might be stored in an odd-numbered cache line, and the | |
588 | variable B might be stored in an even-numbered cache line. Then, if the | |
589 | even-numbered bank of the reading CPU's cache is extremely busy while the | |
590 | odd-numbered bank is idle, one can see the new value of the pointer P (&B), | |
6bc39274 | 591 | but the old value of the variable B (2). |
108b42b4 DH |
592 | |
593 | ||
2ecf8101 PM |
594 | The data dependency barrier is very important to the RCU system, |
595 | for example. See rcu_assign_pointer() and rcu_dereference() in | |
596 | include/linux/rcupdate.h. This permits the current target of an RCU'd | |
597 | pointer to be replaced with a new modified target, without the replacement | |
598 | target appearing to be incompletely initialised. | |
108b42b4 DH |
599 | |
600 | See also the subsection on "Cache Coherency" for a more thorough example. | |
601 | ||
602 | ||
603 | CONTROL DEPENDENCIES | |
604 | -------------------- | |
605 | ||
ff382810 PM |
606 | A load-load control dependency requires a full read memory barrier, not |
607 | simply a data dependency barrier to make it work correctly. Consider the | |
608 | following bit of code: | |
108b42b4 | 609 | |
9af194ce | 610 | q = READ_ONCE(a); |
18c03c61 PZ |
611 | if (q) { |
612 | <data dependency barrier> /* BUG: No data dependency!!! */ | |
9af194ce | 613 | p = READ_ONCE(b); |
45c8a36a | 614 | } |
108b42b4 DH |
615 | |
616 | This will not have the desired effect because there is no actual data | |
2ecf8101 PM |
617 | dependency, but rather a control dependency that the CPU may short-circuit |
618 | by attempting to predict the outcome in advance, so that other CPUs see | |
619 | the load from b as having happened before the load from a. In such a | |
620 | case what's actually required is: | |
108b42b4 | 621 | |
9af194ce | 622 | q = READ_ONCE(a); |
18c03c61 | 623 | if (q) { |
45c8a36a | 624 | <read barrier> |
9af194ce | 625 | p = READ_ONCE(b); |
45c8a36a | 626 | } |
18c03c61 PZ |
627 | |
628 | However, stores are not speculated. This means that ordering -is- provided | |
ff382810 | 629 | for load-store control dependencies, as in the following example: |
18c03c61 | 630 | |
105ff3cb | 631 | q = READ_ONCE(a); |
18c03c61 | 632 | if (q) { |
9af194ce | 633 | WRITE_ONCE(b, p); |
18c03c61 PZ |
634 | } |
635 | ||
5af4692a | 636 | Control dependencies pair normally with other types of barriers. That |
105ff3cb LT |
637 | said, please note that READ_ONCE() is not optional! Without the |
638 | READ_ONCE(), the compiler might combine the load from 'a' with other | |
639 | loads from 'a', and the store to 'b' with other stores to 'b', with | |
640 | possible highly counterintuitive effects on ordering. | |
18c03c61 PZ |
641 | |
642 | Worse yet, if the compiler is able to prove (say) that the value of | |
643 | variable 'a' is always non-zero, it would be well within its rights | |
644 | to optimize the original example by eliminating the "if" statement | |
645 | as follows: | |
646 | ||
647 | q = a; | |
2456d2a6 PM |
648 | b = p; /* BUG: Compiler and CPU can both reorder!!! */ |
649 | ||
105ff3cb | 650 | So don't leave out the READ_ONCE(). |
18c03c61 | 651 | |
2456d2a6 PM |
652 | It is tempting to try to enforce ordering on identical stores on both |
653 | branches of the "if" statement as follows: | |
18c03c61 | 654 | |
105ff3cb | 655 | q = READ_ONCE(a); |
18c03c61 | 656 | if (q) { |
9b2b3bf5 | 657 | barrier(); |
9af194ce | 658 | WRITE_ONCE(b, p); |
18c03c61 PZ |
659 | do_something(); |
660 | } else { | |
9b2b3bf5 | 661 | barrier(); |
9af194ce | 662 | WRITE_ONCE(b, p); |
18c03c61 PZ |
663 | do_something_else(); |
664 | } | |
665 | ||
2456d2a6 PM |
666 | Unfortunately, current compilers will transform this as follows at high |
667 | optimization levels: | |
18c03c61 | 668 | |
105ff3cb | 669 | q = READ_ONCE(a); |
2456d2a6 | 670 | barrier(); |
9af194ce | 671 | WRITE_ONCE(b, p); /* BUG: No ordering vs. load from a!!! */ |
18c03c61 | 672 | if (q) { |
9af194ce | 673 | /* WRITE_ONCE(b, p); -- moved up, BUG!!! */ |
18c03c61 PZ |
674 | do_something(); |
675 | } else { | |
9af194ce | 676 | /* WRITE_ONCE(b, p); -- moved up, BUG!!! */ |
18c03c61 PZ |
677 | do_something_else(); |
678 | } | |
679 | ||
2456d2a6 PM |
680 | Now there is no conditional between the load from 'a' and the store to |
681 | 'b', which means that the CPU is within its rights to reorder them: | |
682 | The conditional is absolutely required, and must be present in the | |
683 | assembly code even after all compiler optimizations have been applied. | |
684 | Therefore, if you need ordering in this example, you need explicit | |
685 | memory barriers, for example, smp_store_release(): | |
18c03c61 | 686 | |
9af194ce | 687 | q = READ_ONCE(a); |
2456d2a6 PM |
688 | if (q) { |
689 | smp_store_release(&b, p); | |
18c03c61 PZ |
690 | do_something(); |
691 | } else { | |
2456d2a6 | 692 | smp_store_release(&b, p); |
18c03c61 PZ |
693 | do_something_else(); |
694 | } | |
695 | ||
2456d2a6 PM |
696 | In contrast, without explicit memory barriers, two-legged-if control |
697 | ordering is guaranteed only when the stores differ, for example: | |
698 | ||
105ff3cb | 699 | q = READ_ONCE(a); |
2456d2a6 | 700 | if (q) { |
9af194ce | 701 | WRITE_ONCE(b, p); |
2456d2a6 PM |
702 | do_something(); |
703 | } else { | |
9af194ce | 704 | WRITE_ONCE(b, r); |
2456d2a6 PM |
705 | do_something_else(); |
706 | } | |
707 | ||
105ff3cb LT |
708 | The initial READ_ONCE() is still required to prevent the compiler from |
709 | proving the value of 'a'. | |
18c03c61 PZ |
710 | |
711 | In addition, you need to be careful what you do with the local variable 'q', | |
712 | otherwise the compiler might be able to guess the value and again remove | |
713 | the needed conditional. For example: | |
714 | ||
105ff3cb | 715 | q = READ_ONCE(a); |
18c03c61 | 716 | if (q % MAX) { |
9af194ce | 717 | WRITE_ONCE(b, p); |
18c03c61 PZ |
718 | do_something(); |
719 | } else { | |
9af194ce | 720 | WRITE_ONCE(b, r); |
18c03c61 PZ |
721 | do_something_else(); |
722 | } | |
723 | ||
724 | If MAX is defined to be 1, then the compiler knows that (q % MAX) is | |
725 | equal to zero, in which case the compiler is within its rights to | |
726 | transform the above code into the following: | |
727 | ||
105ff3cb | 728 | q = READ_ONCE(a); |
9af194ce | 729 | WRITE_ONCE(b, p); |
18c03c61 PZ |
730 | do_something_else(); |
731 | ||
2456d2a6 PM |
732 | Given this transformation, the CPU is not required to respect the ordering |
733 | between the load from variable 'a' and the store to variable 'b'. It is | |
734 | tempting to add a barrier(), but this does not help. The conditional | |
735 | is gone, and the barrier won't bring it back. Therefore, if you are | |
736 | relying on this ordering, you should make sure that MAX is greater than | |
737 | one, perhaps as follows: | |
18c03c61 | 738 | |
105ff3cb | 739 | q = READ_ONCE(a); |
18c03c61 PZ |
740 | BUILD_BUG_ON(MAX <= 1); /* Order load from a with store to b. */ |
741 | if (q % MAX) { | |
9af194ce | 742 | WRITE_ONCE(b, p); |
18c03c61 PZ |
743 | do_something(); |
744 | } else { | |
9af194ce | 745 | WRITE_ONCE(b, r); |
18c03c61 PZ |
746 | do_something_else(); |
747 | } | |
748 | ||
2456d2a6 PM |
749 | Please note once again that the stores to 'b' differ. If they were |
750 | identical, as noted earlier, the compiler could pull this store outside | |
751 | of the 'if' statement. | |
752 | ||
8b19d1de PM |
753 | You must also be careful not to rely too much on boolean short-circuit |
754 | evaluation. Consider this example: | |
755 | ||
105ff3cb | 756 | q = READ_ONCE(a); |
57aecae9 | 757 | if (q || 1 > 0) |
9af194ce | 758 | WRITE_ONCE(b, 1); |
8b19d1de | 759 | |
5af4692a PM |
760 | Because the first condition cannot fault and the second condition is |
761 | always true, the compiler can transform this example as following, | |
762 | defeating control dependency: | |
8b19d1de | 763 | |
105ff3cb | 764 | q = READ_ONCE(a); |
9af194ce | 765 | WRITE_ONCE(b, 1); |
8b19d1de PM |
766 | |
767 | This example underscores the need to ensure that the compiler cannot | |
9af194ce | 768 | out-guess your code. More generally, although READ_ONCE() does force |
8b19d1de PM |
769 | the compiler to actually emit code for a given load, it does not force |
770 | the compiler to use the results. | |
771 | ||
18c03c61 | 772 | Finally, control dependencies do -not- provide transitivity. This is |
5646f7ac PM |
773 | demonstrated by two related examples, with the initial values of |
774 | x and y both being zero: | |
18c03c61 PZ |
775 | |
776 | CPU 0 CPU 1 | |
5af4692a | 777 | ======================= ======================= |
105ff3cb | 778 | r1 = READ_ONCE(x); r2 = READ_ONCE(y); |
5646f7ac | 779 | if (r1 > 0) if (r2 > 0) |
9af194ce | 780 | WRITE_ONCE(y, 1); WRITE_ONCE(x, 1); |
18c03c61 PZ |
781 | |
782 | assert(!(r1 == 1 && r2 == 1)); | |
783 | ||
784 | The above two-CPU example will never trigger the assert(). However, | |
785 | if control dependencies guaranteed transitivity (which they do not), | |
5646f7ac | 786 | then adding the following CPU would guarantee a related assertion: |
18c03c61 | 787 | |
5646f7ac PM |
788 | CPU 2 |
789 | ===================== | |
9af194ce | 790 | WRITE_ONCE(x, 2); |
5646f7ac PM |
791 | |
792 | assert(!(r1 == 2 && r2 == 1 && x == 2)); /* FAILS!!! */ | |
18c03c61 | 793 | |
5646f7ac PM |
794 | But because control dependencies do -not- provide transitivity, the above |
795 | assertion can fail after the combined three-CPU example completes. If you | |
796 | need the three-CPU example to provide ordering, you will need smp_mb() | |
797 | between the loads and stores in the CPU 0 and CPU 1 code fragments, | |
5af4692a PM |
798 | that is, just before or just after the "if" statements. Furthermore, |
799 | the original two-CPU example is very fragile and should be avoided. | |
18c03c61 | 800 | |
5646f7ac PM |
801 | These two examples are the LB and WWC litmus tests from this paper: |
802 | http://www.cl.cam.ac.uk/users/pes20/ppc-supplemental/test6.pdf and this | |
803 | site: https://www.cl.cam.ac.uk/~pes20/ppcmem/index.html. | |
18c03c61 PZ |
804 | |
805 | In summary: | |
806 | ||
807 | (*) Control dependencies can order prior loads against later stores. | |
808 | However, they do -not- guarantee any other sort of ordering: | |
809 | Not prior loads against later loads, nor prior stores against | |
810 | later anything. If you need these other forms of ordering, | |
d87510c5 | 811 | use smp_rmb(), smp_wmb(), or, in the case of prior stores and |
18c03c61 PZ |
812 | later loads, smp_mb(). |
813 | ||
7817b799 PM |
814 | (*) If both legs of the "if" statement begin with identical stores to |
815 | the same variable, then those stores must be ordered, either by | |
816 | preceding both of them with smp_mb() or by using smp_store_release() | |
817 | to carry out the stores. Please note that it is -not- sufficient | |
a5052657 PM |
818 | to use barrier() at beginning of each leg of the "if" statement |
819 | because, as shown by the example above, optimizing compilers can | |
820 | destroy the control dependency while respecting the letter of the | |
821 | barrier() law. | |
9b2b3bf5 | 822 | |
18c03c61 | 823 | (*) Control dependencies require at least one run-time conditional |
586dd56a | 824 | between the prior load and the subsequent store, and this |
9af194ce PM |
825 | conditional must involve the prior load. If the compiler is able |
826 | to optimize the conditional away, it will have also optimized | |
105ff3cb LT |
827 | away the ordering. Careful use of READ_ONCE() and WRITE_ONCE() |
828 | can help to preserve the needed conditional. | |
18c03c61 PZ |
829 | |
830 | (*) Control dependencies require that the compiler avoid reordering the | |
105ff3cb LT |
831 | dependency into nonexistence. Careful use of READ_ONCE() or |
832 | atomic{,64}_read() can help to preserve your control dependency. | |
895f5542 | 833 | Please see the COMPILER BARRIER section for more information. |
18c03c61 | 834 | |
ff382810 PM |
835 | (*) Control dependencies pair normally with other types of barriers. |
836 | ||
18c03c61 PZ |
837 | (*) Control dependencies do -not- provide transitivity. If you |
838 | need transitivity, use smp_mb(). | |
108b42b4 DH |
839 | |
840 | ||
841 | SMP BARRIER PAIRING | |
842 | ------------------- | |
843 | ||
844 | When dealing with CPU-CPU interactions, certain types of memory barrier should | |
845 | always be paired. A lack of appropriate pairing is almost certainly an error. | |
846 | ||
ff382810 PM |
847 | General barriers pair with each other, though they also pair with most |
848 | other types of barriers, albeit without transitivity. An acquire barrier | |
849 | pairs with a release barrier, but both may also pair with other barriers, | |
850 | including of course general barriers. A write barrier pairs with a data | |
851 | dependency barrier, a control dependency, an acquire barrier, a release | |
852 | barrier, a read barrier, or a general barrier. Similarly a read barrier, | |
853 | control dependency, or a data dependency barrier pairs with a write | |
854 | barrier, an acquire barrier, a release barrier, or a general barrier: | |
108b42b4 | 855 | |
2ecf8101 PM |
856 | CPU 1 CPU 2 |
857 | =============== =============== | |
9af194ce | 858 | WRITE_ONCE(a, 1); |
108b42b4 | 859 | <write barrier> |
9af194ce | 860 | WRITE_ONCE(b, 2); x = READ_ONCE(b); |
2ecf8101 | 861 | <read barrier> |
9af194ce | 862 | y = READ_ONCE(a); |
108b42b4 DH |
863 | |
864 | Or: | |
865 | ||
2ecf8101 PM |
866 | CPU 1 CPU 2 |
867 | =============== =============================== | |
108b42b4 DH |
868 | a = 1; |
869 | <write barrier> | |
9af194ce | 870 | WRITE_ONCE(b, &a); x = READ_ONCE(b); |
2ecf8101 PM |
871 | <data dependency barrier> |
872 | y = *x; | |
108b42b4 | 873 | |
ff382810 PM |
874 | Or even: |
875 | ||
876 | CPU 1 CPU 2 | |
877 | =============== =============================== | |
9af194ce | 878 | r1 = READ_ONCE(y); |
ff382810 | 879 | <general barrier> |
9af194ce | 880 | WRITE_ONCE(y, 1); if (r2 = READ_ONCE(x)) { |
ff382810 | 881 | <implicit control dependency> |
9af194ce | 882 | WRITE_ONCE(y, 1); |
ff382810 PM |
883 | } |
884 | ||
885 | assert(r1 == 0 || r2 == 0); | |
886 | ||
108b42b4 DH |
887 | Basically, the read barrier always has to be there, even though it can be of |
888 | the "weaker" type. | |
889 | ||
670bd95e | 890 | [!] Note that the stores before the write barrier would normally be expected to |
81fc6323 | 891 | match the loads after the read barrier or the data dependency barrier, and vice |
670bd95e DH |
892 | versa: |
893 | ||
2ecf8101 PM |
894 | CPU 1 CPU 2 |
895 | =================== =================== | |
9af194ce PM |
896 | WRITE_ONCE(a, 1); }---- --->{ v = READ_ONCE(c); |
897 | WRITE_ONCE(b, 2); } \ / { w = READ_ONCE(d); | |
2ecf8101 | 898 | <write barrier> \ <read barrier> |
9af194ce PM |
899 | WRITE_ONCE(c, 3); } / \ { x = READ_ONCE(a); |
900 | WRITE_ONCE(d, 4); }---- --->{ y = READ_ONCE(b); | |
670bd95e | 901 | |
108b42b4 DH |
902 | |
903 | EXAMPLES OF MEMORY BARRIER SEQUENCES | |
904 | ------------------------------------ | |
905 | ||
81fc6323 | 906 | Firstly, write barriers act as partial orderings on store operations. |
108b42b4 DH |
907 | Consider the following sequence of events: |
908 | ||
909 | CPU 1 | |
910 | ======================= | |
911 | STORE A = 1 | |
912 | STORE B = 2 | |
913 | STORE C = 3 | |
914 | <write barrier> | |
915 | STORE D = 4 | |
916 | STORE E = 5 | |
917 | ||
918 | This sequence of events is committed to the memory coherence system in an order | |
919 | that the rest of the system might perceive as the unordered set of { STORE A, | |
80f7228b | 920 | STORE B, STORE C } all occurring before the unordered set of { STORE D, STORE E |
108b42b4 DH |
921 | }: |
922 | ||
923 | +-------+ : : | |
924 | | | +------+ | |
925 | | |------>| C=3 | } /\ | |
81fc6323 JP |
926 | | | : +------+ }----- \ -----> Events perceptible to |
927 | | | : | A=1 | } \/ the rest of the system | |
108b42b4 DH |
928 | | | : +------+ } |
929 | | CPU 1 | : | B=2 | } | |
930 | | | +------+ } | |
931 | | | wwwwwwwwwwwwwwww } <--- At this point the write barrier | |
932 | | | +------+ } requires all stores prior to the | |
933 | | | : | E=5 | } barrier to be committed before | |
81fc6323 | 934 | | | : +------+ } further stores may take place |
108b42b4 DH |
935 | | |------>| D=4 | } |
936 | | | +------+ | |
937 | +-------+ : : | |
938 | | | |
670bd95e DH |
939 | | Sequence in which stores are committed to the |
940 | | memory system by CPU 1 | |
108b42b4 DH |
941 | V |
942 | ||
943 | ||
81fc6323 | 944 | Secondly, data dependency barriers act as partial orderings on data-dependent |
108b42b4 DH |
945 | loads. Consider the following sequence of events: |
946 | ||
947 | CPU 1 CPU 2 | |
948 | ======================= ======================= | |
c14038c3 | 949 | { B = 7; X = 9; Y = 8; C = &Y } |
108b42b4 DH |
950 | STORE A = 1 |
951 | STORE B = 2 | |
952 | <write barrier> | |
953 | STORE C = &B LOAD X | |
954 | STORE D = 4 LOAD C (gets &B) | |
955 | LOAD *C (reads B) | |
956 | ||
957 | Without intervention, CPU 2 may perceive the events on CPU 1 in some | |
958 | effectively random order, despite the write barrier issued by CPU 1: | |
959 | ||
960 | +-------+ : : : : | |
961 | | | +------+ +-------+ | Sequence of update | |
962 | | |------>| B=2 |----- --->| Y->8 | | of perception on | |
963 | | | : +------+ \ +-------+ | CPU 2 | |
964 | | CPU 1 | : | A=1 | \ --->| C->&Y | V | |
965 | | | +------+ | +-------+ | |
966 | | | wwwwwwwwwwwwwwww | : : | |
967 | | | +------+ | : : | |
968 | | | : | C=&B |--- | : : +-------+ | |
969 | | | : +------+ \ | +-------+ | | | |
970 | | |------>| D=4 | ----------->| C->&B |------>| | | |
971 | | | +------+ | +-------+ | | | |
972 | +-------+ : : | : : | | | |
973 | | : : | | | |
974 | | : : | CPU 2 | | |
975 | | +-------+ | | | |
976 | Apparently incorrect ---> | | B->7 |------>| | | |
977 | perception of B (!) | +-------+ | | | |
978 | | : : | | | |
979 | | +-------+ | | | |
980 | The load of X holds ---> \ | X->9 |------>| | | |
981 | up the maintenance \ +-------+ | | | |
982 | of coherence of B ----->| B->2 | +-------+ | |
983 | +-------+ | |
984 | : : | |
985 | ||
986 | ||
987 | In the above example, CPU 2 perceives that B is 7, despite the load of *C | |
670e9f34 | 988 | (which would be B) coming after the LOAD of C. |
108b42b4 DH |
989 | |
990 | If, however, a data dependency barrier were to be placed between the load of C | |
c14038c3 DH |
991 | and the load of *C (ie: B) on CPU 2: |
992 | ||
993 | CPU 1 CPU 2 | |
994 | ======================= ======================= | |
995 | { B = 7; X = 9; Y = 8; C = &Y } | |
996 | STORE A = 1 | |
997 | STORE B = 2 | |
998 | <write barrier> | |
999 | STORE C = &B LOAD X | |
1000 | STORE D = 4 LOAD C (gets &B) | |
1001 | <data dependency barrier> | |
1002 | LOAD *C (reads B) | |
1003 | ||
1004 | then the following will occur: | |
108b42b4 DH |
1005 | |
1006 | +-------+ : : : : | |
1007 | | | +------+ +-------+ | |
1008 | | |------>| B=2 |----- --->| Y->8 | | |
1009 | | | : +------+ \ +-------+ | |
1010 | | CPU 1 | : | A=1 | \ --->| C->&Y | | |
1011 | | | +------+ | +-------+ | |
1012 | | | wwwwwwwwwwwwwwww | : : | |
1013 | | | +------+ | : : | |
1014 | | | : | C=&B |--- | : : +-------+ | |
1015 | | | : +------+ \ | +-------+ | | | |
1016 | | |------>| D=4 | ----------->| C->&B |------>| | | |
1017 | | | +------+ | +-------+ | | | |
1018 | +-------+ : : | : : | | | |
1019 | | : : | | | |
1020 | | : : | CPU 2 | | |
1021 | | +-------+ | | | |
670bd95e DH |
1022 | | | X->9 |------>| | |
1023 | | +-------+ | | | |
1024 | Makes sure all effects ---> \ ddddddddddddddddd | | | |
1025 | prior to the store of C \ +-------+ | | | |
1026 | are perceptible to ----->| B->2 |------>| | | |
1027 | subsequent loads +-------+ | | | |
108b42b4 DH |
1028 | : : +-------+ |
1029 | ||
1030 | ||
1031 | And thirdly, a read barrier acts as a partial order on loads. Consider the | |
1032 | following sequence of events: | |
1033 | ||
1034 | CPU 1 CPU 2 | |
1035 | ======================= ======================= | |
670bd95e | 1036 | { A = 0, B = 9 } |
108b42b4 | 1037 | STORE A=1 |
108b42b4 | 1038 | <write barrier> |
670bd95e | 1039 | STORE B=2 |
108b42b4 | 1040 | LOAD B |
670bd95e | 1041 | LOAD A |
108b42b4 DH |
1042 | |
1043 | Without intervention, CPU 2 may then choose to perceive the events on CPU 1 in | |
1044 | some effectively random order, despite the write barrier issued by CPU 1: | |
1045 | ||
670bd95e DH |
1046 | +-------+ : : : : |
1047 | | | +------+ +-------+ | |
1048 | | |------>| A=1 |------ --->| A->0 | | |
1049 | | | +------+ \ +-------+ | |
1050 | | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 | | |
1051 | | | +------+ | +-------+ | |
1052 | | |------>| B=2 |--- | : : | |
1053 | | | +------+ \ | : : +-------+ | |
1054 | +-------+ : : \ | +-------+ | | | |
1055 | ---------->| B->2 |------>| | | |
1056 | | +-------+ | CPU 2 | | |
1057 | | | A->0 |------>| | | |
1058 | | +-------+ | | | |
1059 | | : : +-------+ | |
1060 | \ : : | |
1061 | \ +-------+ | |
1062 | ---->| A->1 | | |
1063 | +-------+ | |
1064 | : : | |
108b42b4 | 1065 | |
670bd95e | 1066 | |
6bc39274 | 1067 | If, however, a read barrier were to be placed between the load of B and the |
670bd95e DH |
1068 | load of A on CPU 2: |
1069 | ||
1070 | CPU 1 CPU 2 | |
1071 | ======================= ======================= | |
1072 | { A = 0, B = 9 } | |
1073 | STORE A=1 | |
1074 | <write barrier> | |
1075 | STORE B=2 | |
1076 | LOAD B | |
1077 | <read barrier> | |
1078 | LOAD A | |
1079 | ||
1080 | then the partial ordering imposed by CPU 1 will be perceived correctly by CPU | |
1081 | 2: | |
1082 | ||
1083 | +-------+ : : : : | |
1084 | | | +------+ +-------+ | |
1085 | | |------>| A=1 |------ --->| A->0 | | |
1086 | | | +------+ \ +-------+ | |
1087 | | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 | | |
1088 | | | +------+ | +-------+ | |
1089 | | |------>| B=2 |--- | : : | |
1090 | | | +------+ \ | : : +-------+ | |
1091 | +-------+ : : \ | +-------+ | | | |
1092 | ---------->| B->2 |------>| | | |
1093 | | +-------+ | CPU 2 | | |
1094 | | : : | | | |
1095 | | : : | | | |
1096 | At this point the read ----> \ rrrrrrrrrrrrrrrrr | | | |
1097 | barrier causes all effects \ +-------+ | | | |
1098 | prior to the storage of B ---->| A->1 |------>| | | |
1099 | to be perceptible to CPU 2 +-------+ | | | |
1100 | : : +-------+ | |
1101 | ||
1102 | ||
1103 | To illustrate this more completely, consider what could happen if the code | |
1104 | contained a load of A either side of the read barrier: | |
1105 | ||
1106 | CPU 1 CPU 2 | |
1107 | ======================= ======================= | |
1108 | { A = 0, B = 9 } | |
1109 | STORE A=1 | |
1110 | <write barrier> | |
1111 | STORE B=2 | |
1112 | LOAD B | |
1113 | LOAD A [first load of A] | |
1114 | <read barrier> | |
1115 | LOAD A [second load of A] | |
1116 | ||
1117 | Even though the two loads of A both occur after the load of B, they may both | |
1118 | come up with different values: | |
1119 | ||
1120 | +-------+ : : : : | |
1121 | | | +------+ +-------+ | |
1122 | | |------>| A=1 |------ --->| A->0 | | |
1123 | | | +------+ \ +-------+ | |
1124 | | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 | | |
1125 | | | +------+ | +-------+ | |
1126 | | |------>| B=2 |--- | : : | |
1127 | | | +------+ \ | : : +-------+ | |
1128 | +-------+ : : \ | +-------+ | | | |
1129 | ---------->| B->2 |------>| | | |
1130 | | +-------+ | CPU 2 | | |
1131 | | : : | | | |
1132 | | : : | | | |
1133 | | +-------+ | | | |
1134 | | | A->0 |------>| 1st | | |
1135 | | +-------+ | | | |
1136 | At this point the read ----> \ rrrrrrrrrrrrrrrrr | | | |
1137 | barrier causes all effects \ +-------+ | | | |
1138 | prior to the storage of B ---->| A->1 |------>| 2nd | | |
1139 | to be perceptible to CPU 2 +-------+ | | | |
1140 | : : +-------+ | |
1141 | ||
1142 | ||
1143 | But it may be that the update to A from CPU 1 becomes perceptible to CPU 2 | |
1144 | before the read barrier completes anyway: | |
1145 | ||
1146 | +-------+ : : : : | |
1147 | | | +------+ +-------+ | |
1148 | | |------>| A=1 |------ --->| A->0 | | |
1149 | | | +------+ \ +-------+ | |
1150 | | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 | | |
1151 | | | +------+ | +-------+ | |
1152 | | |------>| B=2 |--- | : : | |
1153 | | | +------+ \ | : : +-------+ | |
1154 | +-------+ : : \ | +-------+ | | | |
1155 | ---------->| B->2 |------>| | | |
1156 | | +-------+ | CPU 2 | | |
1157 | | : : | | | |
1158 | \ : : | | | |
1159 | \ +-------+ | | | |
1160 | ---->| A->1 |------>| 1st | | |
1161 | +-------+ | | | |
1162 | rrrrrrrrrrrrrrrrr | | | |
1163 | +-------+ | | | |
1164 | | A->1 |------>| 2nd | | |
1165 | +-------+ | | | |
1166 | : : +-------+ | |
1167 | ||
1168 | ||
1169 | The guarantee is that the second load will always come up with A == 1 if the | |
1170 | load of B came up with B == 2. No such guarantee exists for the first load of | |
1171 | A; that may come up with either A == 0 or A == 1. | |
1172 | ||
1173 | ||
1174 | READ MEMORY BARRIERS VS LOAD SPECULATION | |
1175 | ---------------------------------------- | |
1176 | ||
1177 | Many CPUs speculate with loads: that is they see that they will need to load an | |
1178 | item from memory, and they find a time where they're not using the bus for any | |
1179 | other loads, and so do the load in advance - even though they haven't actually | |
1180 | got to that point in the instruction execution flow yet. This permits the | |
1181 | actual load instruction to potentially complete immediately because the CPU | |
1182 | already has the value to hand. | |
1183 | ||
1184 | It may turn out that the CPU didn't actually need the value - perhaps because a | |
1185 | branch circumvented the load - in which case it can discard the value or just | |
1186 | cache it for later use. | |
1187 | ||
1188 | Consider: | |
1189 | ||
e0edc78f | 1190 | CPU 1 CPU 2 |
670bd95e | 1191 | ======================= ======================= |
e0edc78f IM |
1192 | LOAD B |
1193 | DIVIDE } Divide instructions generally | |
1194 | DIVIDE } take a long time to perform | |
1195 | LOAD A | |
670bd95e DH |
1196 | |
1197 | Which might appear as this: | |
1198 | ||
1199 | : : +-------+ | |
1200 | +-------+ | | | |
1201 | --->| B->2 |------>| | | |
1202 | +-------+ | CPU 2 | | |
1203 | : :DIVIDE | | | |
1204 | +-------+ | | | |
1205 | The CPU being busy doing a ---> --->| A->0 |~~~~ | | | |
1206 | division speculates on the +-------+ ~ | | | |
1207 | LOAD of A : : ~ | | | |
1208 | : :DIVIDE | | | |
1209 | : : ~ | | | |
1210 | Once the divisions are complete --> : : ~-->| | | |
1211 | the CPU can then perform the : : | | | |
1212 | LOAD with immediate effect : : +-------+ | |
1213 | ||
1214 | ||
1215 | Placing a read barrier or a data dependency barrier just before the second | |
1216 | load: | |
1217 | ||
e0edc78f | 1218 | CPU 1 CPU 2 |
670bd95e | 1219 | ======================= ======================= |
e0edc78f IM |
1220 | LOAD B |
1221 | DIVIDE | |
1222 | DIVIDE | |
670bd95e | 1223 | <read barrier> |
e0edc78f | 1224 | LOAD A |
670bd95e DH |
1225 | |
1226 | will force any value speculatively obtained to be reconsidered to an extent | |
1227 | dependent on the type of barrier used. If there was no change made to the | |
1228 | speculated memory location, then the speculated value will just be used: | |
1229 | ||
1230 | : : +-------+ | |
1231 | +-------+ | | | |
1232 | --->| B->2 |------>| | | |
1233 | +-------+ | CPU 2 | | |
1234 | : :DIVIDE | | | |
1235 | +-------+ | | | |
1236 | The CPU being busy doing a ---> --->| A->0 |~~~~ | | | |
1237 | division speculates on the +-------+ ~ | | | |
1238 | LOAD of A : : ~ | | | |
1239 | : :DIVIDE | | | |
1240 | : : ~ | | | |
1241 | : : ~ | | | |
1242 | rrrrrrrrrrrrrrrr~ | | | |
1243 | : : ~ | | | |
1244 | : : ~-->| | | |
1245 | : : | | | |
1246 | : : +-------+ | |
1247 | ||
1248 | ||
1249 | but if there was an update or an invalidation from another CPU pending, then | |
1250 | the speculation will be cancelled and the value reloaded: | |
1251 | ||
1252 | : : +-------+ | |
1253 | +-------+ | | | |
1254 | --->| B->2 |------>| | | |
1255 | +-------+ | CPU 2 | | |
1256 | : :DIVIDE | | | |
1257 | +-------+ | | | |
1258 | The CPU being busy doing a ---> --->| A->0 |~~~~ | | | |
1259 | division speculates on the +-------+ ~ | | | |
1260 | LOAD of A : : ~ | | | |
1261 | : :DIVIDE | | | |
1262 | : : ~ | | | |
1263 | : : ~ | | | |
1264 | rrrrrrrrrrrrrrrrr | | | |
1265 | +-------+ | | | |
1266 | The speculation is discarded ---> --->| A->1 |------>| | | |
1267 | and an updated value is +-------+ | | | |
1268 | retrieved : : +-------+ | |
108b42b4 DH |
1269 | |
1270 | ||
241e6663 PM |
1271 | TRANSITIVITY |
1272 | ------------ | |
1273 | ||
1274 | Transitivity is a deeply intuitive notion about ordering that is not | |
1275 | always provided by real computer systems. The following example | |
f36fe1e7 | 1276 | demonstrates transitivity: |
241e6663 PM |
1277 | |
1278 | CPU 1 CPU 2 CPU 3 | |
1279 | ======================= ======================= ======================= | |
1280 | { X = 0, Y = 0 } | |
1281 | STORE X=1 LOAD X STORE Y=1 | |
1282 | <general barrier> <general barrier> | |
1283 | LOAD Y LOAD X | |
1284 | ||
1285 | Suppose that CPU 2's load from X returns 1 and its load from Y returns 0. | |
1286 | This indicates that CPU 2's load from X in some sense follows CPU 1's | |
1287 | store to X and that CPU 2's load from Y in some sense preceded CPU 3's | |
1288 | store to Y. The question is then "Can CPU 3's load from X return 0?" | |
1289 | ||
1290 | Because CPU 2's load from X in some sense came after CPU 1's store, it | |
1291 | is natural to expect that CPU 3's load from X must therefore return 1. | |
1292 | This expectation is an example of transitivity: if a load executing on | |
1293 | CPU A follows a load from the same variable executing on CPU B, then | |
1294 | CPU A's load must either return the same value that CPU B's load did, | |
1295 | or must return some later value. | |
1296 | ||
1297 | In the Linux kernel, use of general memory barriers guarantees | |
1298 | transitivity. Therefore, in the above example, if CPU 2's load from X | |
1299 | returns 1 and its load from Y returns 0, then CPU 3's load from X must | |
1300 | also return 1. | |
1301 | ||
1302 | However, transitivity is -not- guaranteed for read or write barriers. | |
1303 | For example, suppose that CPU 2's general barrier in the above example | |
1304 | is changed to a read barrier as shown below: | |
1305 | ||
1306 | CPU 1 CPU 2 CPU 3 | |
1307 | ======================= ======================= ======================= | |
1308 | { X = 0, Y = 0 } | |
1309 | STORE X=1 LOAD X STORE Y=1 | |
1310 | <read barrier> <general barrier> | |
1311 | LOAD Y LOAD X | |
1312 | ||
1313 | This substitution destroys transitivity: in this example, it is perfectly | |
1314 | legal for CPU 2's load from X to return 1, its load from Y to return 0, | |
1315 | and CPU 3's load from X to return 0. | |
1316 | ||
1317 | The key point is that although CPU 2's read barrier orders its pair | |
1318 | of loads, it does not guarantee to order CPU 1's store. Therefore, if | |
1319 | this example runs on a system where CPUs 1 and 2 share a store buffer | |
1320 | or a level of cache, CPU 2 might have early access to CPU 1's writes. | |
1321 | General barriers are therefore required to ensure that all CPUs agree | |
1322 | on the combined order of CPU 1's and CPU 2's accesses. | |
1323 | ||
c535cc92 PM |
1324 | General barriers provide "global transitivity", so that all CPUs will |
1325 | agree on the order of operations. In contrast, a chain of release-acquire | |
1326 | pairs provides only "local transitivity", so that only those CPUs on | |
1327 | the chain are guaranteed to agree on the combined order of the accesses. | |
1328 | For example, switching to C code in deference to Herman Hollerith: | |
1329 | ||
1330 | int u, v, x, y, z; | |
1331 | ||
1332 | void cpu0(void) | |
1333 | { | |
1334 | r0 = smp_load_acquire(&x); | |
1335 | WRITE_ONCE(u, 1); | |
1336 | smp_store_release(&y, 1); | |
1337 | } | |
1338 | ||
1339 | void cpu1(void) | |
1340 | { | |
1341 | r1 = smp_load_acquire(&y); | |
1342 | r4 = READ_ONCE(v); | |
1343 | r5 = READ_ONCE(u); | |
1344 | smp_store_release(&z, 1); | |
1345 | } | |
1346 | ||
1347 | void cpu2(void) | |
1348 | { | |
1349 | r2 = smp_load_acquire(&z); | |
1350 | smp_store_release(&x, 1); | |
1351 | } | |
1352 | ||
1353 | void cpu3(void) | |
1354 | { | |
1355 | WRITE_ONCE(v, 1); | |
1356 | smp_mb(); | |
1357 | r3 = READ_ONCE(u); | |
1358 | } | |
1359 | ||
1360 | Because cpu0(), cpu1(), and cpu2() participate in a local transitive | |
1361 | chain of smp_store_release()/smp_load_acquire() pairs, the following | |
1362 | outcome is prohibited: | |
1363 | ||
1364 | r0 == 1 && r1 == 1 && r2 == 1 | |
1365 | ||
1366 | Furthermore, because of the release-acquire relationship between cpu0() | |
1367 | and cpu1(), cpu1() must see cpu0()'s writes, so that the following | |
1368 | outcome is prohibited: | |
1369 | ||
1370 | r1 == 1 && r5 == 0 | |
1371 | ||
1372 | However, the transitivity of release-acquire is local to the participating | |
1373 | CPUs and does not apply to cpu3(). Therefore, the following outcome | |
1374 | is possible: | |
1375 | ||
1376 | r0 == 0 && r1 == 1 && r2 == 1 && r3 == 0 && r4 == 0 | |
1377 | ||
37ef0341 PM |
1378 | As an aside, the following outcome is also possible: |
1379 | ||
1380 | r0 == 0 && r1 == 1 && r2 == 1 && r3 == 0 && r4 == 0 && r5 == 1 | |
1381 | ||
c535cc92 PM |
1382 | Although cpu0(), cpu1(), and cpu2() will see their respective reads and |
1383 | writes in order, CPUs not involved in the release-acquire chain might | |
1384 | well disagree on the order. This disagreement stems from the fact that | |
1385 | the weak memory-barrier instructions used to implement smp_load_acquire() | |
1386 | and smp_store_release() are not required to order prior stores against | |
1387 | subsequent loads in all cases. This means that cpu3() can see cpu0()'s | |
1388 | store to u as happening -after- cpu1()'s load from v, even though | |
1389 | both cpu0() and cpu1() agree that these two operations occurred in the | |
1390 | intended order. | |
1391 | ||
1392 | However, please keep in mind that smp_load_acquire() is not magic. | |
1393 | In particular, it simply reads from its argument with ordering. It does | |
1394 | -not- ensure that any particular value will be read. Therefore, the | |
1395 | following outcome is possible: | |
1396 | ||
1397 | r0 == 0 && r1 == 0 && r2 == 0 && r5 == 0 | |
1398 | ||
1399 | Note that this outcome can happen even on a mythical sequentially | |
1400 | consistent system where nothing is ever reordered. | |
1401 | ||
1402 | To reiterate, if your code requires global transitivity, use general | |
1403 | barriers throughout. | |
241e6663 PM |
1404 | |
1405 | ||
108b42b4 DH |
1406 | ======================== |
1407 | EXPLICIT KERNEL BARRIERS | |
1408 | ======================== | |
1409 | ||
1410 | The Linux kernel has a variety of different barriers that act at different | |
1411 | levels: | |
1412 | ||
1413 | (*) Compiler barrier. | |
1414 | ||
1415 | (*) CPU memory barriers. | |
1416 | ||
1417 | (*) MMIO write barrier. | |
1418 | ||
1419 | ||
1420 | COMPILER BARRIER | |
1421 | ---------------- | |
1422 | ||
1423 | The Linux kernel has an explicit compiler barrier function that prevents the | |
1424 | compiler from moving the memory accesses either side of it to the other side: | |
1425 | ||
1426 | barrier(); | |
1427 | ||
9af194ce PM |
1428 | This is a general barrier -- there are no read-read or write-write |
1429 | variants of barrier(). However, READ_ONCE() and WRITE_ONCE() can be | |
1430 | thought of as weak forms of barrier() that affect only the specific | |
1431 | accesses flagged by the READ_ONCE() or WRITE_ONCE(). | |
108b42b4 | 1432 | |
692118da PM |
1433 | The barrier() function has the following effects: |
1434 | ||
1435 | (*) Prevents the compiler from reordering accesses following the | |
1436 | barrier() to precede any accesses preceding the barrier(). | |
1437 | One example use for this property is to ease communication between | |
1438 | interrupt-handler code and the code that was interrupted. | |
1439 | ||
1440 | (*) Within a loop, forces the compiler to load the variables used | |
1441 | in that loop's conditional on each pass through that loop. | |
1442 | ||
9af194ce PM |
1443 | The READ_ONCE() and WRITE_ONCE() functions can prevent any number of |
1444 | optimizations that, while perfectly safe in single-threaded code, can | |
1445 | be fatal in concurrent code. Here are some examples of these sorts | |
1446 | of optimizations: | |
692118da | 1447 | |
449f7413 PM |
1448 | (*) The compiler is within its rights to reorder loads and stores |
1449 | to the same variable, and in some cases, the CPU is within its | |
1450 | rights to reorder loads to the same variable. This means that | |
1451 | the following code: | |
1452 | ||
1453 | a[0] = x; | |
1454 | a[1] = x; | |
1455 | ||
1456 | Might result in an older value of x stored in a[1] than in a[0]. | |
1457 | Prevent both the compiler and the CPU from doing this as follows: | |
1458 | ||
9af194ce PM |
1459 | a[0] = READ_ONCE(x); |
1460 | a[1] = READ_ONCE(x); | |
449f7413 | 1461 | |
9af194ce PM |
1462 | In short, READ_ONCE() and WRITE_ONCE() provide cache coherence for |
1463 | accesses from multiple CPUs to a single variable. | |
449f7413 | 1464 | |
692118da PM |
1465 | (*) The compiler is within its rights to merge successive loads from |
1466 | the same variable. Such merging can cause the compiler to "optimize" | |
1467 | the following code: | |
1468 | ||
1469 | while (tmp = a) | |
1470 | do_something_with(tmp); | |
1471 | ||
1472 | into the following code, which, although in some sense legitimate | |
1473 | for single-threaded code, is almost certainly not what the developer | |
1474 | intended: | |
1475 | ||
1476 | if (tmp = a) | |
1477 | for (;;) | |
1478 | do_something_with(tmp); | |
1479 | ||
9af194ce | 1480 | Use READ_ONCE() to prevent the compiler from doing this to you: |
692118da | 1481 | |
9af194ce | 1482 | while (tmp = READ_ONCE(a)) |
692118da PM |
1483 | do_something_with(tmp); |
1484 | ||
1485 | (*) The compiler is within its rights to reload a variable, for example, | |
1486 | in cases where high register pressure prevents the compiler from | |
1487 | keeping all data of interest in registers. The compiler might | |
1488 | therefore optimize the variable 'tmp' out of our previous example: | |
1489 | ||
1490 | while (tmp = a) | |
1491 | do_something_with(tmp); | |
1492 | ||
1493 | This could result in the following code, which is perfectly safe in | |
1494 | single-threaded code, but can be fatal in concurrent code: | |
1495 | ||
1496 | while (a) | |
1497 | do_something_with(a); | |
1498 | ||
1499 | For example, the optimized version of this code could result in | |
1500 | passing a zero to do_something_with() in the case where the variable | |
1501 | a was modified by some other CPU between the "while" statement and | |
1502 | the call to do_something_with(). | |
1503 | ||
9af194ce | 1504 | Again, use READ_ONCE() to prevent the compiler from doing this: |
692118da | 1505 | |
9af194ce | 1506 | while (tmp = READ_ONCE(a)) |
692118da PM |
1507 | do_something_with(tmp); |
1508 | ||
1509 | Note that if the compiler runs short of registers, it might save | |
1510 | tmp onto the stack. The overhead of this saving and later restoring | |
1511 | is why compilers reload variables. Doing so is perfectly safe for | |
1512 | single-threaded code, so you need to tell the compiler about cases | |
1513 | where it is not safe. | |
1514 | ||
1515 | (*) The compiler is within its rights to omit a load entirely if it knows | |
1516 | what the value will be. For example, if the compiler can prove that | |
1517 | the value of variable 'a' is always zero, it can optimize this code: | |
1518 | ||
1519 | while (tmp = a) | |
1520 | do_something_with(tmp); | |
1521 | ||
1522 | Into this: | |
1523 | ||
1524 | do { } while (0); | |
1525 | ||
9af194ce PM |
1526 | This transformation is a win for single-threaded code because it |
1527 | gets rid of a load and a branch. The problem is that the compiler | |
1528 | will carry out its proof assuming that the current CPU is the only | |
1529 | one updating variable 'a'. If variable 'a' is shared, then the | |
1530 | compiler's proof will be erroneous. Use READ_ONCE() to tell the | |
1531 | compiler that it doesn't know as much as it thinks it does: | |
692118da | 1532 | |
9af194ce | 1533 | while (tmp = READ_ONCE(a)) |
692118da PM |
1534 | do_something_with(tmp); |
1535 | ||
1536 | But please note that the compiler is also closely watching what you | |
9af194ce | 1537 | do with the value after the READ_ONCE(). For example, suppose you |
692118da PM |
1538 | do the following and MAX is a preprocessor macro with the value 1: |
1539 | ||
9af194ce | 1540 | while ((tmp = READ_ONCE(a)) % MAX) |
692118da PM |
1541 | do_something_with(tmp); |
1542 | ||
1543 | Then the compiler knows that the result of the "%" operator applied | |
1544 | to MAX will always be zero, again allowing the compiler to optimize | |
1545 | the code into near-nonexistence. (It will still load from the | |
1546 | variable 'a'.) | |
1547 | ||
1548 | (*) Similarly, the compiler is within its rights to omit a store entirely | |
1549 | if it knows that the variable already has the value being stored. | |
1550 | Again, the compiler assumes that the current CPU is the only one | |
1551 | storing into the variable, which can cause the compiler to do the | |
1552 | wrong thing for shared variables. For example, suppose you have | |
1553 | the following: | |
1554 | ||
1555 | a = 0; | |
65f95ff2 | 1556 | ... Code that does not store to variable a ... |
692118da PM |
1557 | a = 0; |
1558 | ||
1559 | The compiler sees that the value of variable 'a' is already zero, so | |
1560 | it might well omit the second store. This would come as a fatal | |
1561 | surprise if some other CPU might have stored to variable 'a' in the | |
1562 | meantime. | |
1563 | ||
9af194ce | 1564 | Use WRITE_ONCE() to prevent the compiler from making this sort of |
692118da PM |
1565 | wrong guess: |
1566 | ||
9af194ce | 1567 | WRITE_ONCE(a, 0); |
65f95ff2 | 1568 | ... Code that does not store to variable a ... |
9af194ce | 1569 | WRITE_ONCE(a, 0); |
692118da PM |
1570 | |
1571 | (*) The compiler is within its rights to reorder memory accesses unless | |
1572 | you tell it not to. For example, consider the following interaction | |
1573 | between process-level code and an interrupt handler: | |
1574 | ||
1575 | void process_level(void) | |
1576 | { | |
1577 | msg = get_message(); | |
1578 | flag = true; | |
1579 | } | |
1580 | ||
1581 | void interrupt_handler(void) | |
1582 | { | |
1583 | if (flag) | |
1584 | process_message(msg); | |
1585 | } | |
1586 | ||
df5cbb27 | 1587 | There is nothing to prevent the compiler from transforming |
692118da PM |
1588 | process_level() to the following, in fact, this might well be a |
1589 | win for single-threaded code: | |
1590 | ||
1591 | void process_level(void) | |
1592 | { | |
1593 | flag = true; | |
1594 | msg = get_message(); | |
1595 | } | |
1596 | ||
1597 | If the interrupt occurs between these two statement, then | |
9af194ce | 1598 | interrupt_handler() might be passed a garbled msg. Use WRITE_ONCE() |
692118da PM |
1599 | to prevent this as follows: |
1600 | ||
1601 | void process_level(void) | |
1602 | { | |
9af194ce PM |
1603 | WRITE_ONCE(msg, get_message()); |
1604 | WRITE_ONCE(flag, true); | |
692118da PM |
1605 | } |
1606 | ||
1607 | void interrupt_handler(void) | |
1608 | { | |
9af194ce PM |
1609 | if (READ_ONCE(flag)) |
1610 | process_message(READ_ONCE(msg)); | |
692118da PM |
1611 | } |
1612 | ||
9af194ce PM |
1613 | Note that the READ_ONCE() and WRITE_ONCE() wrappers in |
1614 | interrupt_handler() are needed if this interrupt handler can itself | |
1615 | be interrupted by something that also accesses 'flag' and 'msg', | |
1616 | for example, a nested interrupt or an NMI. Otherwise, READ_ONCE() | |
1617 | and WRITE_ONCE() are not needed in interrupt_handler() other than | |
1618 | for documentation purposes. (Note also that nested interrupts | |
1619 | do not typically occur in modern Linux kernels, in fact, if an | |
1620 | interrupt handler returns with interrupts enabled, you will get a | |
1621 | WARN_ONCE() splat.) | |
1622 | ||
1623 | You should assume that the compiler can move READ_ONCE() and | |
1624 | WRITE_ONCE() past code not containing READ_ONCE(), WRITE_ONCE(), | |
1625 | barrier(), or similar primitives. | |
1626 | ||
1627 | This effect could also be achieved using barrier(), but READ_ONCE() | |
1628 | and WRITE_ONCE() are more selective: With READ_ONCE() and | |
1629 | WRITE_ONCE(), the compiler need only forget the contents of the | |
1630 | indicated memory locations, while with barrier() the compiler must | |
1631 | discard the value of all memory locations that it has currented | |
1632 | cached in any machine registers. Of course, the compiler must also | |
1633 | respect the order in which the READ_ONCE()s and WRITE_ONCE()s occur, | |
1634 | though the CPU of course need not do so. | |
692118da PM |
1635 | |
1636 | (*) The compiler is within its rights to invent stores to a variable, | |
1637 | as in the following example: | |
1638 | ||
1639 | if (a) | |
1640 | b = a; | |
1641 | else | |
1642 | b = 42; | |
1643 | ||
1644 | The compiler might save a branch by optimizing this as follows: | |
1645 | ||
1646 | b = 42; | |
1647 | if (a) | |
1648 | b = a; | |
1649 | ||
1650 | In single-threaded code, this is not only safe, but also saves | |
1651 | a branch. Unfortunately, in concurrent code, this optimization | |
1652 | could cause some other CPU to see a spurious value of 42 -- even | |
1653 | if variable 'a' was never zero -- when loading variable 'b'. | |
9af194ce | 1654 | Use WRITE_ONCE() to prevent this as follows: |
692118da PM |
1655 | |
1656 | if (a) | |
9af194ce | 1657 | WRITE_ONCE(b, a); |
692118da | 1658 | else |
9af194ce | 1659 | WRITE_ONCE(b, 42); |
692118da PM |
1660 | |
1661 | The compiler can also invent loads. These are usually less | |
1662 | damaging, but they can result in cache-line bouncing and thus in | |
9af194ce | 1663 | poor performance and scalability. Use READ_ONCE() to prevent |
692118da PM |
1664 | invented loads. |
1665 | ||
1666 | (*) For aligned memory locations whose size allows them to be accessed | |
1667 | with a single memory-reference instruction, prevents "load tearing" | |
1668 | and "store tearing," in which a single large access is replaced by | |
1669 | multiple smaller accesses. For example, given an architecture having | |
1670 | 16-bit store instructions with 7-bit immediate fields, the compiler | |
1671 | might be tempted to use two 16-bit store-immediate instructions to | |
1672 | implement the following 32-bit store: | |
1673 | ||
1674 | p = 0x00010002; | |
1675 | ||
1676 | Please note that GCC really does use this sort of optimization, | |
1677 | which is not surprising given that it would likely take more | |
1678 | than two instructions to build the constant and then store it. | |
1679 | This optimization can therefore be a win in single-threaded code. | |
1680 | In fact, a recent bug (since fixed) caused GCC to incorrectly use | |
1681 | this optimization in a volatile store. In the absence of such bugs, | |
9af194ce | 1682 | use of WRITE_ONCE() prevents store tearing in the following example: |
692118da | 1683 | |
9af194ce | 1684 | WRITE_ONCE(p, 0x00010002); |
692118da PM |
1685 | |
1686 | Use of packed structures can also result in load and store tearing, | |
1687 | as in this example: | |
1688 | ||
1689 | struct __attribute__((__packed__)) foo { | |
1690 | short a; | |
1691 | int b; | |
1692 | short c; | |
1693 | }; | |
1694 | struct foo foo1, foo2; | |
1695 | ... | |
1696 | ||
1697 | foo2.a = foo1.a; | |
1698 | foo2.b = foo1.b; | |
1699 | foo2.c = foo1.c; | |
1700 | ||
9af194ce PM |
1701 | Because there are no READ_ONCE() or WRITE_ONCE() wrappers and no |
1702 | volatile markings, the compiler would be well within its rights to | |
1703 | implement these three assignment statements as a pair of 32-bit | |
1704 | loads followed by a pair of 32-bit stores. This would result in | |
1705 | load tearing on 'foo1.b' and store tearing on 'foo2.b'. READ_ONCE() | |
1706 | and WRITE_ONCE() again prevent tearing in this example: | |
692118da PM |
1707 | |
1708 | foo2.a = foo1.a; | |
9af194ce | 1709 | WRITE_ONCE(foo2.b, READ_ONCE(foo1.b)); |
692118da PM |
1710 | foo2.c = foo1.c; |
1711 | ||
9af194ce PM |
1712 | All that aside, it is never necessary to use READ_ONCE() and |
1713 | WRITE_ONCE() on a variable that has been marked volatile. For example, | |
1714 | because 'jiffies' is marked volatile, it is never necessary to | |
1715 | say READ_ONCE(jiffies). The reason for this is that READ_ONCE() and | |
1716 | WRITE_ONCE() are implemented as volatile casts, which has no effect when | |
1717 | its argument is already marked volatile. | |
692118da PM |
1718 | |
1719 | Please note that these compiler barriers have no direct effect on the CPU, | |
1720 | which may then reorder things however it wishes. | |
108b42b4 DH |
1721 | |
1722 | ||
1723 | CPU MEMORY BARRIERS | |
1724 | ------------------- | |
1725 | ||
1726 | The Linux kernel has eight basic CPU memory barriers: | |
1727 | ||
1728 | TYPE MANDATORY SMP CONDITIONAL | |
1729 | =============== ======================= =========================== | |
1730 | GENERAL mb() smp_mb() | |
1731 | WRITE wmb() smp_wmb() | |
1732 | READ rmb() smp_rmb() | |
1733 | DATA DEPENDENCY read_barrier_depends() smp_read_barrier_depends() | |
1734 | ||
1735 | ||
73f10281 | 1736 | All memory barriers except the data dependency barriers imply a compiler |
0b6fa347 | 1737 | barrier. Data dependencies do not impose any additional compiler ordering. |
73f10281 | 1738 | |
9af194ce PM |
1739 | Aside: In the case of data dependencies, the compiler would be expected |
1740 | to issue the loads in the correct order (eg. `a[b]` would have to load | |
1741 | the value of b before loading a[b]), however there is no guarantee in | |
1742 | the C specification that the compiler may not speculate the value of b | |
1743 | (eg. is equal to 1) and load a before b (eg. tmp = a[1]; if (b != 1) | |
0b6fa347 SP |
1744 | tmp = a[b]; ). There is also the problem of a compiler reloading b after |
1745 | having loaded a[b], thus having a newer copy of b than a[b]. A consensus | |
9af194ce PM |
1746 | has not yet been reached about these problems, however the READ_ONCE() |
1747 | macro is a good place to start looking. | |
108b42b4 DH |
1748 | |
1749 | SMP memory barriers are reduced to compiler barriers on uniprocessor compiled | |
81fc6323 | 1750 | systems because it is assumed that a CPU will appear to be self-consistent, |
108b42b4 | 1751 | and will order overlapping accesses correctly with respect to itself. |
6a65d263 | 1752 | However, see the subsection on "Virtual Machine Guests" below. |
108b42b4 DH |
1753 | |
1754 | [!] Note that SMP memory barriers _must_ be used to control the ordering of | |
1755 | references to shared memory on SMP systems, though the use of locking instead | |
1756 | is sufficient. | |
1757 | ||
1758 | Mandatory barriers should not be used to control SMP effects, since mandatory | |
6a65d263 MT |
1759 | barriers impose unnecessary overhead on both SMP and UP systems. They may, |
1760 | however, be used to control MMIO effects on accesses through relaxed memory I/O | |
1761 | windows. These barriers are required even on non-SMP systems as they affect | |
1762 | the order in which memory operations appear to a device by prohibiting both the | |
1763 | compiler and the CPU from reordering them. | |
108b42b4 DH |
1764 | |
1765 | ||
1766 | There are some more advanced barrier functions: | |
1767 | ||
b92b8b35 | 1768 | (*) smp_store_mb(var, value) |
108b42b4 | 1769 | |
75b2bd55 | 1770 | This assigns the value to the variable and then inserts a full memory |
2d142e59 DB |
1771 | barrier after it. It isn't guaranteed to insert anything more than a |
1772 | compiler barrier in a UP compilation. | |
108b42b4 DH |
1773 | |
1774 | ||
1b15611e PZ |
1775 | (*) smp_mb__before_atomic(); |
1776 | (*) smp_mb__after_atomic(); | |
108b42b4 | 1777 | |
1b15611e PZ |
1778 | These are for use with atomic (such as add, subtract, increment and |
1779 | decrement) functions that don't return a value, especially when used for | |
1780 | reference counting. These functions do not imply memory barriers. | |
1781 | ||
1782 | These are also used for atomic bitop functions that do not return a | |
1783 | value (such as set_bit and clear_bit). | |
108b42b4 DH |
1784 | |
1785 | As an example, consider a piece of code that marks an object as being dead | |
1786 | and then decrements the object's reference count: | |
1787 | ||
1788 | obj->dead = 1; | |
1b15611e | 1789 | smp_mb__before_atomic(); |
108b42b4 DH |
1790 | atomic_dec(&obj->ref_count); |
1791 | ||
1792 | This makes sure that the death mark on the object is perceived to be set | |
1793 | *before* the reference counter is decremented. | |
1794 | ||
1795 | See Documentation/atomic_ops.txt for more information. See the "Atomic | |
1796 | operations" subsection for information on where to use these. | |
1797 | ||
1798 | ||
ad2ad5d3 | 1799 | (*) lockless_dereference(); |
0b6fa347 | 1800 | |
ad2ad5d3 PM |
1801 | This can be thought of as a pointer-fetch wrapper around the |
1802 | smp_read_barrier_depends() data-dependency barrier. | |
1803 | ||
1804 | This is also similar to rcu_dereference(), but in cases where | |
1805 | object lifetime is handled by some mechanism other than RCU, for | |
1806 | example, when the objects removed only when the system goes down. | |
1807 | In addition, lockless_dereference() is used in some data structures | |
1808 | that can be used both with and without RCU. | |
1809 | ||
1810 | ||
1077fa36 AD |
1811 | (*) dma_wmb(); |
1812 | (*) dma_rmb(); | |
1813 | ||
1814 | These are for use with consistent memory to guarantee the ordering | |
1815 | of writes or reads of shared memory accessible to both the CPU and a | |
1816 | DMA capable device. | |
1817 | ||
1818 | For example, consider a device driver that shares memory with a device | |
1819 | and uses a descriptor status value to indicate if the descriptor belongs | |
1820 | to the device or the CPU, and a doorbell to notify it when new | |
1821 | descriptors are available: | |
1822 | ||
1823 | if (desc->status != DEVICE_OWN) { | |
1824 | /* do not read data until we own descriptor */ | |
1825 | dma_rmb(); | |
1826 | ||
1827 | /* read/modify data */ | |
1828 | read_data = desc->data; | |
1829 | desc->data = write_data; | |
1830 | ||
1831 | /* flush modifications before status update */ | |
1832 | dma_wmb(); | |
1833 | ||
1834 | /* assign ownership */ | |
1835 | desc->status = DEVICE_OWN; | |
1836 | ||
1837 | /* force memory to sync before notifying device via MMIO */ | |
1838 | wmb(); | |
1839 | ||
1840 | /* notify device of new descriptors */ | |
1841 | writel(DESC_NOTIFY, doorbell); | |
1842 | } | |
1843 | ||
1844 | The dma_rmb() allows us guarantee the device has released ownership | |
7a458007 | 1845 | before we read the data from the descriptor, and the dma_wmb() allows |
1077fa36 AD |
1846 | us to guarantee the data is written to the descriptor before the device |
1847 | can see it now has ownership. The wmb() is needed to guarantee that the | |
1848 | cache coherent memory writes have completed before attempting a write to | |
1849 | the cache incoherent MMIO region. | |
1850 | ||
1851 | See Documentation/DMA-API.txt for more information on consistent memory. | |
1852 | ||
108b42b4 DH |
1853 | MMIO WRITE BARRIER |
1854 | ------------------ | |
1855 | ||
1856 | The Linux kernel also has a special barrier for use with memory-mapped I/O | |
1857 | writes: | |
1858 | ||
1859 | mmiowb(); | |
1860 | ||
1861 | This is a variation on the mandatory write barrier that causes writes to weakly | |
1862 | ordered I/O regions to be partially ordered. Its effects may go beyond the | |
1863 | CPU->Hardware interface and actually affect the hardware at some level. | |
1864 | ||
166bda71 | 1865 | See the subsection "Acquires vs I/O accesses" for more information. |
108b42b4 DH |
1866 | |
1867 | ||
1868 | =============================== | |
1869 | IMPLICIT KERNEL MEMORY BARRIERS | |
1870 | =============================== | |
1871 | ||
1872 | Some of the other functions in the linux kernel imply memory barriers, amongst | |
670bd95e | 1873 | which are locking and scheduling functions. |
108b42b4 DH |
1874 | |
1875 | This specification is a _minimum_ guarantee; any particular architecture may | |
1876 | provide more substantial guarantees, but these may not be relied upon outside | |
1877 | of arch specific code. | |
1878 | ||
1879 | ||
166bda71 SP |
1880 | LOCK ACQUISITION FUNCTIONS |
1881 | -------------------------- | |
108b42b4 DH |
1882 | |
1883 | The Linux kernel has a number of locking constructs: | |
1884 | ||
1885 | (*) spin locks | |
1886 | (*) R/W spin locks | |
1887 | (*) mutexes | |
1888 | (*) semaphores | |
1889 | (*) R/W semaphores | |
108b42b4 | 1890 | |
2e4f5382 | 1891 | In all cases there are variants on "ACQUIRE" operations and "RELEASE" operations |
108b42b4 DH |
1892 | for each construct. These operations all imply certain barriers: |
1893 | ||
2e4f5382 | 1894 | (1) ACQUIRE operation implication: |
108b42b4 | 1895 | |
2e4f5382 PZ |
1896 | Memory operations issued after the ACQUIRE will be completed after the |
1897 | ACQUIRE operation has completed. | |
108b42b4 | 1898 | |
8dd853d7 PM |
1899 | Memory operations issued before the ACQUIRE may be completed after |
1900 | the ACQUIRE operation has completed. An smp_mb__before_spinlock(), | |
d956028e | 1901 | combined with a following ACQUIRE, orders prior stores against |
0b6fa347 | 1902 | subsequent loads and stores. Note that this is weaker than smp_mb()! |
d956028e | 1903 | The smp_mb__before_spinlock() primitive is free on many architectures. |
108b42b4 | 1904 | |
2e4f5382 | 1905 | (2) RELEASE operation implication: |
108b42b4 | 1906 | |
2e4f5382 PZ |
1907 | Memory operations issued before the RELEASE will be completed before the |
1908 | RELEASE operation has completed. | |
108b42b4 | 1909 | |
2e4f5382 PZ |
1910 | Memory operations issued after the RELEASE may be completed before the |
1911 | RELEASE operation has completed. | |
108b42b4 | 1912 | |
2e4f5382 | 1913 | (3) ACQUIRE vs ACQUIRE implication: |
108b42b4 | 1914 | |
2e4f5382 PZ |
1915 | All ACQUIRE operations issued before another ACQUIRE operation will be |
1916 | completed before that ACQUIRE operation. | |
108b42b4 | 1917 | |
2e4f5382 | 1918 | (4) ACQUIRE vs RELEASE implication: |
108b42b4 | 1919 | |
2e4f5382 PZ |
1920 | All ACQUIRE operations issued before a RELEASE operation will be |
1921 | completed before the RELEASE operation. | |
108b42b4 | 1922 | |
2e4f5382 | 1923 | (5) Failed conditional ACQUIRE implication: |
108b42b4 | 1924 | |
2e4f5382 PZ |
1925 | Certain locking variants of the ACQUIRE operation may fail, either due to |
1926 | being unable to get the lock immediately, or due to receiving an unblocked | |
108b42b4 DH |
1927 | signal whilst asleep waiting for the lock to become available. Failed |
1928 | locks do not imply any sort of barrier. | |
1929 | ||
2e4f5382 PZ |
1930 | [!] Note: one of the consequences of lock ACQUIREs and RELEASEs being only |
1931 | one-way barriers is that the effects of instructions outside of a critical | |
1932 | section may seep into the inside of the critical section. | |
108b42b4 | 1933 | |
2e4f5382 PZ |
1934 | An ACQUIRE followed by a RELEASE may not be assumed to be full memory barrier |
1935 | because it is possible for an access preceding the ACQUIRE to happen after the | |
1936 | ACQUIRE, and an access following the RELEASE to happen before the RELEASE, and | |
1937 | the two accesses can themselves then cross: | |
670bd95e DH |
1938 | |
1939 | *A = a; | |
2e4f5382 PZ |
1940 | ACQUIRE M |
1941 | RELEASE M | |
670bd95e DH |
1942 | *B = b; |
1943 | ||
1944 | may occur as: | |
1945 | ||
2e4f5382 | 1946 | ACQUIRE M, STORE *B, STORE *A, RELEASE M |
17eb88e0 | 1947 | |
8dd853d7 PM |
1948 | When the ACQUIRE and RELEASE are a lock acquisition and release, |
1949 | respectively, this same reordering can occur if the lock's ACQUIRE and | |
1950 | RELEASE are to the same lock variable, but only from the perspective of | |
1951 | another CPU not holding that lock. In short, a ACQUIRE followed by an | |
1952 | RELEASE may -not- be assumed to be a full memory barrier. | |
1953 | ||
12d560f4 PM |
1954 | Similarly, the reverse case of a RELEASE followed by an ACQUIRE does |
1955 | not imply a full memory barrier. Therefore, the CPU's execution of the | |
1956 | critical sections corresponding to the RELEASE and the ACQUIRE can cross, | |
1957 | so that: | |
17eb88e0 PM |
1958 | |
1959 | *A = a; | |
2e4f5382 PZ |
1960 | RELEASE M |
1961 | ACQUIRE N | |
17eb88e0 PM |
1962 | *B = b; |
1963 | ||
1964 | could occur as: | |
1965 | ||
2e4f5382 | 1966 | ACQUIRE N, STORE *B, STORE *A, RELEASE M |
17eb88e0 | 1967 | |
8dd853d7 PM |
1968 | It might appear that this reordering could introduce a deadlock. |
1969 | However, this cannot happen because if such a deadlock threatened, | |
1970 | the RELEASE would simply complete, thereby avoiding the deadlock. | |
1971 | ||
1972 | Why does this work? | |
1973 | ||
1974 | One key point is that we are only talking about the CPU doing | |
1975 | the reordering, not the compiler. If the compiler (or, for | |
1976 | that matter, the developer) switched the operations, deadlock | |
1977 | -could- occur. | |
1978 | ||
1979 | But suppose the CPU reordered the operations. In this case, | |
1980 | the unlock precedes the lock in the assembly code. The CPU | |
1981 | simply elected to try executing the later lock operation first. | |
1982 | If there is a deadlock, this lock operation will simply spin (or | |
1983 | try to sleep, but more on that later). The CPU will eventually | |
1984 | execute the unlock operation (which preceded the lock operation | |
1985 | in the assembly code), which will unravel the potential deadlock, | |
1986 | allowing the lock operation to succeed. | |
1987 | ||
1988 | But what if the lock is a sleeplock? In that case, the code will | |
1989 | try to enter the scheduler, where it will eventually encounter | |
1990 | a memory barrier, which will force the earlier unlock operation | |
1991 | to complete, again unraveling the deadlock. There might be | |
1992 | a sleep-unlock race, but the locking primitive needs to resolve | |
1993 | such races properly in any case. | |
1994 | ||
108b42b4 DH |
1995 | Locks and semaphores may not provide any guarantee of ordering on UP compiled |
1996 | systems, and so cannot be counted on in such a situation to actually achieve | |
1997 | anything at all - especially with respect to I/O accesses - unless combined | |
1998 | with interrupt disabling operations. | |
1999 | ||
2000 | See also the section on "Inter-CPU locking barrier effects". | |
2001 | ||
2002 | ||
2003 | As an example, consider the following: | |
2004 | ||
2005 | *A = a; | |
2006 | *B = b; | |
2e4f5382 | 2007 | ACQUIRE |
108b42b4 DH |
2008 | *C = c; |
2009 | *D = d; | |
2e4f5382 | 2010 | RELEASE |
108b42b4 DH |
2011 | *E = e; |
2012 | *F = f; | |
2013 | ||
2014 | The following sequence of events is acceptable: | |
2015 | ||
2e4f5382 | 2016 | ACQUIRE, {*F,*A}, *E, {*C,*D}, *B, RELEASE |
108b42b4 DH |
2017 | |
2018 | [+] Note that {*F,*A} indicates a combined access. | |
2019 | ||
2020 | But none of the following are: | |
2021 | ||
2e4f5382 PZ |
2022 | {*F,*A}, *B, ACQUIRE, *C, *D, RELEASE, *E |
2023 | *A, *B, *C, ACQUIRE, *D, RELEASE, *E, *F | |
2024 | *A, *B, ACQUIRE, *C, RELEASE, *D, *E, *F | |
2025 | *B, ACQUIRE, *C, *D, RELEASE, {*F,*A}, *E | |
108b42b4 DH |
2026 | |
2027 | ||
2028 | ||
2029 | INTERRUPT DISABLING FUNCTIONS | |
2030 | ----------------------------- | |
2031 | ||
2e4f5382 PZ |
2032 | Functions that disable interrupts (ACQUIRE equivalent) and enable interrupts |
2033 | (RELEASE equivalent) will act as compiler barriers only. So if memory or I/O | |
108b42b4 DH |
2034 | barriers are required in such a situation, they must be provided from some |
2035 | other means. | |
2036 | ||
2037 | ||
50fa610a DH |
2038 | SLEEP AND WAKE-UP FUNCTIONS |
2039 | --------------------------- | |
2040 | ||
2041 | Sleeping and waking on an event flagged in global data can be viewed as an | |
2042 | interaction between two pieces of data: the task state of the task waiting for | |
2043 | the event and the global data used to indicate the event. To make sure that | |
2044 | these appear to happen in the right order, the primitives to begin the process | |
2045 | of going to sleep, and the primitives to initiate a wake up imply certain | |
2046 | barriers. | |
2047 | ||
2048 | Firstly, the sleeper normally follows something like this sequence of events: | |
2049 | ||
2050 | for (;;) { | |
2051 | set_current_state(TASK_UNINTERRUPTIBLE); | |
2052 | if (event_indicated) | |
2053 | break; | |
2054 | schedule(); | |
2055 | } | |
2056 | ||
2057 | A general memory barrier is interpolated automatically by set_current_state() | |
2058 | after it has altered the task state: | |
2059 | ||
2060 | CPU 1 | |
2061 | =============================== | |
2062 | set_current_state(); | |
b92b8b35 | 2063 | smp_store_mb(); |
50fa610a DH |
2064 | STORE current->state |
2065 | <general barrier> | |
2066 | LOAD event_indicated | |
2067 | ||
2068 | set_current_state() may be wrapped by: | |
2069 | ||
2070 | prepare_to_wait(); | |
2071 | prepare_to_wait_exclusive(); | |
2072 | ||
2073 | which therefore also imply a general memory barrier after setting the state. | |
2074 | The whole sequence above is available in various canned forms, all of which | |
2075 | interpolate the memory barrier in the right place: | |
2076 | ||
2077 | wait_event(); | |
2078 | wait_event_interruptible(); | |
2079 | wait_event_interruptible_exclusive(); | |
2080 | wait_event_interruptible_timeout(); | |
2081 | wait_event_killable(); | |
2082 | wait_event_timeout(); | |
2083 | wait_on_bit(); | |
2084 | wait_on_bit_lock(); | |
2085 | ||
2086 | ||
2087 | Secondly, code that performs a wake up normally follows something like this: | |
2088 | ||
2089 | event_indicated = 1; | |
2090 | wake_up(&event_wait_queue); | |
2091 | ||
2092 | or: | |
2093 | ||
2094 | event_indicated = 1; | |
2095 | wake_up_process(event_daemon); | |
2096 | ||
0b6fa347 SP |
2097 | A write memory barrier is implied by wake_up() and co. if and only if they |
2098 | wake something up. The barrier occurs before the task state is cleared, and so | |
2099 | sits between the STORE to indicate the event and the STORE to set TASK_RUNNING: | |
50fa610a DH |
2100 | |
2101 | CPU 1 CPU 2 | |
2102 | =============================== =============================== | |
2103 | set_current_state(); STORE event_indicated | |
b92b8b35 | 2104 | smp_store_mb(); wake_up(); |
50fa610a DH |
2105 | STORE current->state <write barrier> |
2106 | <general barrier> STORE current->state | |
2107 | LOAD event_indicated | |
2108 | ||
5726ce06 PM |
2109 | To repeat, this write memory barrier is present if and only if something |
2110 | is actually awakened. To see this, consider the following sequence of | |
2111 | events, where X and Y are both initially zero: | |
2112 | ||
2113 | CPU 1 CPU 2 | |
2114 | =============================== =============================== | |
2115 | X = 1; STORE event_indicated | |
2116 | smp_mb(); wake_up(); | |
2117 | Y = 1; wait_event(wq, Y == 1); | |
2118 | wake_up(); load from Y sees 1, no memory barrier | |
2119 | load from X might see 0 | |
2120 | ||
2121 | In contrast, if a wakeup does occur, CPU 2's load from X would be guaranteed | |
2122 | to see 1. | |
2123 | ||
50fa610a DH |
2124 | The available waker functions include: |
2125 | ||
2126 | complete(); | |
2127 | wake_up(); | |
2128 | wake_up_all(); | |
2129 | wake_up_bit(); | |
2130 | wake_up_interruptible(); | |
2131 | wake_up_interruptible_all(); | |
2132 | wake_up_interruptible_nr(); | |
2133 | wake_up_interruptible_poll(); | |
2134 | wake_up_interruptible_sync(); | |
2135 | wake_up_interruptible_sync_poll(); | |
2136 | wake_up_locked(); | |
2137 | wake_up_locked_poll(); | |
2138 | wake_up_nr(); | |
2139 | wake_up_poll(); | |
2140 | wake_up_process(); | |
2141 | ||
2142 | ||
2143 | [!] Note that the memory barriers implied by the sleeper and the waker do _not_ | |
2144 | order multiple stores before the wake-up with respect to loads of those stored | |
2145 | values after the sleeper has called set_current_state(). For instance, if the | |
2146 | sleeper does: | |
2147 | ||
2148 | set_current_state(TASK_INTERRUPTIBLE); | |
2149 | if (event_indicated) | |
2150 | break; | |
2151 | __set_current_state(TASK_RUNNING); | |
2152 | do_something(my_data); | |
2153 | ||
2154 | and the waker does: | |
2155 | ||
2156 | my_data = value; | |
2157 | event_indicated = 1; | |
2158 | wake_up(&event_wait_queue); | |
2159 | ||
2160 | there's no guarantee that the change to event_indicated will be perceived by | |
2161 | the sleeper as coming after the change to my_data. In such a circumstance, the | |
2162 | code on both sides must interpolate its own memory barriers between the | |
2163 | separate data accesses. Thus the above sleeper ought to do: | |
2164 | ||
2165 | set_current_state(TASK_INTERRUPTIBLE); | |
2166 | if (event_indicated) { | |
2167 | smp_rmb(); | |
2168 | do_something(my_data); | |
2169 | } | |
2170 | ||
2171 | and the waker should do: | |
2172 | ||
2173 | my_data = value; | |
2174 | smp_wmb(); | |
2175 | event_indicated = 1; | |
2176 | wake_up(&event_wait_queue); | |
2177 | ||
2178 | ||
108b42b4 DH |
2179 | MISCELLANEOUS FUNCTIONS |
2180 | ----------------------- | |
2181 | ||
2182 | Other functions that imply barriers: | |
2183 | ||
2184 | (*) schedule() and similar imply full memory barriers. | |
2185 | ||
108b42b4 | 2186 | |
2e4f5382 PZ |
2187 | =================================== |
2188 | INTER-CPU ACQUIRING BARRIER EFFECTS | |
2189 | =================================== | |
108b42b4 DH |
2190 | |
2191 | On SMP systems locking primitives give a more substantial form of barrier: one | |
2192 | that does affect memory access ordering on other CPUs, within the context of | |
2193 | conflict on any particular lock. | |
2194 | ||
2195 | ||
2e4f5382 PZ |
2196 | ACQUIRES VS MEMORY ACCESSES |
2197 | --------------------------- | |
108b42b4 | 2198 | |
79afecfa | 2199 | Consider the following: the system has a pair of spinlocks (M) and (Q), and |
108b42b4 DH |
2200 | three CPUs; then should the following sequence of events occur: |
2201 | ||
2202 | CPU 1 CPU 2 | |
2203 | =============================== =============================== | |
9af194ce | 2204 | WRITE_ONCE(*A, a); WRITE_ONCE(*E, e); |
2e4f5382 | 2205 | ACQUIRE M ACQUIRE Q |
9af194ce PM |
2206 | WRITE_ONCE(*B, b); WRITE_ONCE(*F, f); |
2207 | WRITE_ONCE(*C, c); WRITE_ONCE(*G, g); | |
2e4f5382 | 2208 | RELEASE M RELEASE Q |
9af194ce | 2209 | WRITE_ONCE(*D, d); WRITE_ONCE(*H, h); |
108b42b4 | 2210 | |
81fc6323 | 2211 | Then there is no guarantee as to what order CPU 3 will see the accesses to *A |
108b42b4 | 2212 | through *H occur in, other than the constraints imposed by the separate locks |
0b6fa347 | 2213 | on the separate CPUs. It might, for example, see: |
108b42b4 | 2214 | |
2e4f5382 | 2215 | *E, ACQUIRE M, ACQUIRE Q, *G, *C, *F, *A, *B, RELEASE Q, *D, *H, RELEASE M |
108b42b4 DH |
2216 | |
2217 | But it won't see any of: | |
2218 | ||
2e4f5382 PZ |
2219 | *B, *C or *D preceding ACQUIRE M |
2220 | *A, *B or *C following RELEASE M | |
2221 | *F, *G or *H preceding ACQUIRE Q | |
2222 | *E, *F or *G following RELEASE Q | |
108b42b4 DH |
2223 | |
2224 | ||
108b42b4 | 2225 | |
2e4f5382 PZ |
2226 | ACQUIRES VS I/O ACCESSES |
2227 | ------------------------ | |
108b42b4 DH |
2228 | |
2229 | Under certain circumstances (especially involving NUMA), I/O accesses within | |
2230 | two spinlocked sections on two different CPUs may be seen as interleaved by the | |
2231 | PCI bridge, because the PCI bridge does not necessarily participate in the | |
2232 | cache-coherence protocol, and is therefore incapable of issuing the required | |
2233 | read memory barriers. | |
2234 | ||
2235 | For example: | |
2236 | ||
2237 | CPU 1 CPU 2 | |
2238 | =============================== =============================== | |
2239 | spin_lock(Q) | |
2240 | writel(0, ADDR) | |
2241 | writel(1, DATA); | |
2242 | spin_unlock(Q); | |
2243 | spin_lock(Q); | |
2244 | writel(4, ADDR); | |
2245 | writel(5, DATA); | |
2246 | spin_unlock(Q); | |
2247 | ||
2248 | may be seen by the PCI bridge as follows: | |
2249 | ||
2250 | STORE *ADDR = 0, STORE *ADDR = 4, STORE *DATA = 1, STORE *DATA = 5 | |
2251 | ||
2252 | which would probably cause the hardware to malfunction. | |
2253 | ||
2254 | ||
2255 | What is necessary here is to intervene with an mmiowb() before dropping the | |
2256 | spinlock, for example: | |
2257 | ||
2258 | CPU 1 CPU 2 | |
2259 | =============================== =============================== | |
2260 | spin_lock(Q) | |
2261 | writel(0, ADDR) | |
2262 | writel(1, DATA); | |
2263 | mmiowb(); | |
2264 | spin_unlock(Q); | |
2265 | spin_lock(Q); | |
2266 | writel(4, ADDR); | |
2267 | writel(5, DATA); | |
2268 | mmiowb(); | |
2269 | spin_unlock(Q); | |
2270 | ||
81fc6323 JP |
2271 | this will ensure that the two stores issued on CPU 1 appear at the PCI bridge |
2272 | before either of the stores issued on CPU 2. | |
108b42b4 DH |
2273 | |
2274 | ||
81fc6323 JP |
2275 | Furthermore, following a store by a load from the same device obviates the need |
2276 | for the mmiowb(), because the load forces the store to complete before the load | |
108b42b4 DH |
2277 | is performed: |
2278 | ||
2279 | CPU 1 CPU 2 | |
2280 | =============================== =============================== | |
2281 | spin_lock(Q) | |
2282 | writel(0, ADDR) | |
2283 | a = readl(DATA); | |
2284 | spin_unlock(Q); | |
2285 | spin_lock(Q); | |
2286 | writel(4, ADDR); | |
2287 | b = readl(DATA); | |
2288 | spin_unlock(Q); | |
2289 | ||
2290 | ||
2291 | See Documentation/DocBook/deviceiobook.tmpl for more information. | |
2292 | ||
2293 | ||
2294 | ================================= | |
2295 | WHERE ARE MEMORY BARRIERS NEEDED? | |
2296 | ================================= | |
2297 | ||
2298 | Under normal operation, memory operation reordering is generally not going to | |
2299 | be a problem as a single-threaded linear piece of code will still appear to | |
50fa610a | 2300 | work correctly, even if it's in an SMP kernel. There are, however, four |
108b42b4 DH |
2301 | circumstances in which reordering definitely _could_ be a problem: |
2302 | ||
2303 | (*) Interprocessor interaction. | |
2304 | ||
2305 | (*) Atomic operations. | |
2306 | ||
81fc6323 | 2307 | (*) Accessing devices. |
108b42b4 DH |
2308 | |
2309 | (*) Interrupts. | |
2310 | ||
2311 | ||
2312 | INTERPROCESSOR INTERACTION | |
2313 | -------------------------- | |
2314 | ||
2315 | When there's a system with more than one processor, more than one CPU in the | |
2316 | system may be working on the same data set at the same time. This can cause | |
2317 | synchronisation problems, and the usual way of dealing with them is to use | |
2318 | locks. Locks, however, are quite expensive, and so it may be preferable to | |
2319 | operate without the use of a lock if at all possible. In such a case | |
2320 | operations that affect both CPUs may have to be carefully ordered to prevent | |
2321 | a malfunction. | |
2322 | ||
2323 | Consider, for example, the R/W semaphore slow path. Here a waiting process is | |
2324 | queued on the semaphore, by virtue of it having a piece of its stack linked to | |
2325 | the semaphore's list of waiting processes: | |
2326 | ||
2327 | struct rw_semaphore { | |
2328 | ... | |
2329 | spinlock_t lock; | |
2330 | struct list_head waiters; | |
2331 | }; | |
2332 | ||
2333 | struct rwsem_waiter { | |
2334 | struct list_head list; | |
2335 | struct task_struct *task; | |
2336 | }; | |
2337 | ||
2338 | To wake up a particular waiter, the up_read() or up_write() functions have to: | |
2339 | ||
2340 | (1) read the next pointer from this waiter's record to know as to where the | |
2341 | next waiter record is; | |
2342 | ||
81fc6323 | 2343 | (2) read the pointer to the waiter's task structure; |
108b42b4 DH |
2344 | |
2345 | (3) clear the task pointer to tell the waiter it has been given the semaphore; | |
2346 | ||
2347 | (4) call wake_up_process() on the task; and | |
2348 | ||
2349 | (5) release the reference held on the waiter's task struct. | |
2350 | ||
81fc6323 | 2351 | In other words, it has to perform this sequence of events: |
108b42b4 DH |
2352 | |
2353 | LOAD waiter->list.next; | |
2354 | LOAD waiter->task; | |
2355 | STORE waiter->task; | |
2356 | CALL wakeup | |
2357 | RELEASE task | |
2358 | ||
2359 | and if any of these steps occur out of order, then the whole thing may | |
2360 | malfunction. | |
2361 | ||
2362 | Once it has queued itself and dropped the semaphore lock, the waiter does not | |
2363 | get the lock again; it instead just waits for its task pointer to be cleared | |
2364 | before proceeding. Since the record is on the waiter's stack, this means that | |
2365 | if the task pointer is cleared _before_ the next pointer in the list is read, | |
2366 | another CPU might start processing the waiter and might clobber the waiter's | |
2367 | stack before the up*() function has a chance to read the next pointer. | |
2368 | ||
2369 | Consider then what might happen to the above sequence of events: | |
2370 | ||
2371 | CPU 1 CPU 2 | |
2372 | =============================== =============================== | |
2373 | down_xxx() | |
2374 | Queue waiter | |
2375 | Sleep | |
2376 | up_yyy() | |
2377 | LOAD waiter->task; | |
2378 | STORE waiter->task; | |
2379 | Woken up by other event | |
2380 | <preempt> | |
2381 | Resume processing | |
2382 | down_xxx() returns | |
2383 | call foo() | |
2384 | foo() clobbers *waiter | |
2385 | </preempt> | |
2386 | LOAD waiter->list.next; | |
2387 | --- OOPS --- | |
2388 | ||
2389 | This could be dealt with using the semaphore lock, but then the down_xxx() | |
2390 | function has to needlessly get the spinlock again after being woken up. | |
2391 | ||
2392 | The way to deal with this is to insert a general SMP memory barrier: | |
2393 | ||
2394 | LOAD waiter->list.next; | |
2395 | LOAD waiter->task; | |
2396 | smp_mb(); | |
2397 | STORE waiter->task; | |
2398 | CALL wakeup | |
2399 | RELEASE task | |
2400 | ||
2401 | In this case, the barrier makes a guarantee that all memory accesses before the | |
2402 | barrier will appear to happen before all the memory accesses after the barrier | |
2403 | with respect to the other CPUs on the system. It does _not_ guarantee that all | |
2404 | the memory accesses before the barrier will be complete by the time the barrier | |
2405 | instruction itself is complete. | |
2406 | ||
2407 | On a UP system - where this wouldn't be a problem - the smp_mb() is just a | |
2408 | compiler barrier, thus making sure the compiler emits the instructions in the | |
6bc39274 DH |
2409 | right order without actually intervening in the CPU. Since there's only one |
2410 | CPU, that CPU's dependency ordering logic will take care of everything else. | |
108b42b4 DH |
2411 | |
2412 | ||
2413 | ATOMIC OPERATIONS | |
2414 | ----------------- | |
2415 | ||
dbc8700e DH |
2416 | Whilst they are technically interprocessor interaction considerations, atomic |
2417 | operations are noted specially as some of them imply full memory barriers and | |
2418 | some don't, but they're very heavily relied on as a group throughout the | |
2419 | kernel. | |
2420 | ||
2421 | Any atomic operation that modifies some state in memory and returns information | |
2422 | about the state (old or new) implies an SMP-conditional general memory barrier | |
26333576 NP |
2423 | (smp_mb()) on each side of the actual operation (with the exception of |
2424 | explicit lock operations, described later). These include: | |
108b42b4 DH |
2425 | |
2426 | xchg(); | |
fb2b5819 | 2427 | atomic_xchg(); atomic_long_xchg(); |
fb2b5819 PM |
2428 | atomic_inc_return(); atomic_long_inc_return(); |
2429 | atomic_dec_return(); atomic_long_dec_return(); | |
2430 | atomic_add_return(); atomic_long_add_return(); | |
2431 | atomic_sub_return(); atomic_long_sub_return(); | |
2432 | atomic_inc_and_test(); atomic_long_inc_and_test(); | |
2433 | atomic_dec_and_test(); atomic_long_dec_and_test(); | |
2434 | atomic_sub_and_test(); atomic_long_sub_and_test(); | |
2435 | atomic_add_negative(); atomic_long_add_negative(); | |
dbc8700e DH |
2436 | test_and_set_bit(); |
2437 | test_and_clear_bit(); | |
2438 | test_and_change_bit(); | |
2439 | ||
ed2de9f7 WD |
2440 | /* when succeeds */ |
2441 | cmpxchg(); | |
2442 | atomic_cmpxchg(); atomic_long_cmpxchg(); | |
fb2b5819 PM |
2443 | atomic_add_unless(); atomic_long_add_unless(); |
2444 | ||
2e4f5382 | 2445 | These are used for such things as implementing ACQUIRE-class and RELEASE-class |
dbc8700e DH |
2446 | operations and adjusting reference counters towards object destruction, and as |
2447 | such the implicit memory barrier effects are necessary. | |
108b42b4 | 2448 | |
108b42b4 | 2449 | |
81fc6323 | 2450 | The following operations are potential problems as they do _not_ imply memory |
2e4f5382 | 2451 | barriers, but might be used for implementing such things as RELEASE-class |
dbc8700e | 2452 | operations: |
108b42b4 | 2453 | |
dbc8700e | 2454 | atomic_set(); |
108b42b4 DH |
2455 | set_bit(); |
2456 | clear_bit(); | |
2457 | change_bit(); | |
dbc8700e DH |
2458 | |
2459 | With these the appropriate explicit memory barrier should be used if necessary | |
1b15611e | 2460 | (smp_mb__before_atomic() for instance). |
108b42b4 DH |
2461 | |
2462 | ||
dbc8700e | 2463 | The following also do _not_ imply memory barriers, and so may require explicit |
1b15611e | 2464 | memory barriers under some circumstances (smp_mb__before_atomic() for |
81fc6323 | 2465 | instance): |
108b42b4 DH |
2466 | |
2467 | atomic_add(); | |
2468 | atomic_sub(); | |
2469 | atomic_inc(); | |
2470 | atomic_dec(); | |
2471 | ||
2472 | If they're used for statistics generation, then they probably don't need memory | |
2473 | barriers, unless there's a coupling between statistical data. | |
2474 | ||
2475 | If they're used for reference counting on an object to control its lifetime, | |
2476 | they probably don't need memory barriers because either the reference count | |
2477 | will be adjusted inside a locked section, or the caller will already hold | |
2478 | sufficient references to make the lock, and thus a memory barrier unnecessary. | |
2479 | ||
2480 | If they're used for constructing a lock of some description, then they probably | |
2481 | do need memory barriers as a lock primitive generally has to do things in a | |
2482 | specific order. | |
2483 | ||
108b42b4 | 2484 | Basically, each usage case has to be carefully considered as to whether memory |
dbc8700e DH |
2485 | barriers are needed or not. |
2486 | ||
26333576 NP |
2487 | The following operations are special locking primitives: |
2488 | ||
2489 | test_and_set_bit_lock(); | |
2490 | clear_bit_unlock(); | |
2491 | __clear_bit_unlock(); | |
2492 | ||
0b6fa347 SP |
2493 | These implement ACQUIRE-class and RELEASE-class operations. These should be |
2494 | used in preference to other operations when implementing locking primitives, | |
2495 | because their implementations can be optimised on many architectures. | |
26333576 | 2496 | |
dbc8700e DH |
2497 | [!] Note that special memory barrier primitives are available for these |
2498 | situations because on some CPUs the atomic instructions used imply full memory | |
2499 | barriers, and so barrier instructions are superfluous in conjunction with them, | |
2500 | and in such cases the special barrier primitives will be no-ops. | |
108b42b4 DH |
2501 | |
2502 | See Documentation/atomic_ops.txt for more information. | |
2503 | ||
2504 | ||
2505 | ACCESSING DEVICES | |
2506 | ----------------- | |
2507 | ||
2508 | Many devices can be memory mapped, and so appear to the CPU as if they're just | |
2509 | a set of memory locations. To control such a device, the driver usually has to | |
2510 | make the right memory accesses in exactly the right order. | |
2511 | ||
2512 | However, having a clever CPU or a clever compiler creates a potential problem | |
2513 | in that the carefully sequenced accesses in the driver code won't reach the | |
2514 | device in the requisite order if the CPU or the compiler thinks it is more | |
2515 | efficient to reorder, combine or merge accesses - something that would cause | |
2516 | the device to malfunction. | |
2517 | ||
2518 | Inside of the Linux kernel, I/O should be done through the appropriate accessor | |
2519 | routines - such as inb() or writel() - which know how to make such accesses | |
2520 | appropriately sequential. Whilst this, for the most part, renders the explicit | |
2521 | use of memory barriers unnecessary, there are a couple of situations where they | |
2522 | might be needed: | |
2523 | ||
2524 | (1) On some systems, I/O stores are not strongly ordered across all CPUs, and | |
2525 | so for _all_ general drivers locks should be used and mmiowb() must be | |
2526 | issued prior to unlocking the critical section. | |
2527 | ||
2528 | (2) If the accessor functions are used to refer to an I/O memory window with | |
2529 | relaxed memory access properties, then _mandatory_ memory barriers are | |
2530 | required to enforce ordering. | |
2531 | ||
2532 | See Documentation/DocBook/deviceiobook.tmpl for more information. | |
2533 | ||
2534 | ||
2535 | INTERRUPTS | |
2536 | ---------- | |
2537 | ||
2538 | A driver may be interrupted by its own interrupt service routine, and thus the | |
2539 | two parts of the driver may interfere with each other's attempts to control or | |
2540 | access the device. | |
2541 | ||
2542 | This may be alleviated - at least in part - by disabling local interrupts (a | |
2543 | form of locking), such that the critical operations are all contained within | |
2544 | the interrupt-disabled section in the driver. Whilst the driver's interrupt | |
2545 | routine is executing, the driver's core may not run on the same CPU, and its | |
2546 | interrupt is not permitted to happen again until the current interrupt has been | |
2547 | handled, thus the interrupt handler does not need to lock against that. | |
2548 | ||
2549 | However, consider a driver that was talking to an ethernet card that sports an | |
2550 | address register and a data register. If that driver's core talks to the card | |
2551 | under interrupt-disablement and then the driver's interrupt handler is invoked: | |
2552 | ||
2553 | LOCAL IRQ DISABLE | |
2554 | writew(ADDR, 3); | |
2555 | writew(DATA, y); | |
2556 | LOCAL IRQ ENABLE | |
2557 | <interrupt> | |
2558 | writew(ADDR, 4); | |
2559 | q = readw(DATA); | |
2560 | </interrupt> | |
2561 | ||
2562 | The store to the data register might happen after the second store to the | |
2563 | address register if ordering rules are sufficiently relaxed: | |
2564 | ||
2565 | STORE *ADDR = 3, STORE *ADDR = 4, STORE *DATA = y, q = LOAD *DATA | |
2566 | ||
2567 | ||
2568 | If ordering rules are relaxed, it must be assumed that accesses done inside an | |
2569 | interrupt disabled section may leak outside of it and may interleave with | |
2570 | accesses performed in an interrupt - and vice versa - unless implicit or | |
2571 | explicit barriers are used. | |
2572 | ||
2573 | Normally this won't be a problem because the I/O accesses done inside such | |
2574 | sections will include synchronous load operations on strictly ordered I/O | |
0b6fa347 | 2575 | registers that form implicit I/O barriers. If this isn't sufficient then an |
108b42b4 DH |
2576 | mmiowb() may need to be used explicitly. |
2577 | ||
2578 | ||
2579 | A similar situation may occur between an interrupt routine and two routines | |
0b6fa347 | 2580 | running on separate CPUs that communicate with each other. If such a case is |
108b42b4 DH |
2581 | likely, then interrupt-disabling locks should be used to guarantee ordering. |
2582 | ||
2583 | ||
2584 | ========================== | |
2585 | KERNEL I/O BARRIER EFFECTS | |
2586 | ========================== | |
2587 | ||
2588 | When accessing I/O memory, drivers should use the appropriate accessor | |
2589 | functions: | |
2590 | ||
2591 | (*) inX(), outX(): | |
2592 | ||
2593 | These are intended to talk to I/O space rather than memory space, but | |
0b6fa347 SP |
2594 | that's primarily a CPU-specific concept. The i386 and x86_64 processors |
2595 | do indeed have special I/O space access cycles and instructions, but many | |
108b42b4 DH |
2596 | CPUs don't have such a concept. |
2597 | ||
81fc6323 JP |
2598 | The PCI bus, amongst others, defines an I/O space concept which - on such |
2599 | CPUs as i386 and x86_64 - readily maps to the CPU's concept of I/O | |
6bc39274 DH |
2600 | space. However, it may also be mapped as a virtual I/O space in the CPU's |
2601 | memory map, particularly on those CPUs that don't support alternate I/O | |
2602 | spaces. | |
108b42b4 DH |
2603 | |
2604 | Accesses to this space may be fully synchronous (as on i386), but | |
2605 | intermediary bridges (such as the PCI host bridge) may not fully honour | |
2606 | that. | |
2607 | ||
2608 | They are guaranteed to be fully ordered with respect to each other. | |
2609 | ||
2610 | They are not guaranteed to be fully ordered with respect to other types of | |
2611 | memory and I/O operation. | |
2612 | ||
2613 | (*) readX(), writeX(): | |
2614 | ||
2615 | Whether these are guaranteed to be fully ordered and uncombined with | |
2616 | respect to each other on the issuing CPU depends on the characteristics | |
0b6fa347 | 2617 | defined for the memory window through which they're accessing. On later |
108b42b4 DH |
2618 | i386 architecture machines, for example, this is controlled by way of the |
2619 | MTRR registers. | |
2620 | ||
81fc6323 | 2621 | Ordinarily, these will be guaranteed to be fully ordered and uncombined, |
108b42b4 DH |
2622 | provided they're not accessing a prefetchable device. |
2623 | ||
2624 | However, intermediary hardware (such as a PCI bridge) may indulge in | |
2625 | deferral if it so wishes; to flush a store, a load from the same location | |
2626 | is preferred[*], but a load from the same device or from configuration | |
2627 | space should suffice for PCI. | |
2628 | ||
2629 | [*] NOTE! attempting to load from the same location as was written to may | |
e0edc78f IM |
2630 | cause a malfunction - consider the 16550 Rx/Tx serial registers for |
2631 | example. | |
108b42b4 DH |
2632 | |
2633 | Used with prefetchable I/O memory, an mmiowb() barrier may be required to | |
2634 | force stores to be ordered. | |
2635 | ||
2636 | Please refer to the PCI specification for more information on interactions | |
2637 | between PCI transactions. | |
2638 | ||
a8e0aead WD |
2639 | (*) readX_relaxed(), writeX_relaxed() |
2640 | ||
2641 | These are similar to readX() and writeX(), but provide weaker memory | |
0b6fa347 | 2642 | ordering guarantees. Specifically, they do not guarantee ordering with |
a8e0aead | 2643 | respect to normal memory accesses (e.g. DMA buffers) nor do they guarantee |
0b6fa347 SP |
2644 | ordering with respect to LOCK or UNLOCK operations. If the latter is |
2645 | required, an mmiowb() barrier can be used. Note that relaxed accesses to | |
a8e0aead WD |
2646 | the same peripheral are guaranteed to be ordered with respect to each |
2647 | other. | |
108b42b4 DH |
2648 | |
2649 | (*) ioreadX(), iowriteX() | |
2650 | ||
81fc6323 | 2651 | These will perform appropriately for the type of access they're actually |
108b42b4 DH |
2652 | doing, be it inX()/outX() or readX()/writeX(). |
2653 | ||
2654 | ||
2655 | ======================================== | |
2656 | ASSUMED MINIMUM EXECUTION ORDERING MODEL | |
2657 | ======================================== | |
2658 | ||
2659 | It has to be assumed that the conceptual CPU is weakly-ordered but that it will | |
2660 | maintain the appearance of program causality with respect to itself. Some CPUs | |
2661 | (such as i386 or x86_64) are more constrained than others (such as powerpc or | |
2662 | frv), and so the most relaxed case (namely DEC Alpha) must be assumed outside | |
2663 | of arch-specific code. | |
2664 | ||
2665 | This means that it must be considered that the CPU will execute its instruction | |
2666 | stream in any order it feels like - or even in parallel - provided that if an | |
81fc6323 | 2667 | instruction in the stream depends on an earlier instruction, then that |
108b42b4 DH |
2668 | earlier instruction must be sufficiently complete[*] before the later |
2669 | instruction may proceed; in other words: provided that the appearance of | |
2670 | causality is maintained. | |
2671 | ||
2672 | [*] Some instructions have more than one effect - such as changing the | |
2673 | condition codes, changing registers or changing memory - and different | |
2674 | instructions may depend on different effects. | |
2675 | ||
2676 | A CPU may also discard any instruction sequence that winds up having no | |
2677 | ultimate effect. For example, if two adjacent instructions both load an | |
2678 | immediate value into the same register, the first may be discarded. | |
2679 | ||
2680 | ||
2681 | Similarly, it has to be assumed that compiler might reorder the instruction | |
2682 | stream in any way it sees fit, again provided the appearance of causality is | |
2683 | maintained. | |
2684 | ||
2685 | ||
2686 | ============================ | |
2687 | THE EFFECTS OF THE CPU CACHE | |
2688 | ============================ | |
2689 | ||
2690 | The way cached memory operations are perceived across the system is affected to | |
2691 | a certain extent by the caches that lie between CPUs and memory, and by the | |
2692 | memory coherence system that maintains the consistency of state in the system. | |
2693 | ||
2694 | As far as the way a CPU interacts with another part of the system through the | |
2695 | caches goes, the memory system has to include the CPU's caches, and memory | |
2696 | barriers for the most part act at the interface between the CPU and its cache | |
2697 | (memory barriers logically act on the dotted line in the following diagram): | |
2698 | ||
2699 | <--- CPU ---> : <----------- Memory -----------> | |
2700 | : | |
2701 | +--------+ +--------+ : +--------+ +-----------+ | |
2702 | | | | | : | | | | +--------+ | |
e0edc78f IM |
2703 | | CPU | | Memory | : | CPU | | | | | |
2704 | | Core |--->| Access |----->| Cache |<-->| | | | | |
108b42b4 | 2705 | | | | Queue | : | | | |--->| Memory | |
e0edc78f IM |
2706 | | | | | : | | | | | | |
2707 | +--------+ +--------+ : +--------+ | | | | | |
108b42b4 DH |
2708 | : | Cache | +--------+ |
2709 | : | Coherency | | |
2710 | : | Mechanism | +--------+ | |
2711 | +--------+ +--------+ : +--------+ | | | | | |
2712 | | | | | : | | | | | | | |
2713 | | CPU | | Memory | : | CPU | | |--->| Device | | |
e0edc78f IM |
2714 | | Core |--->| Access |----->| Cache |<-->| | | | |
2715 | | | | Queue | : | | | | | | | |
108b42b4 DH |
2716 | | | | | : | | | | +--------+ |
2717 | +--------+ +--------+ : +--------+ +-----------+ | |
2718 | : | |
2719 | : | |
2720 | ||
2721 | Although any particular load or store may not actually appear outside of the | |
2722 | CPU that issued it since it may have been satisfied within the CPU's own cache, | |
2723 | it will still appear as if the full memory access had taken place as far as the | |
2724 | other CPUs are concerned since the cache coherency mechanisms will migrate the | |
2725 | cacheline over to the accessing CPU and propagate the effects upon conflict. | |
2726 | ||
2727 | The CPU core may execute instructions in any order it deems fit, provided the | |
2728 | expected program causality appears to be maintained. Some of the instructions | |
2729 | generate load and store operations which then go into the queue of memory | |
2730 | accesses to be performed. The core may place these in the queue in any order | |
2731 | it wishes, and continue execution until it is forced to wait for an instruction | |
2732 | to complete. | |
2733 | ||
2734 | What memory barriers are concerned with is controlling the order in which | |
2735 | accesses cross from the CPU side of things to the memory side of things, and | |
2736 | the order in which the effects are perceived to happen by the other observers | |
2737 | in the system. | |
2738 | ||
2739 | [!] Memory barriers are _not_ needed within a given CPU, as CPUs always see | |
2740 | their own loads and stores as if they had happened in program order. | |
2741 | ||
2742 | [!] MMIO or other device accesses may bypass the cache system. This depends on | |
2743 | the properties of the memory window through which devices are accessed and/or | |
2744 | the use of any special device communication instructions the CPU may have. | |
2745 | ||
2746 | ||
2747 | CACHE COHERENCY | |
2748 | --------------- | |
2749 | ||
2750 | Life isn't quite as simple as it may appear above, however: for while the | |
2751 | caches are expected to be coherent, there's no guarantee that that coherency | |
2752 | will be ordered. This means that whilst changes made on one CPU will | |
2753 | eventually become visible on all CPUs, there's no guarantee that they will | |
2754 | become apparent in the same order on those other CPUs. | |
2755 | ||
2756 | ||
81fc6323 JP |
2757 | Consider dealing with a system that has a pair of CPUs (1 & 2), each of which |
2758 | has a pair of parallel data caches (CPU 1 has A/B, and CPU 2 has C/D): | |
108b42b4 DH |
2759 | |
2760 | : | |
2761 | : +--------+ | |
2762 | : +---------+ | | | |
2763 | +--------+ : +--->| Cache A |<------->| | | |
2764 | | | : | +---------+ | | | |
2765 | | CPU 1 |<---+ | | | |
2766 | | | : | +---------+ | | | |
2767 | +--------+ : +--->| Cache B |<------->| | | |
2768 | : +---------+ | | | |
2769 | : | Memory | | |
2770 | : +---------+ | System | | |
2771 | +--------+ : +--->| Cache C |<------->| | | |
2772 | | | : | +---------+ | | | |
2773 | | CPU 2 |<---+ | | | |
2774 | | | : | +---------+ | | | |
2775 | +--------+ : +--->| Cache D |<------->| | | |
2776 | : +---------+ | | | |
2777 | : +--------+ | |
2778 | : | |
2779 | ||
2780 | Imagine the system has the following properties: | |
2781 | ||
2782 | (*) an odd-numbered cache line may be in cache A, cache C or it may still be | |
2783 | resident in memory; | |
2784 | ||
2785 | (*) an even-numbered cache line may be in cache B, cache D or it may still be | |
2786 | resident in memory; | |
2787 | ||
2788 | (*) whilst the CPU core is interrogating one cache, the other cache may be | |
2789 | making use of the bus to access the rest of the system - perhaps to | |
2790 | displace a dirty cacheline or to do a speculative load; | |
2791 | ||
2792 | (*) each cache has a queue of operations that need to be applied to that cache | |
2793 | to maintain coherency with the rest of the system; | |
2794 | ||
2795 | (*) the coherency queue is not flushed by normal loads to lines already | |
2796 | present in the cache, even though the contents of the queue may | |
81fc6323 | 2797 | potentially affect those loads. |
108b42b4 DH |
2798 | |
2799 | Imagine, then, that two writes are made on the first CPU, with a write barrier | |
2800 | between them to guarantee that they will appear to reach that CPU's caches in | |
2801 | the requisite order: | |
2802 | ||
2803 | CPU 1 CPU 2 COMMENT | |
2804 | =============== =============== ======================================= | |
2805 | u == 0, v == 1 and p == &u, q == &u | |
2806 | v = 2; | |
81fc6323 | 2807 | smp_wmb(); Make sure change to v is visible before |
108b42b4 DH |
2808 | change to p |
2809 | <A:modify v=2> v is now in cache A exclusively | |
2810 | p = &v; | |
2811 | <B:modify p=&v> p is now in cache B exclusively | |
2812 | ||
2813 | The write memory barrier forces the other CPUs in the system to perceive that | |
2814 | the local CPU's caches have apparently been updated in the correct order. But | |
81fc6323 | 2815 | now imagine that the second CPU wants to read those values: |
108b42b4 DH |
2816 | |
2817 | CPU 1 CPU 2 COMMENT | |
2818 | =============== =============== ======================================= | |
2819 | ... | |
2820 | q = p; | |
2821 | x = *q; | |
2822 | ||
81fc6323 | 2823 | The above pair of reads may then fail to happen in the expected order, as the |
108b42b4 DH |
2824 | cacheline holding p may get updated in one of the second CPU's caches whilst |
2825 | the update to the cacheline holding v is delayed in the other of the second | |
2826 | CPU's caches by some other cache event: | |
2827 | ||
2828 | CPU 1 CPU 2 COMMENT | |
2829 | =============== =============== ======================================= | |
2830 | u == 0, v == 1 and p == &u, q == &u | |
2831 | v = 2; | |
2832 | smp_wmb(); | |
2833 | <A:modify v=2> <C:busy> | |
2834 | <C:queue v=2> | |
79afecfa | 2835 | p = &v; q = p; |
108b42b4 DH |
2836 | <D:request p> |
2837 | <B:modify p=&v> <D:commit p=&v> | |
e0edc78f | 2838 | <D:read p> |
108b42b4 DH |
2839 | x = *q; |
2840 | <C:read *q> Reads from v before v updated in cache | |
2841 | <C:unbusy> | |
2842 | <C:commit v=2> | |
2843 | ||
2844 | Basically, whilst both cachelines will be updated on CPU 2 eventually, there's | |
2845 | no guarantee that, without intervention, the order of update will be the same | |
2846 | as that committed on CPU 1. | |
2847 | ||
2848 | ||
2849 | To intervene, we need to interpolate a data dependency barrier or a read | |
2850 | barrier between the loads. This will force the cache to commit its coherency | |
2851 | queue before processing any further requests: | |
2852 | ||
2853 | CPU 1 CPU 2 COMMENT | |
2854 | =============== =============== ======================================= | |
2855 | u == 0, v == 1 and p == &u, q == &u | |
2856 | v = 2; | |
2857 | smp_wmb(); | |
2858 | <A:modify v=2> <C:busy> | |
2859 | <C:queue v=2> | |
3fda982c | 2860 | p = &v; q = p; |
108b42b4 DH |
2861 | <D:request p> |
2862 | <B:modify p=&v> <D:commit p=&v> | |
e0edc78f | 2863 | <D:read p> |
108b42b4 DH |
2864 | smp_read_barrier_depends() |
2865 | <C:unbusy> | |
2866 | <C:commit v=2> | |
2867 | x = *q; | |
2868 | <C:read *q> Reads from v after v updated in cache | |
2869 | ||
2870 | ||
2871 | This sort of problem can be encountered on DEC Alpha processors as they have a | |
2872 | split cache that improves performance by making better use of the data bus. | |
2873 | Whilst most CPUs do imply a data dependency barrier on the read when a memory | |
2874 | access depends on a read, not all do, so it may not be relied on. | |
2875 | ||
2876 | Other CPUs may also have split caches, but must coordinate between the various | |
3f6dee9b | 2877 | cachelets for normal memory accesses. The semantics of the Alpha removes the |
81fc6323 | 2878 | need for coordination in the absence of memory barriers. |
108b42b4 DH |
2879 | |
2880 | ||
2881 | CACHE COHERENCY VS DMA | |
2882 | ---------------------- | |
2883 | ||
2884 | Not all systems maintain cache coherency with respect to devices doing DMA. In | |
2885 | such cases, a device attempting DMA may obtain stale data from RAM because | |
2886 | dirty cache lines may be resident in the caches of various CPUs, and may not | |
2887 | have been written back to RAM yet. To deal with this, the appropriate part of | |
2888 | the kernel must flush the overlapping bits of cache on each CPU (and maybe | |
2889 | invalidate them as well). | |
2890 | ||
2891 | In addition, the data DMA'd to RAM by a device may be overwritten by dirty | |
2892 | cache lines being written back to RAM from a CPU's cache after the device has | |
81fc6323 JP |
2893 | installed its own data, or cache lines present in the CPU's cache may simply |
2894 | obscure the fact that RAM has been updated, until at such time as the cacheline | |
2895 | is discarded from the CPU's cache and reloaded. To deal with this, the | |
2896 | appropriate part of the kernel must invalidate the overlapping bits of the | |
108b42b4 DH |
2897 | cache on each CPU. |
2898 | ||
2899 | See Documentation/cachetlb.txt for more information on cache management. | |
2900 | ||
2901 | ||
2902 | CACHE COHERENCY VS MMIO | |
2903 | ----------------------- | |
2904 | ||
2905 | Memory mapped I/O usually takes place through memory locations that are part of | |
81fc6323 | 2906 | a window in the CPU's memory space that has different properties assigned than |
108b42b4 DH |
2907 | the usual RAM directed window. |
2908 | ||
2909 | Amongst these properties is usually the fact that such accesses bypass the | |
2910 | caching entirely and go directly to the device buses. This means MMIO accesses | |
2911 | may, in effect, overtake accesses to cached memory that were emitted earlier. | |
2912 | A memory barrier isn't sufficient in such a case, but rather the cache must be | |
2913 | flushed between the cached memory write and the MMIO access if the two are in | |
2914 | any way dependent. | |
2915 | ||
2916 | ||
2917 | ========================= | |
2918 | THE THINGS CPUS GET UP TO | |
2919 | ========================= | |
2920 | ||
2921 | A programmer might take it for granted that the CPU will perform memory | |
81fc6323 | 2922 | operations in exactly the order specified, so that if the CPU is, for example, |
108b42b4 DH |
2923 | given the following piece of code to execute: |
2924 | ||
9af194ce PM |
2925 | a = READ_ONCE(*A); |
2926 | WRITE_ONCE(*B, b); | |
2927 | c = READ_ONCE(*C); | |
2928 | d = READ_ONCE(*D); | |
2929 | WRITE_ONCE(*E, e); | |
108b42b4 | 2930 | |
81fc6323 | 2931 | they would then expect that the CPU will complete the memory operation for each |
108b42b4 DH |
2932 | instruction before moving on to the next one, leading to a definite sequence of |
2933 | operations as seen by external observers in the system: | |
2934 | ||
2935 | LOAD *A, STORE *B, LOAD *C, LOAD *D, STORE *E. | |
2936 | ||
2937 | ||
2938 | Reality is, of course, much messier. With many CPUs and compilers, the above | |
2939 | assumption doesn't hold because: | |
2940 | ||
2941 | (*) loads are more likely to need to be completed immediately to permit | |
2942 | execution progress, whereas stores can often be deferred without a | |
2943 | problem; | |
2944 | ||
2945 | (*) loads may be done speculatively, and the result discarded should it prove | |
2946 | to have been unnecessary; | |
2947 | ||
81fc6323 JP |
2948 | (*) loads may be done speculatively, leading to the result having been fetched |
2949 | at the wrong time in the expected sequence of events; | |
108b42b4 DH |
2950 | |
2951 | (*) the order of the memory accesses may be rearranged to promote better use | |
2952 | of the CPU buses and caches; | |
2953 | ||
2954 | (*) loads and stores may be combined to improve performance when talking to | |
2955 | memory or I/O hardware that can do batched accesses of adjacent locations, | |
2956 | thus cutting down on transaction setup costs (memory and PCI devices may | |
2957 | both be able to do this); and | |
2958 | ||
2959 | (*) the CPU's data cache may affect the ordering, and whilst cache-coherency | |
2960 | mechanisms may alleviate this - once the store has actually hit the cache | |
2961 | - there's no guarantee that the coherency management will be propagated in | |
2962 | order to other CPUs. | |
2963 | ||
2964 | So what another CPU, say, might actually observe from the above piece of code | |
2965 | is: | |
2966 | ||
2967 | LOAD *A, ..., LOAD {*C,*D}, STORE *E, STORE *B | |
2968 | ||
2969 | (Where "LOAD {*C,*D}" is a combined load) | |
2970 | ||
2971 | ||
2972 | However, it is guaranteed that a CPU will be self-consistent: it will see its | |
2973 | _own_ accesses appear to be correctly ordered, without the need for a memory | |
2974 | barrier. For instance with the following code: | |
2975 | ||
9af194ce PM |
2976 | U = READ_ONCE(*A); |
2977 | WRITE_ONCE(*A, V); | |
2978 | WRITE_ONCE(*A, W); | |
2979 | X = READ_ONCE(*A); | |
2980 | WRITE_ONCE(*A, Y); | |
2981 | Z = READ_ONCE(*A); | |
108b42b4 DH |
2982 | |
2983 | and assuming no intervention by an external influence, it can be assumed that | |
2984 | the final result will appear to be: | |
2985 | ||
2986 | U == the original value of *A | |
2987 | X == W | |
2988 | Z == Y | |
2989 | *A == Y | |
2990 | ||
2991 | The code above may cause the CPU to generate the full sequence of memory | |
2992 | accesses: | |
2993 | ||
2994 | U=LOAD *A, STORE *A=V, STORE *A=W, X=LOAD *A, STORE *A=Y, Z=LOAD *A | |
2995 | ||
2996 | in that order, but, without intervention, the sequence may have almost any | |
9af194ce PM |
2997 | combination of elements combined or discarded, provided the program's view |
2998 | of the world remains consistent. Note that READ_ONCE() and WRITE_ONCE() | |
2999 | are -not- optional in the above example, as there are architectures | |
3000 | where a given CPU might reorder successive loads to the same location. | |
3001 | On such architectures, READ_ONCE() and WRITE_ONCE() do whatever is | |
3002 | necessary to prevent this, for example, on Itanium the volatile casts | |
3003 | used by READ_ONCE() and WRITE_ONCE() cause GCC to emit the special ld.acq | |
3004 | and st.rel instructions (respectively) that prevent such reordering. | |
108b42b4 DH |
3005 | |
3006 | The compiler may also combine, discard or defer elements of the sequence before | |
3007 | the CPU even sees them. | |
3008 | ||
3009 | For instance: | |
3010 | ||
3011 | *A = V; | |
3012 | *A = W; | |
3013 | ||
3014 | may be reduced to: | |
3015 | ||
3016 | *A = W; | |
3017 | ||
9af194ce | 3018 | since, without either a write barrier or an WRITE_ONCE(), it can be |
2ecf8101 | 3019 | assumed that the effect of the storage of V to *A is lost. Similarly: |
108b42b4 DH |
3020 | |
3021 | *A = Y; | |
3022 | Z = *A; | |
3023 | ||
9af194ce PM |
3024 | may, without a memory barrier or an READ_ONCE() and WRITE_ONCE(), be |
3025 | reduced to: | |
108b42b4 DH |
3026 | |
3027 | *A = Y; | |
3028 | Z = Y; | |
3029 | ||
3030 | and the LOAD operation never appear outside of the CPU. | |
3031 | ||
3032 | ||
3033 | AND THEN THERE'S THE ALPHA | |
3034 | -------------------------- | |
3035 | ||
3036 | The DEC Alpha CPU is one of the most relaxed CPUs there is. Not only that, | |
3037 | some versions of the Alpha CPU have a split data cache, permitting them to have | |
81fc6323 | 3038 | two semantically-related cache lines updated at separate times. This is where |
108b42b4 DH |
3039 | the data dependency barrier really becomes necessary as this synchronises both |
3040 | caches with the memory coherence system, thus making it seem like pointer | |
3041 | changes vs new data occur in the right order. | |
3042 | ||
81fc6323 | 3043 | The Alpha defines the Linux kernel's memory barrier model. |
108b42b4 DH |
3044 | |
3045 | See the subsection on "Cache Coherency" above. | |
3046 | ||
0b6fa347 | 3047 | |
6a65d263 | 3048 | VIRTUAL MACHINE GUESTS |
3dbf0913 | 3049 | ---------------------- |
6a65d263 MT |
3050 | |
3051 | Guests running within virtual machines might be affected by SMP effects even if | |
3052 | the guest itself is compiled without SMP support. This is an artifact of | |
3053 | interfacing with an SMP host while running an UP kernel. Using mandatory | |
3054 | barriers for this use-case would be possible but is often suboptimal. | |
3055 | ||
3056 | To handle this case optimally, low-level virt_mb() etc macros are available. | |
3057 | These have the same effect as smp_mb() etc when SMP is enabled, but generate | |
0b6fa347 | 3058 | identical code for SMP and non-SMP systems. For example, virtual machine guests |
6a65d263 MT |
3059 | should use virt_mb() rather than smp_mb() when synchronizing against a |
3060 | (possibly SMP) host. | |
3061 | ||
3062 | These are equivalent to smp_mb() etc counterparts in all other respects, | |
3063 | in particular, they do not control MMIO effects: to control | |
3064 | MMIO effects, use mandatory barriers. | |
108b42b4 | 3065 | |
0b6fa347 | 3066 | |
90fddabf DH |
3067 | ============ |
3068 | EXAMPLE USES | |
3069 | ============ | |
3070 | ||
3071 | CIRCULAR BUFFERS | |
3072 | ---------------- | |
3073 | ||
3074 | Memory barriers can be used to implement circular buffering without the need | |
3075 | of a lock to serialise the producer with the consumer. See: | |
3076 | ||
3077 | Documentation/circular-buffers.txt | |
3078 | ||
3079 | for details. | |
3080 | ||
3081 | ||
108b42b4 DH |
3082 | ========== |
3083 | REFERENCES | |
3084 | ========== | |
3085 | ||
3086 | Alpha AXP Architecture Reference Manual, Second Edition (Sites & Witek, | |
3087 | Digital Press) | |
3088 | Chapter 5.2: Physical Address Space Characteristics | |
3089 | Chapter 5.4: Caches and Write Buffers | |
3090 | Chapter 5.5: Data Sharing | |
3091 | Chapter 5.6: Read/Write Ordering | |
3092 | ||
3093 | AMD64 Architecture Programmer's Manual Volume 2: System Programming | |
3094 | Chapter 7.1: Memory-Access Ordering | |
3095 | Chapter 7.4: Buffering and Combining Memory Writes | |
3096 | ||
3097 | IA-32 Intel Architecture Software Developer's Manual, Volume 3: | |
3098 | System Programming Guide | |
3099 | Chapter 7.1: Locked Atomic Operations | |
3100 | Chapter 7.2: Memory Ordering | |
3101 | Chapter 7.4: Serializing Instructions | |
3102 | ||
3103 | The SPARC Architecture Manual, Version 9 | |
3104 | Chapter 8: Memory Models | |
3105 | Appendix D: Formal Specification of the Memory Models | |
3106 | Appendix J: Programming with the Memory Models | |
3107 | ||
3108 | UltraSPARC Programmer Reference Manual | |
3109 | Chapter 5: Memory Accesses and Cacheability | |
3110 | Chapter 15: Sparc-V9 Memory Models | |
3111 | ||
3112 | UltraSPARC III Cu User's Manual | |
3113 | Chapter 9: Memory Models | |
3114 | ||
3115 | UltraSPARC IIIi Processor User's Manual | |
3116 | Chapter 8: Memory Models | |
3117 | ||
3118 | UltraSPARC Architecture 2005 | |
3119 | Chapter 9: Memory | |
3120 | Appendix D: Formal Specifications of the Memory Models | |
3121 | ||
3122 | UltraSPARC T1 Supplement to the UltraSPARC Architecture 2005 | |
3123 | Chapter 8: Memory Models | |
3124 | Appendix F: Caches and Cache Coherency | |
3125 | ||
3126 | Solaris Internals, Core Kernel Architecture, p63-68: | |
3127 | Chapter 3.3: Hardware Considerations for Locks and | |
3128 | Synchronization | |
3129 | ||
3130 | Unix Systems for Modern Architectures, Symmetric Multiprocessing and Caching | |
3131 | for Kernel Programmers: | |
3132 | Chapter 13: Other Memory Models | |
3133 | ||
3134 | Intel Itanium Architecture Software Developer's Manual: Volume 1: | |
3135 | Section 2.6: Speculation | |
3136 | Section 4.4: Memory Access |