| 105 |
|
* successful atomic operation per enq/deq pair. But it also |
| 106 |
|
* enables lower cost variants of queue maintenance mechanics. (A |
| 107 |
|
* variation of this idea applies even for non-dual queues that |
| 108 |
< |
* support deletion of embedded elements, such as |
| 108 |
> |
* support deletion of interior elements, such as |
| 109 |
|
* j.u.c.ConcurrentLinkedQueue.) |
| 110 |
|
* |
| 111 |
< |
* Once a node is matched, its item can never again change. We |
| 112 |
< |
* may thus arrange that the linked list of them contains a prefix |
| 113 |
< |
* of zero or more matched nodes, followed by a suffix of zero or |
| 114 |
< |
* more unmatched nodes. (Note that we allow both the prefix and |
| 115 |
< |
* suffix to be zero length, which in turn means that we do not |
| 116 |
< |
* use a dummy header.) If we were not concerned with either time |
| 117 |
< |
* or space efficiency, we could correctly perform enqueue and |
| 118 |
< |
* dequeue operations by traversing from a pointer to the initial |
| 119 |
< |
* node; CASing the item of the first unmatched node on match and |
| 120 |
< |
* CASing the next field of the trailing node on appends. While |
| 121 |
< |
* this would be a terrible idea in itself, it does have the |
| 122 |
< |
* benefit of not requiring ANY atomic updates on head/tail |
| 123 |
< |
* fields. |
| 111 |
> |
* Once a node is matched, its match status can never again |
| 112 |
> |
* change. We may thus arrange that the linked list of them |
| 113 |
> |
* contain a prefix of zero or more matched nodes, followed by a |
| 114 |
> |
* suffix of zero or more unmatched nodes. (Note that we allow |
| 115 |
> |
* both the prefix and suffix to be zero length, which in turn |
| 116 |
> |
* means that we do not use a dummy header.) If we were not |
| 117 |
> |
* concerned with either time or space efficiency, we could |
| 118 |
> |
* correctly perform enqueue and dequeue operations by traversing |
| 119 |
> |
* from a pointer to the initial node; CASing the item of the |
| 120 |
> |
* first unmatched node on match and CASing the next field of the |
| 121 |
> |
* trailing node on appends. (Plus some special-casing when |
| 122 |
> |
* initially empty). While this would be a terrible idea in |
| 123 |
> |
* itself, it does have the benefit of not requiring ANY atomic |
| 124 |
> |
* updates on head/tail fields. |
| 125 |
|
* |
| 126 |
|
* We introduce here an approach that lies between the extremes of |
| 127 |
< |
* never versus always updating queue (head and tail) pointers |
| 128 |
< |
* that reflects the tradeoff of sometimes requiring extra traversal |
| 129 |
< |
* steps to locate the first and/or last unmatched nodes, versus |
| 130 |
< |
* the reduced overhead and contention of fewer updates to queue |
| 131 |
< |
* pointers. For example, a possible snapshot of a queue is: |
| 127 |
> |
* never versus always updating queue (head and tail) pointers. |
| 128 |
> |
* This offers a tradeoff between sometimes requiring extra |
| 129 |
> |
* traversal steps to locate the first and/or last unmatched |
| 130 |
> |
* nodes, versus the reduced overhead and contention of fewer |
| 131 |
> |
* updates to queue pointers. For example, a possible snapshot of |
| 132 |
> |
* a queue is: |
| 133 |
|
* |
| 134 |
|
* head tail |
| 135 |
|
* | | |
| 141 |
|
* similarly for "tail") is an empirical matter. We have found |
| 142 |
|
* that using very small constants in the range of 1-3 work best |
| 143 |
|
* over a range of platforms. Larger values introduce increasing |
| 144 |
< |
* costs of cache misses and risks of long traversal chains. |
| 144 |
> |
* costs of cache misses and risks of long traversal chains, while |
| 145 |
> |
* smaller values increase CAS contentiona and overhead. |
| 146 |
|
* |
| 147 |
|
* Dual queues with slack differ from plain M&S dual queues by |
| 148 |
|
* virtue of only sometimes updating head or tail pointers when |
| 206 |
|
* additional GC bookkeeping ("write barriers") that are sometimes |
| 207 |
|
* more costly than the writes themselves because of contention). |
| 208 |
|
* |
| 209 |
< |
* Removal of internal nodes (due to timed out or interrupted |
| 209 |
> |
* Removal of interior nodes (due to timed out or interrupted |
| 210 |
|
* waits, or calls to remove or Iterator.remove) uses a scheme |
| 211 |
< |
* roughly similar to that in Scherer, Lea, and Scott |
| 211 |
> |
* roughly similar to that in Scherer, Lea, and Scott's |
| 212 |
|
* SynchronousQueue. Given a predecessor, we can unsplice any node |
| 213 |
|
* except the (actual) tail of the queue. To avoid build-up of |
| 214 |
|
* cancelled trailing nodes, upon a request to remove a trailing |
| 215 |
< |
* node, it is placed in field "cleanMe" to be unspliced later. |
| 215 |
> |
* node, it is placed in field "cleanMe" to be unspliced upon the |
| 216 |
> |
* next call to unsplice any other node. Situations needing such |
| 217 |
> |
* mechanics are not common but do occur in practice; for example |
| 218 |
> |
* when an unbounded series of short timed calls to poll |
| 219 |
> |
* repeatedly time out but never otherwise fall off the list |
| 220 |
> |
* because of an untimed call to take at the front of the |
| 221 |
> |
* queue. (Note that maintaining field cleanMe does not otherwise |
| 222 |
> |
* much impact garbage retention even if never cleared by some |
| 223 |
> |
* other call because the held node will eventually either |
| 224 |
> |
* directly or indirectly lead to a self-link once off the list.) |
| 225 |
|
* |
| 226 |
|
* *** Overview of implementation *** |
| 227 |
|
* |
| 229 |
|
* slack of two. The slack value is hard-wired: a path greater |
| 230 |
|
* than one is naturally implemented by checking equality of |
| 231 |
|
* traversal pointers except when the list has only one element, |
| 232 |
< |
* in which case we keep max slack at one. Avoiding tracking |
| 233 |
< |
* explicit counts across situations slightly simplifies an |
| 232 |
> |
* in which case we keep target slack at one. Avoiding tracking |
| 233 |
> |
* explicit counts across method calls slightly simplifies an |
| 234 |
|
* already-messy implementation. Using randomization would |
| 235 |
|
* probably work better if there were a low-quality dirt-cheap |
| 236 |
|
* per-thread one available, but even ThreadLocalRandom is too |
| 237 |
|
* heavy for these purposes. |
| 238 |
|
* |
| 239 |
< |
* With such a small slack value, path short-circuiting is rarely |
| 240 |
< |
* worthwhile. However, it is used (in awaitMatch) immediately |
| 241 |
< |
* before a waiting thread starts to block, as a final bit of |
| 242 |
< |
* helping at a point when contention with others is extremely |
| 243 |
< |
* unlikely (since if other threads that could release it are |
| 244 |
< |
* operating, then the current thread wouldn't be blocking). |
| 239 |
> |
* With such a small target slack value, it is rarely worthwhile |
| 240 |
> |
* to augment this with path short-circuiting; i.e., unsplicing |
| 241 |
> |
* nodes between head and the first unmatched node, or similarly |
| 242 |
> |
* for tail, rather than advancing head or tail proper. However, |
| 243 |
> |
* it is used (in awaitMatch) immediately before a waiting thread |
| 244 |
> |
* starts to block, as a final bit of helping at a point when |
| 245 |
> |
* contention with others is extremely unlikely (since if other |
| 246 |
> |
* threads that could release it are operating, then the current |
| 247 |
> |
* thread wouldn't be blocking). |
| 248 |
> |
* |
| 249 |
> |
* We allow both the head and tail fields to be null before any |
| 250 |
> |
* nodes are enqueued; initializing upon first append. This |
| 251 |
> |
* simplifies some other logic, as well as providing more |
| 252 |
> |
* efficient explicit control paths instead of letting JVMs insert |
| 253 |
> |
* implicit NullPointerExceptions when they are null. While not |
| 254 |
> |
* currently fully implemented, we also leave open the possibility |
| 255 |
> |
* of re-nulling these fields when empty (which is is complicated |
| 256 |
> |
* to arrange, for little benefit.) |
| 257 |
|
* |
| 258 |
|
* All enqueue/dequeue operations are handled by the single method |
| 259 |
|
* "xfer" with parameters indicating whether to act as some form |
| 273 |
|
* case matching it and returning, also if necessary updating |
| 274 |
|
* head to one past the matched node (or the node itself if the |
| 275 |
|
* list has no other unmatched nodes). If the CAS misses, then |
| 276 |
< |
* a retry loops until the slack is at most two. Traversals |
| 277 |
< |
* also check if the initial head is now off-list, in which |
| 278 |
< |
* case they start at the new head. |
| 276 |
> |
* a loop retries advancing head by two steps until either |
| 277 |
> |
* success or the slack is at most two. By requiring that each |
| 278 |
> |
* attempt advances head by two (if applicable), we ensure that |
| 279 |
> |
* the slack does not grow without bound. Traversals also check |
| 280 |
> |
* if the initial head is now off-list, in which case they |
| 281 |
> |
* start at the new head. |
| 282 |
|
* |
| 283 |
|
* If no candidates are found and the call was untimed |
| 284 |
|
* poll/offer, (argument "how" is NOW) return. |
| 533 |
|
/** |
| 534 |
|
* Tries to append node s as tail. |
| 535 |
|
* |
| 509 |
– |
* @param haveData true if appending in data mode |
| 536 |
|
* @param s the node to append |
| 537 |
+ |
* @param haveData true if appending in data mode |
| 538 |
|
* @return null on failure due to losing race with append in |
| 539 |
|
* different mode, else s's predecessor, or s itself if no |
| 540 |
|
* predecessor |
| 541 |
|
*/ |
| 542 |
|
private Node tryAppend(Node s, boolean haveData) { |
| 543 |
< |
for (Node t = tail, p = t;;) { // move p to actual tail and append |
| 543 |
> |
for (Node t = tail, p = t;;) { // move p to last node and append |
| 544 |
|
Node n, u; // temps for reads of next & tail |
| 545 |
|
if (p == null && (p = head) == null) { |
| 546 |
|
if (casHead(null, s)) |
| 548 |
|
} |
| 549 |
|
else if (p.cannotPrecede(haveData)) |
| 550 |
|
return null; // lost race vs opposite mode |
| 551 |
< |
else if ((n = p.next) != null) // Not tail; keep traversing |
| 551 |
> |
else if ((n = p.next) != null) // not last; keep traversing |
| 552 |
|
p = p != t && t != (u = tail) ? (t = u) : // stale tail |
| 553 |
|
(p != n) ? n : null; // restart if off list |
| 554 |
|
else if (!p.casNext(null, s)) |
| 555 |
|
p = p.next; // re-read on CAS failure |
| 556 |
|
else { |
| 557 |
< |
if (p != t) { // Update if slack now >= 2 |
| 557 |
> |
if (p != t) { // update if slack now >= 2 |
| 558 |
|
while ((tail != t || !casTail(t, s)) && |
| 559 |
|
(t = tail) != null && |
| 560 |
|
(s = t.next) != null && // advance and retry |
| 568 |
|
/** |
| 569 |
|
* Spins/yields/blocks until node s is matched or caller gives up. |
| 570 |
|
* |
| 571 |
< |
* @param pred the predecessor of s or s or null if none |
| 571 |
> |
* @param pred the predecessor of s, or s or null if none |
| 572 |
|
* @param s the waiting node |
| 573 |
|
* @param e the comparison value for checking match |
| 574 |
|
* @param how either SYNC or TIMEOUT |
| 781 |
|
s.forgetContents(); // clear unneeded fields |
| 782 |
|
/* |
| 783 |
|
* At any given time, exactly one node on list cannot be |
| 784 |
< |
* deleted -- the last inserted node. To accommodate this, if |
| 785 |
< |
* we cannot delete s, we save its predecessor as "cleanMe", |
| 784 |
> |
* unlinked -- the last inserted node. To accommodate this, if |
| 785 |
> |
* we cannot unlink s, we save its predecessor as "cleanMe", |
| 786 |
|
* processing the previously saved version first. Because only |
| 787 |
|
* one node in the list can have a null next, at least one of |
| 788 |
|
* node s or the node previously saved can always be |
