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lru_cache.h
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1 /*
2  lru_cache.c
3 
4  This file is part of DRBD by Philipp Reisner and Lars Ellenberg.
5 
6  Copyright (C) 2003-2008, LINBIT Information Technologies GmbH.
7  Copyright (C) 2003-2008, Philipp Reisner <[email protected]>.
8  Copyright (C) 2003-2008, Lars Ellenberg <[email protected]>.
9 
10  drbd is free software; you can redistribute it and/or modify
11  it under the terms of the GNU General Public License as published by
12  the Free Software Foundation; either version 2, or (at your option)
13  any later version.
14 
15  drbd is distributed in the hope that it will be useful,
16  but WITHOUT ANY WARRANTY; without even the implied warranty of
17  MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the
18  GNU General Public License for more details.
19 
20  You should have received a copy of the GNU General Public License
21  along with drbd; see the file COPYING. If not, write to
22  the Free Software Foundation, 675 Mass Ave, Cambridge, MA 02139, USA.
23 
24  */
25 
26 #ifndef LRU_CACHE_H
27 #define LRU_CACHE_H
28 
29 #include <linux/list.h>
30 #include <linux/slab.h>
31 #include <linux/bitops.h>
32 #include <linux/string.h> /* for memset */
33 #include <linux/seq_file.h>
34 
35 /*
36 This header file (and its .c file; kernel-doc of functions see there)
37  define a helper framework to easily keep track of index:label associations,
38  and changes to an "active set" of objects, as well as pending transactions,
39  to persistently record those changes.
40 
41  We use an LRU policy if it is necessary to "cool down" a region currently in
42  the active set before we can "heat" a previously unused region.
43 
44  Because of this later property, it is called "lru_cache".
45  As it actually Tracks Objects in an Active SeT, we could also call it
46  toast (incidentally that is what may happen to the data on the
47  backend storage uppon next resync, if we don't get it right).
48 
49 What for?
50 
51 We replicate IO (more or less synchronously) to local and remote disk.
52 
53 For crash recovery after replication node failure,
54  we need to resync all regions that have been target of in-flight WRITE IO
55  (in use, or "hot", regions), as we don't know wether or not those WRITEs have
56  made it to stable storage.
57 
58  To avoid a "full resync", we need to persistently track these regions.
59 
60  This is known as "write intent log", and can be implemented as on-disk
61  (coarse or fine grained) bitmap, or other meta data.
62 
63  To avoid the overhead of frequent extra writes to this meta data area,
64  usually the condition is softened to regions that _may_ have been target of
65  in-flight WRITE IO, e.g. by only lazily clearing the on-disk write-intent
66  bitmap, trading frequency of meta data transactions against amount of
67  (possibly unnecessary) resync traffic.
68 
69  If we set a hard limit on the area that may be "hot" at any given time, we
70  limit the amount of resync traffic needed for crash recovery.
71 
72 For recovery after replication link failure,
73  we need to resync all blocks that have been changed on the other replica
74  in the mean time, or, if both replica have been changed independently [*],
75  all blocks that have been changed on either replica in the mean time.
76  [*] usually as a result of a cluster split-brain and insufficient protection.
77  but there are valid use cases to do this on purpose.
78 
79  Tracking those blocks can be implemented as "dirty bitmap".
80  Having it fine-grained reduces the amount of resync traffic.
81  It should also be persistent, to allow for reboots (or crashes)
82  while the replication link is down.
83 
84 There are various possible implementations for persistently storing
85 write intent log information, three of which are mentioned here.
86 
87 "Chunk dirtying"
88  The on-disk "dirty bitmap" may be re-used as "write-intent" bitmap as well.
89  To reduce the frequency of bitmap updates for write-intent log purposes,
90  one could dirty "chunks" (of some size) at a time of the (fine grained)
91  on-disk bitmap, while keeping the in-memory "dirty" bitmap as clean as
92  possible, flushing it to disk again when a previously "hot" (and on-disk
93  dirtied as full chunk) area "cools down" again (no IO in flight anymore,
94  and none expected in the near future either).
95 
96 "Explicit (coarse) write intent bitmap"
97  An other implementation could chose a (probably coarse) explicit bitmap,
98  for write-intent log purposes, additionally to the fine grained dirty bitmap.
99 
100 "Activity log"
101  Yet an other implementation may keep track of the hot regions, by starting
102  with an empty set, and writing down a journal of region numbers that have
103  become "hot", or have "cooled down" again.
104 
105  To be able to use a ring buffer for this journal of changes to the active
106  set, we not only record the actual changes to that set, but also record the
107  not changing members of the set in a round robin fashion. To do so, we use a
108  fixed (but configurable) number of slots which we can identify by index, and
109  associate region numbers (labels) with these indices.
110  For each transaction recording a change to the active set, we record the
111  change itself (index: -old_label, +new_label), and which index is associated
112  with which label (index: current_label) within a certain sliding window that
113  is moved further over the available indices with each such transaction.
114 
115  Thus, for crash recovery, if the ringbuffer is sufficiently large, we can
116  accurately reconstruct the active set.
117 
118  Sufficiently large depends only on maximum number of active objects, and the
119  size of the sliding window recording "index: current_label" associations within
120  each transaction.
121 
122  This is what we call the "activity log".
123 
124  Currently we need one activity log transaction per single label change, which
125  does not give much benefit over the "dirty chunks of bitmap" approach, other
126  than potentially less seeks.
127 
128  We plan to change the transaction format to support multiple changes per
129  transaction, which then would reduce several (disjoint, "random") updates to
130  the bitmap into one transaction to the activity log ring buffer.
131 */
132 
133 /* this defines an element in a tracked set
134  * .colision is for hash table lookup.
135  * When we process a new IO request, we know its sector, thus can deduce the
136  * region number (label) easily. To do the label -> object lookup without a
137  * full list walk, we use a simple hash table.
138  *
139  * .list is on one of three lists:
140  * in_use: currently in use (refcnt > 0, lc_number != LC_FREE)
141  * lru: unused but ready to be reused or recycled
142  * (lc_refcnt == 0, lc_number != LC_FREE),
143  * free: unused but ready to be recycled
144  * (lc_refcnt == 0, lc_number == LC_FREE),
145  *
146  * an element is said to be "in the active set",
147  * if either on "in_use" or "lru", i.e. lc_number != LC_FREE.
148  *
149  * DRBD currently (May 2009) only uses 61 elements on the resync lru_cache
150  * (total memory usage 2 pages), and up to 3833 elements on the act_log
151  * lru_cache, totalling ~215 kB for 64bit architecture, ~53 pages.
152  *
153  * We usually do not actually free these objects again, but only "recycle"
154  * them, as the change "index: -old_label, +LC_FREE" would need a transaction
155  * as well. Which also means that using a kmem_cache to allocate the objects
156  * from wastes some resources.
157  * But it avoids high order page allocations in kmalloc.
158  */
159 struct lc_element {
161  struct list_head list; /* LRU list or free list */
162  unsigned refcnt;
163  /* back "pointer" into lc_cache->element[index],
164  * for paranoia, and for "lc_element_to_index" */
165  unsigned lc_index;
166  /* if we want to track a larger set of objects,
167  * it needs to become arch independend u64 */
168  unsigned lc_number;
169 
170  /* special label when on free list */
171 #define LC_FREE (~0U)
172 };
173 
174 struct lru_cache {
175  /* the least recently used item is kept at lru->prev */
176  struct list_head lru;
177  struct list_head free;
179 
180  /* the pre-created kmem cache to allocate the objects from */
182 
183  /* size of tracked objects, used to memset(,0,) them in lc_reset */
184  size_t element_size;
185  /* offset of struct lc_element member in the tracked object */
186  size_t element_off;
187 
188  /* number of elements (indices) */
189  unsigned int nr_elements;
190  /* Arbitrary limit on maximum tracked objects. Practical limit is much
191  * lower due to allocation failures, probably. For typical use cases,
192  * nr_elements should be a few thousand at most.
193  * This also limits the maximum value of lc_element.lc_index, allowing the
194  * 8 high bits of .lc_index to be overloaded with flags in the future. */
195 #define LC_MAX_ACTIVE (1<<24)
196 
197  /* statistics */
198  unsigned used; /* number of lelements currently on in_use list */
199  unsigned long hits, misses, starving, dirty, changed;
200 
201  /* see below: flag-bits for lru_cache */
202  unsigned long flags;
203 
204  /* when changing the label of an index element */
205  unsigned int new_number;
206 
207  /* for paranoia when changing the label of an index element */
209 
210  void *lc_private;
211  const char *name;
212 
213  /* nr_elements there */
216 };
217 
218 
219 /* flag-bits for lru_cache */
220 enum {
221  /* debugging aid, to catch concurrent access early.
222  * user needs to guarantee exclusive access by proper locking! */
224  /* if we need to change the set, but currently there is a changing
225  * transaction pending, we are "dirty", and must deferr further
226  * changing requests */
228  /* if we need to change the set, but currently there is no free nor
229  * unused element available, we are "starving", and must not give out
230  * further references, to guarantee that eventually some refcnt will
231  * drop to zero and we will be able to make progress again, changing
232  * the set, writing the transaction.
233  * if the statistics say we are frequently starving,
234  * nr_elements is too small. */
236 };
237 #define LC_PARANOIA (1<<__LC_PARANOIA)
238 #define LC_DIRTY (1<<__LC_DIRTY)
239 #define LC_STARVING (1<<__LC_STARVING)
240 
241 extern struct lru_cache *lc_create(const char *name, struct kmem_cache *cache,
242  unsigned e_count, size_t e_size, size_t e_off);
243 extern void lc_reset(struct lru_cache *lc);
244 extern void lc_destroy(struct lru_cache *lc);
245 extern void lc_set(struct lru_cache *lc, unsigned int enr, int index);
246 extern void lc_del(struct lru_cache *lc, struct lc_element *element);
247 
248 extern struct lc_element *lc_try_get(struct lru_cache *lc, unsigned int enr);
249 extern struct lc_element *lc_find(struct lru_cache *lc, unsigned int enr);
250 extern struct lc_element *lc_get(struct lru_cache *lc, unsigned int enr);
251 extern unsigned int lc_put(struct lru_cache *lc, struct lc_element *e);
252 extern void lc_changed(struct lru_cache *lc, struct lc_element *e);
253 
254 struct seq_file;
255 extern size_t lc_seq_printf_stats(struct seq_file *seq, struct lru_cache *lc);
256 
257 extern void lc_seq_dump_details(struct seq_file *seq, struct lru_cache *lc, char *utext,
258  void (*detail) (struct seq_file *, struct lc_element *));
259 
267 static inline int lc_try_lock(struct lru_cache *lc)
268 {
269  return !test_and_set_bit(__LC_DIRTY, &lc->flags);
270 }
271 
276 static inline void lc_unlock(struct lru_cache *lc)
277 {
278  clear_bit(__LC_DIRTY, &lc->flags);
280 }
281 
282 static inline int lc_is_used(struct lru_cache *lc, unsigned int enr)
283 {
284  struct lc_element *e = lc_find(lc, enr);
285  return e && e->refcnt;
286 }
287 
288 #define lc_entry(ptr, type, member) \
289  container_of(ptr, type, member)
290 
291 extern struct lc_element *lc_element_by_index(struct lru_cache *lc, unsigned i);
292 extern unsigned int lc_index_of(struct lru_cache *lc, struct lc_element *e);
293 
294 #endif