| 1 | <HTML>
|
|---|
| 2 | <HEAD>
|
|---|
| 3 | <TITLE> Conservative GC Algorithmic Overview </TITLE>
|
|---|
| 4 | <AUTHOR> Hans-J. Boehm, Silicon Graphics</author>
|
|---|
| 5 | </HEAD>
|
|---|
| 6 | <BODY>
|
|---|
| 7 | <H1> <I>This is under construction</i> </h1>
|
|---|
| 8 | <H1> Conservative GC Algorithmic Overview </h1>
|
|---|
| 9 | <P>
|
|---|
| 10 | This is a description of the algorithms and data structures used in our
|
|---|
| 11 | conservative garbage collector. I expect the level of detail to increase
|
|---|
| 12 | with time. For a survey of GC algorithms, see for example
|
|---|
| 13 | <A HREF="ftp://ftp.cs.utexas.edu/pub/garbage/gcsurvey.ps"> Paul Wilson's
|
|---|
| 14 | excellent paper</a>. For an overview of the collector interface,
|
|---|
| 15 | see <A HREF="gcinterface.html">here</a>.
|
|---|
| 16 | <P>
|
|---|
| 17 | This description is targeted primarily at someone trying to understand the
|
|---|
| 18 | source code. It specifically refers to variable and function names.
|
|---|
| 19 | It may also be useful for understanding the algorithms at a higher level.
|
|---|
| 20 | <P>
|
|---|
| 21 | The description here assumes that the collector is used in default mode.
|
|---|
| 22 | In particular, we assume that it used as a garbage collector, and not just
|
|---|
| 23 | a leak detector. We initially assume that it is used in stop-the-world,
|
|---|
| 24 | non-incremental mode, though the presence of the incremental collector
|
|---|
| 25 | will be apparent in the design.
|
|---|
| 26 | We assume the default finalization model, but the code affected by that
|
|---|
| 27 | is very localized.
|
|---|
| 28 | <H2> Introduction </h2>
|
|---|
| 29 | The garbage collector uses a modified mark-sweep algorithm. Conceptually
|
|---|
| 30 | it operates roughly in four phases:
|
|---|
| 31 |
|
|---|
| 32 | <OL>
|
|---|
| 33 |
|
|---|
| 34 | <LI>
|
|---|
| 35 | <I>Preparation</i> Clear all mark bits, indicating that all objects
|
|---|
| 36 | are potentially unreachable.
|
|---|
| 37 |
|
|---|
| 38 | <LI>
|
|---|
| 39 | <I>Mark phase</i> Marks all objects that can be reachable via chains of
|
|---|
| 40 | pointers from variables. Normally the collector has no real information
|
|---|
| 41 | about the location of pointer variables in the heap, so it
|
|---|
| 42 | views all static data areas, stacks and registers as potentially containing
|
|---|
| 43 | containing pointers. Any bit patterns that represent addresses inside
|
|---|
| 44 | heap objects managed by the collector are viewed as pointers.
|
|---|
| 45 | Unless the client program has made heap object layout information
|
|---|
| 46 | available to the collector, any heap objects found to be reachable from
|
|---|
| 47 | variables are again scanned similarly.
|
|---|
| 48 |
|
|---|
| 49 | <LI>
|
|---|
| 50 | <I>Sweep phase</i> Scans the heap for inaccessible, and hence unmarked,
|
|---|
| 51 | objects, and returns them to an appropriate free list for reuse. This is
|
|---|
| 52 | not really a separate phase; even in non incremental mode this is operation
|
|---|
| 53 | is usually performed on demand during an allocation that discovers an empty
|
|---|
| 54 | free list. Thus the sweep phase is very unlikely to touch a page that
|
|---|
| 55 | would not have been touched shortly thereafter anyway.
|
|---|
| 56 |
|
|---|
| 57 | <LI>
|
|---|
| 58 | <I>Finalization phase</i> Unreachable objects which had been registered
|
|---|
| 59 | for finalization are enqueued for finalization outside the collector.
|
|---|
| 60 |
|
|---|
| 61 | </ol>
|
|---|
| 62 |
|
|---|
| 63 | <P>
|
|---|
| 64 | The remaining sections describe the memory allocation data structures,
|
|---|
| 65 | and then the last 3 collection phases in more detail. We conclude by
|
|---|
| 66 | outlining some of the additional features implemented in the collector.
|
|---|
| 67 |
|
|---|
| 68 | <H2>Allocation</h2>
|
|---|
| 69 | The collector includes its own memory allocator. The allocator obtains
|
|---|
| 70 | memory from the system in a platform-dependent way. Under UNIX, it
|
|---|
| 71 | uses either <TT>malloc</tt>, <TT>sbrk</tt>, or <TT>mmap</tt>.
|
|---|
| 72 | <P>
|
|---|
| 73 | Most static data used by the allocator, as well as that needed by the
|
|---|
| 74 | rest of the garbage collector is stored inside the
|
|---|
| 75 | <TT>_GC_arrays</tt> structure.
|
|---|
| 76 | This allows the garbage collector to easily ignore the collectors own
|
|---|
| 77 | data structures when it searches for root pointers. Other allocator
|
|---|
| 78 | and collector internal data structures are allocated dynamically
|
|---|
| 79 | with <TT>GC_scratch_alloc</tt>. <TT>GC_scratch_alloc</tt> does not
|
|---|
| 80 | allow for deallocation, and is therefore used only for permanent data
|
|---|
| 81 | structures.
|
|---|
| 82 | <P>
|
|---|
| 83 | The allocator allocates objects of different <I>kinds</i>.
|
|---|
| 84 | Different kinds are handled somewhat differently by certain parts
|
|---|
| 85 | of the garbage collector. Certain kinds are scanned for pointers,
|
|---|
| 86 | others are not. Some may have per-object type descriptors that
|
|---|
| 87 | determine pointer locations. Or a specific kind may correspond
|
|---|
| 88 | to one specific object layout. Two built-in kinds are uncollectable.
|
|---|
| 89 | One (<TT>STUBBORN</tt>) is immutable without special precautions.
|
|---|
| 90 | In spite of that, it is very likely that most applications currently
|
|---|
| 91 | use at most two kinds: <TT>NORMAL</tt> and <TT>PTRFREE</tt> objects.
|
|---|
| 92 | <P>
|
|---|
| 93 | The collector uses a two level allocator. A large block is defined to
|
|---|
| 94 | be one larger than half of <TT>HBLKSIZE</tt>, which is a power of 2,
|
|---|
| 95 | typically on the order of the page size.
|
|---|
| 96 | <P>
|
|---|
| 97 | Large block sizes are rounded up to
|
|---|
| 98 | the next multiple of <TT>HBLKSIZE</tt> and then allocated by
|
|---|
| 99 | <TT>GC_allochblk</tt>. This uses roughly what Paul Wilson has termed
|
|---|
| 100 | a "next fit" algorithm, i.e. first-fit with a rotating pointer.
|
|---|
| 101 | The implementation does check for a better fitting immediately
|
|---|
| 102 | adjacent block, which gives it somewhat better fragmentation characteristics.
|
|---|
| 103 | I'm now convinced it should use a best fit algorithm. The actual
|
|---|
| 104 | implementation of <TT>GC_allochblk</tt>
|
|---|
| 105 | is significantly complicated by black-listing issues
|
|---|
| 106 | (see below).
|
|---|
| 107 | <P>
|
|---|
| 108 | Small blocks are allocated in blocks of size <TT>HBLKSIZE</tt>.
|
|---|
| 109 | Each block is
|
|---|
| 110 | dedicated to only one object size and kind. The allocator maintains
|
|---|
| 111 | separate free lists for each size and kind of object.
|
|---|
| 112 | <P>
|
|---|
| 113 | In order to avoid allocating blocks for too many distinct object sizes,
|
|---|
| 114 | the collector normally does not directly allocate objects of every possible
|
|---|
| 115 | request size. Instead request are rounded up to one of a smaller number
|
|---|
| 116 | of allocated sizes, for which free lists are maintained. The exact
|
|---|
| 117 | allocated sizes are computed on demand, but subject to the constraint
|
|---|
| 118 | that they increase roughly in geometric progression. Thus objects
|
|---|
| 119 | requested early in the execution are likely to be allocated with exactly
|
|---|
| 120 | the requested size, subject to alignment constraints.
|
|---|
| 121 | See <TT>GC_init_size_map</tt> for details.
|
|---|
| 122 | <P>
|
|---|
| 123 | The actual size rounding operation during small object allocation is
|
|---|
| 124 | implemented as a table lookup in <TT>GC_size_map</tt>.
|
|---|
| 125 | <P>
|
|---|
| 126 | Both collector initialization and computation of allocated sizes are
|
|---|
| 127 | handled carefully so that they do not slow down the small object fast
|
|---|
| 128 | allocation path. An attempt to allocate before the collector is initialized,
|
|---|
| 129 | or before the appropriate <TT>GC_size_map</tt> entry is computed,
|
|---|
| 130 | will take the same path as an allocation attempt with an empty free list.
|
|---|
| 131 | This results in a call to the slow path code (<TT>GC_generic_malloc_inner</tt>)
|
|---|
| 132 | which performs the appropriate initialization checks.
|
|---|
| 133 | <P>
|
|---|
| 134 | In non-incremental mode, we make a decision about whether to garbage collect
|
|---|
| 135 | whenever an allocation would otherwise have failed with the current heap size.
|
|---|
| 136 | If the total amount of allocation since the last collection is less than
|
|---|
| 137 | the heap size divided by <TT>GC_free_space_divisor</tt>, we try to
|
|---|
| 138 | expand the heap. Otherwise, we initiate a garbage collection. This ensures
|
|---|
| 139 | that the amount of garbage collection work per allocated byte remains
|
|---|
| 140 | constant.
|
|---|
| 141 | <P>
|
|---|
| 142 | The above is in fat an oversimplification of the real heap expansion
|
|---|
| 143 | heuristic, which adjusts slightly for root size and certain kinds of
|
|---|
| 144 | fragmentation. In particular, programs with a large root set size and
|
|---|
| 145 | little live heap memory will expand the heap to amortize the cost of
|
|---|
| 146 | scanning the roots.
|
|---|
| 147 | <P>
|
|---|
| 148 | Versions 5.x of the collector actually collect more frequently in
|
|---|
| 149 | nonincremental mode. The large block allocator usually refuses to split
|
|---|
| 150 | large heap blocks once the garbage collection threshold is
|
|---|
| 151 | reached. This often has the effect of collecting well before the
|
|---|
| 152 | heap fills up, thus reducing fragmentation and working set size at the
|
|---|
| 153 | expense of GC time. 6.x will chose an intermediate strategy depending
|
|---|
| 154 | on how much large object allocation has taken place in the past.
|
|---|
| 155 | (If the collector is configured to unmap unused pages, versions 6.x
|
|---|
| 156 | will use the 5.x strategy.)
|
|---|
| 157 | <P>
|
|---|
| 158 | (It has been suggested that this should be adjusted so that we favor
|
|---|
| 159 | expansion if the resulting heap still fits into physical memory.
|
|---|
| 160 | In many cases, that would no doubt help. But it is tricky to do this
|
|---|
| 161 | in a way that remains robust if multiple application are contending
|
|---|
| 162 | for a single pool of physical memory.)
|
|---|
| 163 |
|
|---|
| 164 | <H2>Mark phase</h2>
|
|---|
| 165 |
|
|---|
| 166 | The marker maintains an explicit stack of memory regions that are known
|
|---|
| 167 | to be accessible, but that have not yet been searched for contained pointers.
|
|---|
| 168 | Each stack entry contains the starting address of the block to be scanned,
|
|---|
| 169 | as well as a descriptor of the block. If no layout information is
|
|---|
| 170 | available for the block, then the descriptor is simply a length.
|
|---|
| 171 | (For other possibilities, see <TT>gc_mark.h</tt>.)
|
|---|
| 172 | <P>
|
|---|
| 173 | At the beginning of the mark phase, all root segments are pushed on the
|
|---|
| 174 | stack by <TT>GC_push_roots</tt>. If <TT>ALL_INTERIOR_PTRS</tt> is not
|
|---|
| 175 | defined, then stack roots require special treatment. In this case, the
|
|---|
| 176 | normal marking code ignores interior pointers, but <TT>GC_push_all_stack</tt>
|
|---|
| 177 | explicitly checks for interior pointers and pushes descriptors for target
|
|---|
| 178 | objects.
|
|---|
| 179 | <P>
|
|---|
| 180 | The marker is structured to allow incremental marking.
|
|---|
| 181 | Each call to <TT>GC_mark_some</tt> performs a small amount of
|
|---|
| 182 | work towards marking the heap.
|
|---|
| 183 | It maintains
|
|---|
| 184 | explicit state in the form of <TT>GC_mark_state</tt>, which
|
|---|
| 185 | identifies a particular sub-phase. Some other pieces of state, most
|
|---|
| 186 | notably the mark stack, identify how much work remains to be done
|
|---|
| 187 | in each sub-phase. The normal progression of mark states for
|
|---|
| 188 | a stop-the-world collection is:
|
|---|
| 189 | <OL>
|
|---|
| 190 | <LI> <TT>MS_INVALID</tt> indicating that there may be accessible unmarked
|
|---|
| 191 | objects. In this case <TT>GC_objects_are_marked</tt> will simultaneously
|
|---|
| 192 | be false, so the mark state is advanced to
|
|---|
| 193 | <LI> <TT>MS_PUSH_UNCOLLECTABLE</tt> indicating that it suffices to push
|
|---|
| 194 | uncollectable objects, roots, and then mark everything reachable from them.
|
|---|
| 195 | <TT>Scan_ptr</tt> is advanced through the heap until all uncollectable
|
|---|
| 196 | objects are pushed, and objects reachable from them are marked.
|
|---|
| 197 | At that point, the next call to <TT>GC_mark_some</tt> calls
|
|---|
| 198 | <TT>GC_push_roots</tt> to push the roots. It the advances the
|
|---|
| 199 | mark state to
|
|---|
| 200 | <LI> <TT>MS_ROOTS_PUSHED</tt> asserting that once the mark stack is
|
|---|
| 201 | empty, all reachable objects are marked. Once in this state, we work
|
|---|
| 202 | only on emptying the mark stack. Once this is completed, the state
|
|---|
| 203 | changes to
|
|---|
| 204 | <LI> <TT>MS_NONE</tt> indicating that reachable objects are marked.
|
|---|
| 205 | </ol>
|
|---|
| 206 |
|
|---|
| 207 | The core mark routine <TT>GC_mark_from_mark_stack</tt>, is called
|
|---|
| 208 | repeatedly by several of the sub-phases when the mark stack starts to fill
|
|---|
| 209 | up. It is also called repeatedly in <TT>MS_ROOTS_PUSHED</tt> state
|
|---|
| 210 | to empty the mark stack.
|
|---|
| 211 | The routine is designed to only perform a limited amount of marking at
|
|---|
| 212 | each call, so that it can also be used by the incremental collector.
|
|---|
| 213 | It is fairly carefully tuned, since it usually consumes a large majority
|
|---|
| 214 | of the garbage collection time.
|
|---|
| 215 | <P>
|
|---|
| 216 | The marker correctly handles mark stack overflows. Whenever the mark stack
|
|---|
| 217 | overflows, the mark state is reset to <TT>MS_INVALID</tt>.
|
|---|
| 218 | Since there are already marked objects in the heap,
|
|---|
| 219 | this eventually forces a complete
|
|---|
| 220 | scan of the heap, searching for pointers, during which any unmarked objects
|
|---|
| 221 | referenced by marked objects are again pushed on the mark stack. This
|
|---|
| 222 | process is repeated until the mark phase completes without a stack overflow.
|
|---|
| 223 | Each time the stack overflows, an attempt is made to grow the mark stack.
|
|---|
| 224 | All pieces of the collector that push regions onto the mark stack have to be
|
|---|
| 225 | careful to ensure forward progress, even in case of repeated mark stack
|
|---|
| 226 | overflows. Every mark attempt results in additional marked objects.
|
|---|
| 227 | <P>
|
|---|
| 228 | Each mark stack entry is processed by examining all candidate pointers
|
|---|
| 229 | in the range described by the entry. If the region has no associated
|
|---|
| 230 | type information, then this typically requires that each 4-byte aligned
|
|---|
| 231 | quantity (8-byte aligned with 64-bit pointers) be considered a candidate
|
|---|
| 232 | pointer.
|
|---|
| 233 | <P>
|
|---|
| 234 | We determine whether a candidate pointer is actually the address of
|
|---|
| 235 | a heap block. This is done in the following steps:
|
|---|
| 236 | <NL>
|
|---|
| 237 | <LI> The candidate pointer is checked against rough heap bounds.
|
|---|
| 238 | These heap bounds are maintained such that all actual heap objects
|
|---|
| 239 | fall between them. In order to facilitate black-listing (see below)
|
|---|
| 240 | we also include address regions that the heap is likely to expand into.
|
|---|
| 241 | Most non-pointers fail this initial test.
|
|---|
| 242 | <LI> The candidate pointer is divided into two pieces; the most significant
|
|---|
| 243 | bits identify a <TT>HBLKSIZE</tt>-sized page in the address space, and
|
|---|
| 244 | the least significant bits specify an offset within that page.
|
|---|
| 245 | (A hardware page may actually consist of multiple such pages.
|
|---|
| 246 | HBLKSIZE is usually the page size divided by a small power of two.)
|
|---|
| 247 | <LI>
|
|---|
| 248 | The page address part of the candidate pointer is looked up in a
|
|---|
| 249 | <A HREF="tree.html">table</a>.
|
|---|
| 250 | Each table entry contains either 0, indicating that the page is not part
|
|---|
| 251 | of the garbage collected heap, a small integer <I>n</i>, indicating
|
|---|
| 252 | that the page is part of large object, starting at least <I>n</i> pages
|
|---|
| 253 | back, or a pointer to a descriptor for the page. In the first case,
|
|---|
| 254 | the candidate pointer i not a true pointer and can be safely ignored.
|
|---|
| 255 | In the last two cases, we can obtain a descriptor for the page containing
|
|---|
| 256 | the beginning of the object.
|
|---|
| 257 | <LI>
|
|---|
| 258 | The starting address of the referenced object is computed.
|
|---|
| 259 | The page descriptor contains the size of the object(s)
|
|---|
| 260 | in that page, the object kind, and the necessary mark bits for those
|
|---|
| 261 | objects. The size information can be used to map the candidate pointer
|
|---|
| 262 | to the object starting address. To accelerate this process, the page header
|
|---|
| 263 | also contains a pointer to a precomputed map of page offsets to displacements
|
|---|
| 264 | from the beginning of an object. The use of this map avoids a
|
|---|
| 265 | potentially slow integer remainder operation in computing the object
|
|---|
| 266 | start address.
|
|---|
| 267 | <LI>
|
|---|
| 268 | The mark bit for the target object is checked and set. If the object
|
|---|
| 269 | was previously unmarked, the object is pushed on the mark stack.
|
|---|
| 270 | The descriptor is read from the page descriptor. (This is computed
|
|---|
| 271 | from information <TT>GC_obj_kinds</tt> when the page is first allocated.)
|
|---|
| 272 | </nl>
|
|---|
| 273 | <P>
|
|---|
| 274 | At the end of the mark phase, mark bits for left-over free lists are cleared,
|
|---|
| 275 | in case a free list was accidentally marked due to a stray pointer.
|
|---|
| 276 |
|
|---|
| 277 | <H2>Sweep phase</h2>
|
|---|
| 278 |
|
|---|
| 279 | At the end of the mark phase, all blocks in the heap are examined.
|
|---|
| 280 | Unmarked large objects are immediately returned to the large object free list.
|
|---|
| 281 | Each small object page is checked to see if all mark bits are clear.
|
|---|
| 282 | If so, the entire page is returned to the large object free list.
|
|---|
| 283 | Small object pages containing some reachable object are queued for later
|
|---|
| 284 | sweeping.
|
|---|
| 285 | <P>
|
|---|
| 286 | This initial sweep pass touches only block headers, not
|
|---|
| 287 | the blocks themselves. Thus it does not require significant paging, even
|
|---|
| 288 | if large sections of the heap are not in physical memory.
|
|---|
| 289 | <P>
|
|---|
| 290 | Nonempty small object pages are swept when an allocation attempt
|
|---|
| 291 | encounters an empty free list for that object size and kind.
|
|---|
| 292 | Pages for the correct size and kind are repeatedly swept until at
|
|---|
| 293 | least one empty block is found. Sweeping such a page involves
|
|---|
| 294 | scanning the mark bit array in the page header, and building a free
|
|---|
| 295 | list linked through the first words in the objects themselves.
|
|---|
| 296 | This does involve touching the appropriate data page, but in most cases
|
|---|
| 297 | it will be touched only just before it is used for allocation.
|
|---|
| 298 | Hence any paging is essentially unavoidable.
|
|---|
| 299 | <P>
|
|---|
| 300 | Except in the case of pointer-free objects, we maintain the invariant
|
|---|
| 301 | that any object in a small object free list is cleared (except possibly
|
|---|
| 302 | for the link field). Thus it becomes the burden of the small object
|
|---|
| 303 | sweep routine to clear objects. This has the advantage that we can
|
|---|
| 304 | easily recover from accidentally marking a free list, though that could
|
|---|
| 305 | also be handled by other means. The collector currently spends a fair
|
|---|
| 306 | amount of time clearing objects, and this approach should probably be
|
|---|
| 307 | revisited.
|
|---|
| 308 | <P>
|
|---|
| 309 | In most configurations, we use specialized sweep routines to handle common
|
|---|
| 310 | small object sizes. Since we allocate one mark bit per word, it becomes
|
|---|
| 311 | easier to examine the relevant mark bits if the object size divides
|
|---|
| 312 | the word length evenly. We also suitably unroll the inner sweep loop
|
|---|
| 313 | in each case. (It is conceivable that profile-based procedure cloning
|
|---|
| 314 | in the compiler could make this unnecessary and counterproductive. I
|
|---|
| 315 | know of no existing compiler to which this applies.)
|
|---|
| 316 | <P>
|
|---|
| 317 | The sweeping of small object pages could be avoided completely at the expense
|
|---|
| 318 | of examining mark bits directly in the allocator. This would probably
|
|---|
| 319 | be more expensive, since each allocation call would have to reload
|
|---|
| 320 | a large amount of state (e.g. next object address to be swept, position
|
|---|
| 321 | in mark bit table) before it could do its work. The current scheme
|
|---|
| 322 | keeps the allocator simple and allows useful optimizations in the sweeper.
|
|---|
| 323 |
|
|---|
| 324 | <H2>Finalization</h2>
|
|---|
| 325 | Both <TT>GC_register_disappearing_link</tt> and
|
|---|
| 326 | <TT>GC_register_finalizer</tt> add the request to a corresponding hash
|
|---|
| 327 | table. The hash table is allocated out of collected memory, but
|
|---|
| 328 | the reference to the finalizable object is hidden from the collector.
|
|---|
| 329 | Currently finalization requests are processed non-incrementally at the
|
|---|
| 330 | end of a mark cycle.
|
|---|
| 331 | <P>
|
|---|
| 332 | The collector makes an initial pass over the table of finalizable objects,
|
|---|
| 333 | pushing the contents of unmarked objects onto the mark stack.
|
|---|
| 334 | After pushing each object, the marker is invoked to mark all objects
|
|---|
| 335 | reachable from it. The object itself is not explicitly marked.
|
|---|
| 336 | This assures that objects on which a finalizer depends are neither
|
|---|
| 337 | collected nor finalized.
|
|---|
| 338 | <P>
|
|---|
| 339 | If in the process of marking from an object the
|
|---|
| 340 | object itself becomes marked, we have uncovered
|
|---|
| 341 | a cycle involving the object. This usually results in a warning from the
|
|---|
| 342 | collector. Such objects are not finalized, since it may be
|
|---|
| 343 | unsafe to do so. See the more detailed
|
|---|
| 344 | <A HREF="finalization.html"> discussion of finalization semantics</a>.
|
|---|
| 345 | <P>
|
|---|
| 346 | Any objects remaining unmarked at the end of this process are added to
|
|---|
| 347 | a queue of objects whose finalizers can be run. Depending on collector
|
|---|
| 348 | configuration, finalizers are dequeued and run either implicitly during
|
|---|
| 349 | allocation calls, or explicitly in response to a user request.
|
|---|
| 350 | <P>
|
|---|
| 351 | The collector provides a mechanism for replacing the procedure that is
|
|---|
| 352 | used to mark through objects. This is used both to provide support for
|
|---|
| 353 | Java-style unordered finalization, and to ignore certain kinds of cycles,
|
|---|
| 354 | <I>e.g.</i> those arising from C++ implementations of virtual inheritance.
|
|---|
| 355 |
|
|---|
| 356 | <H2>Generational Collection and Dirty Bits</h2>
|
|---|
| 357 | We basically use the parallel and generational GC algorithm described in
|
|---|
| 358 | <A HREF="papers/pldi91.ps.gz">"Mostly Parallel Garbage Collection"</a>,
|
|---|
| 359 | by Boehm, Demers, and Shenker.
|
|---|
| 360 | <P>
|
|---|
| 361 | The most significant modification is that
|
|---|
| 362 | the collector always runs in the allocating thread.
|
|---|
| 363 | There is no separate garbage collector thread.
|
|---|
| 364 | If an allocation attempt either requests a large object, or encounters
|
|---|
| 365 | an empty small object free list, and notices that there is a collection
|
|---|
| 366 | in progress, it immediately performs a small amount of marking work
|
|---|
| 367 | as described above.
|
|---|
| 368 | <P>
|
|---|
| 369 | This change was made both because we wanted to easily accommodate
|
|---|
| 370 | single-threaded environments, and because a separate GC thread requires
|
|---|
| 371 | very careful control over the scheduler to prevent the mutator from
|
|---|
| 372 | out-running the collector, and hence provoking unneeded heap growth.
|
|---|
| 373 | <P>
|
|---|
| 374 | In incremental mode, the heap is always expanded when we encounter
|
|---|
| 375 | insufficient space for an allocation. Garbage collection is triggered
|
|---|
| 376 | whenever we notice that more than
|
|---|
| 377 | <TT>GC_heap_size</tt>/2 * <TT>GC_free_space_divisor</tt>
|
|---|
| 378 | bytes of allocation have taken place.
|
|---|
| 379 | After <TT>GC_full_freq</tt> minor collections a major collection
|
|---|
| 380 | is started.
|
|---|
| 381 | <P>
|
|---|
| 382 | All collections initially run interrupted until a predetermined
|
|---|
| 383 | amount of time (50 msecs by default) has expired. If this allows
|
|---|
| 384 | the collection to complete entirely, we can avoid correcting
|
|---|
| 385 | for data structure modifications during the collection. If it does
|
|---|
| 386 | not complete, we return control to the mutator, and perform small
|
|---|
| 387 | amounts of additional GC work during those later allocations that
|
|---|
| 388 | cannot be satisfied from small object free lists. When marking completes,
|
|---|
| 389 | the set of modified pages is retrieved, and we mark once again from
|
|---|
| 390 | marked objects on those pages, this time with the mutator stopped.
|
|---|
| 391 | <P>
|
|---|
| 392 | We keep track of modified pages using one of three distinct mechanisms:
|
|---|
| 393 | <OL>
|
|---|
| 394 | <LI>
|
|---|
| 395 | Through explicit mutator cooperation. Currently this requires
|
|---|
| 396 | the use of <TT>GC_malloc_stubborn</tt>.
|
|---|
| 397 | <LI>
|
|---|
| 398 | By write-protecting physical pages and catching write faults. This is
|
|---|
| 399 | implemented for many Unix-like systems and for win32. It is not possible
|
|---|
| 400 | in a few environments.
|
|---|
| 401 | <LI>
|
|---|
| 402 | By retrieving dirty bit information from /proc. (Currently only Sun's
|
|---|
| 403 | Solaris supports this. Though this is considerably cleaner, performance
|
|---|
| 404 | may actually be better with mprotect and signals.)
|
|---|
| 405 | </ol>
|
|---|
| 406 |
|
|---|
| 407 | <H2>Thread support</h2>
|
|---|
| 408 | We support several different threading models. Unfortunately Pthreads,
|
|---|
| 409 | the only reasonably well standardized thread model, supports too narrow
|
|---|
| 410 | an interface for conservative garbage collection. There appears to be
|
|---|
| 411 | no portable way to allow the collector to coexist with various Pthreads
|
|---|
| 412 | implementations. Hence we currently support only a few of the more
|
|---|
| 413 | common Pthreads implementations.
|
|---|
| 414 | <P>
|
|---|
| 415 | In particular, it is very difficult for the collector to stop all other
|
|---|
| 416 | threads in the system and examine the register contents. This is currently
|
|---|
| 417 | accomplished with very different mechanisms for different Pthreads
|
|---|
| 418 | implementations. The Solaris implementation temporarily disables much
|
|---|
| 419 | of the user-level threads implementation by stopping kernel-level threads
|
|---|
| 420 | ("lwp"s). The Irix implementation sends signals to individual Pthreads
|
|---|
| 421 | and has them wait in the signal handler. The Linux implementation
|
|---|
| 422 | is similar in spirit to the Irix one.
|
|---|
| 423 | <P>
|
|---|
| 424 | The Irix implementation uses
|
|---|
| 425 | only documented Pthreads calls, but relies on extensions to their semantics,
|
|---|
| 426 | notably the use of mutexes and condition variables from signal
|
|---|
| 427 | handlers. The Linux implementation should be far closer to
|
|---|
| 428 | portable, though impirically it is not completely portable.
|
|---|
| 429 | <P>
|
|---|
| 430 | All implementations must
|
|---|
| 431 | intercept thread creation and a few other thread-specific calls to allow
|
|---|
| 432 | enumeration of threads and location of thread stacks. This is current
|
|---|
| 433 | accomplished with <TT># define</tt>'s in <TT>gc.h</tt>, or optionally
|
|---|
| 434 | by using ld's function call wrapping mechanism under Linux.
|
|---|
| 435 | <P>
|
|---|
| 436 | Comments are appreciated. Please send mail to
|
|---|
| 437 | <A HREF="mailto:boehm@acm.org"><TT>boehm@acm.org</tt></a>
|
|---|
| 438 | </body>
|
|---|
