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内存分配器

golang内存分配最初是基于tcmalloc的，但是有很大的不同。tcmalloc文章：
参见：http://goog-perftools.sourceforge.net/doc/tcmalloc.html
翻译：https://blog.csdn.net/DERRANTCM/article/details/105342996

主分配器在大量页（runs of pages）中工作。将较小的分配大小（最大为32 kB，包括32 kB）舍入为大约70个大小类别之一，每个类别都有其自己的大小完全相同的空闲对象集。任何空闲的内存页都可以拆分为一个大小类别的对象集，然后使用空闲位图（free bitmap）进行管理。

分配器的数据结构为：

fixalloc：用于固定大小的堆外对象的空闲列表分配器，用于管理分配器使用的存储。
mheap：malloc堆，以页（8192字节）粒度进行管理。
mspan：由mheap管理的一系列使用中的页面。
mcentral：收集给定大小类的所有跨度。
mcache：具有可用空间的mspans的每P个缓存。
mstats：分配统计信息。

分配一个小对象沿用了高速缓存的层次结构：

1、将大小四舍五入为一个较小的类别，然后在此P的mcache中查看相应的mspan。扫描mspan的空闲位图以找到空闲位置（slot）。如果有空闲位置，分配它。这都可以在不获取锁的情况下完成。
2、如果mspan没有可用位置，则从mcentral的具有可用空间的所需size类的mspan列表中获取一个新的mspan。获得整个跨度（span）会摊销锁定mcentral的成本。
3、如果mcentral的mspan列表为空，从mheap获取一系列页以用于mspan。
4、如果mheap为空或没有足够大的页，则从操作系统中分配一组新的页（至少1MB）。分配大量页面将分摊与操作系统进行对话的成本。

清除mspan并释放对象沿用了类似的层次结构：

1、如果响应分配而清除了mspan，则将mspan返还到mcache以满足分配。
2、否则，如果mspan仍有已分配的对象，则将其放在mspan的size类别的mcentral空闲列表上。
3、否则，如果mspan中的所有对象都是空闲的，则mspan的页面将返回到mheap，并且mspan现在已失效。

分配和释放大对象直接使用mheap，而绕过mcache和mcentral。如果mspan.needzero为false，则mspan中的可用对象位置已被清零。否则，如果needzero为true，则在分配对象时将其清零。通过这种方式延迟归零有很多好处：

1、堆栈帧分配可以完全避免置零。
2、它具有更好的时间局部性，因为该程序可能即将写入内存。
3、我们不会将永远不会被重用的页面归零。

虚拟内存布局

堆由一组arena组成，这些arena在64位上为64MB，在32位（heapArenaBytes）上为4MB。每个arena的起始地址也与arena大小对齐。
每个arena都有一个关联的heapArena对象，该对象存储该arena的元数据：arena中所有字（word）的堆位图和arena中所有页的跨度（span）图。它们本身是堆外分配的。
由于arena是对齐的，因此可以将地址空间视为一系列arena帧（frame）。arena映射（mheap_.arenas）从arena帧号映射到*heapArena，对于不由Go堆支持的部分地址空间，映射为nil。arena映射的结构为两层数组，由“L1”arena映射和许多“ L2”arena映射组成；但是，由于arena很大，因此在许多体系结构上，arena映射都由一个大型L2映射组成。
arena地图覆盖了整个可用的地址空间，从而允许Go堆使用地址空间的任何部分。分配器尝试使arena保持连续，以便大跨度（以及大对象）可以跨越arena。

OS内存管理抽象层

在任何给定时间，运行时管理的地址空间区域可能处于四种状态之一：

1）无（None）——未保留和未映射，这是任何区域的默认状态。
2）保留（Reserved）——运行时拥有，但是访问它会导致故障。不计入进程的内存占用。
3）已准备（Prepared）——保留，意在不由物理内存支持（尽管OS可能会延迟实现）。可以有效过渡到就绪。在这样的区域中访问内存是不确定的（可能会出错，可能会返回意外的零等）。
4）就绪（Ready）——可以安全地访问。

这组状态对于支持所有当前受支持的平台而言绝对不是必需的。只需一个“无”，“保留”和“就绪”就可以解决问题。但是，“已准备”状态为我们提供了用于性能目的的灵活性。例如，在POSIX-y操作系统上，“保留”通常是设置了PROT_NONE的私有匿名mmap’d区域，要转换到“就绪”状态，需要设置PROT_READ | PROT_WRITE。但是，Prepared的规格不足使我们仅使用MADV_FREE从Ready过渡到Prepared。因此，在“准备好”状态下，我们可以提早设置一次权限位，我们可以有效地告诉操作系统，当我们严格不需要它们时，可以自由地将页面从我们手中夺走。

对于每个操作系统，都有一组通用的帮助程序，这些帮助程序在这些状态之间转换内存区域。帮助程序如下：

sysAlloc

sysAlloc将OS选择的内存区域从“无”转换为“就绪”。更具体地说，它从操作系统中获取大量的零位内存，通常大约为一百千字节或兆字节。该内存始终可以立即使用。

sysFree

sysFree将内存区域从任何状态转换为“无（Ready）”。因此，它无条件返回内存。如果在分配过程中检测到内存不足错误，或用于划分出地址空间的对齐部分，则使用此方法。仅当sysReserve始终返回与堆分配器的对齐限制对齐的内存区域时，如果sysFree是无操作的，这是可以的。

sysReserve

sysReserve将内存区域从“无（None）”转换为“保留（Reserved）”。它以这样一种方式保留地址空间，即在访问时（通过权限或未提交内存）会导致致命错误。因此，这种保留永远不会受到物理内存的支持。如果传递给它的指针为非nil，则调用者希望在那里保留，但是sysReserve仍然可以选择另一个位置（如果该位置不可用）。

注意：sysReserve返回OS对齐的内存，但是堆分配器可能使用更大的对齐方式，因此调用者必须小心地重新对齐sysReserve获得的内存。

sysMap

sysMap将内存区域从“保留（Reserved）”状态转换为“已准备（Prepared）”状态。它确保可以将存储区域有效地转换为“就绪（Ready）”。

sysUsed

sysUsed将内存区域从“已准备（Prepared）”过渡到“就绪（Ready）”。它通知操作系统需要内存区域，并确保可以安全地访问该区域。在没有明确的提交步骤和严格的过量提交限制的系统上，这通常是不操作的，例如，在Windows上至关重要。

sysUnused

sysUnused将内存区域从“就绪（Ready）”转换为“已准备（Prepared）”。它通知操作系统，不再需要支持该内存区域的物理页，并且可以将其重新用于其他目的。 sysUnused内存区域的内容被认为是没用的，在调用sysUsed之前，不得再次访问该区域。

sysFault

sysFault将内存区域从“就绪（Ready）”或“已准备（Prepared）”转换为“保留（Reserved）”。它标记了一个区域，以便在访问时总是会发生故障。仅用于调试运行时。

状态转换图

hint的选择

64位机器

在64位计算机上，我们选择以下hit因为：

1、从地址空间的中间开始，可以轻松扩展到连续范围，而无需运行其他映射。
2、这使Go堆地址在调试时更容易识别。
3、gccgo中的堆栈扫描仍然很保守，因此将地址与其他数据区分开很重要。

从0x00c0开始意味着有效的内存地址将从0x00c0、0x00c1 … n 小端开始，即c0 00，c1 00，…这些都不是有效的UTF-8序列，否则它们是尽可能远离ff（可能是一个公共字节）。如果失败，我们尝试其他0xXXc0地址。较早的尝试使用0x11f8导致线程分配期间OS X上的内存不足错误。 0x00c0导致与AddressSanitizer发生冲突，后者保留了最多0x0100的所有内存。这些选择减少了保守的垃圾收集器不收集内存的可能性，因为某些非指针内存块具有与内存地址匹配的位模式。

但是，在arm64上，我们忽略了上面的所有建议，并在0x40 << 32处分配，因为当使用具有3级转换缓冲区的4k页面时，
用户地址空间在darwin/arm64上被限制为39位，地址空间更小。

在AIX上，对于64位，mmaps从0x0A00000000000000开始。

32位机器

在32位计算机上，我们更加关注保持可用堆是连续的。
因此：

1、我们为所有的heapArena保留空间，这样它们就不会与heap交错。它们约为258MB，因此还算不错。（如果出现问题，我们可以在前面预留较小的空间。）
2、我们建议堆从二进制文件的末尾开始，因此我们有最大的机会保持其连续性。
3、我们尝试放出一个相当大的初始堆保留。

一些常量

var zerobase uintptr：所有0字节分配的基地址

源码

// Memory allocator.
//
// This was originally based on tcmalloc, but has diverged quite a bit.
// http://goog-perftools.sourceforge.net/doc/tcmalloc.html// The main allocator works in runs of pages.
// Small allocation sizes (up to and including 32 kB) are
// rounded to one of about 70 size classes, each of which
// has its own free set of objects of exactly that size.
// Any free page of memory can be split into a set of objects
// of one size class, which are then managed using a free bitmap.
//
// The allocator's data structures are:
//
//    fixalloc: a free-list allocator for fixed-size off-heap objects,
//        used to manage storage used by the allocator.
//    mheap: the malloc heap, managed at page (8192-byte) granularity.
//    mspan: a run of in-use pages managed by the mheap.
//    mcentral: collects all spans of a given size class.
//    mcache: a per-P cache of mspans with free space.
//    mstats: allocation statistics.
//
// Allocating a small object proceeds up a hierarchy of caches:
//
//    1. Round the size up to one of the small size classes
//       and look in the corresponding mspan in this P's mcache.
//       Scan the mspan's free bitmap to find a free slot.
//       If there is a free slot, allocate it.
//       This can all be done without acquiring a lock.
//
//    2. If the mspan has no free slots, obtain a new mspan
//       from the mcentral's list of mspans of the required size
//       class that have free space.
//       Obtaining a whole span amortizes the cost of locking
//       the mcentral.
//
//    3. If the mcentral's mspan list is empty, obtain a run
//       of pages from the mheap to use for the mspan.
//
//    4. If the mheap is empty or has no page runs large enough,
//       allocate a new group of pages (at least 1MB) from the
//       operating system. Allocating a large run of pages
//       amortizes the cost of talking to the operating system.
//
// Sweeping an mspan and freeing objects on it proceeds up a similar
// hierarchy:
//
//    1. If the mspan is being swept in response to allocation, it
//       is returned to the mcache to satisfy the allocation.
//
//    2. Otherwise, if the mspan still has allocated objects in it,
//       it is placed on the mcentral free list for the mspan's size
//       class.
//
//    3. Otherwise, if all objects in the mspan are free, the mspan's
//       pages are returned to the mheap and the mspan is now dead.
//
// Allocating and freeing a large object uses the mheap
// directly, bypassing the mcache and mcentral.
//
// If mspan.needzero is false, then free object slots in the mspan are
// already zeroed. Otherwise if needzero is true, objects are zeroed as
// they are allocated. There are various benefits to delaying zeroing
// this way:
//
//    1. Stack frame allocation can avoid zeroing altogether.
//
//    2. It exhibits better temporal locality, since the program is
//       probably about to write to the memory.
//
//    3. We don't zero pages that never get reused.// Virtual memory layout
//
// The heap consists of a set of arenas, which are 64MB on 64-bit and
// 4MB on 32-bit (heapArenaBytes). Each arena's start address is also
// aligned to the arena size.
//
// Each arena has an associated heapArena object that stores the
// metadata for that arena: the heap bitmap for all words in the arena
// and the span map for all pages in the arena. heapArena objects are
// themselves allocated off-heap.
//
// Since arenas are aligned, the address space can be viewed as a
// series of arena frames. The arena map (mheap_.arenas) maps from
// arena frame number to *heapArena, or nil for parts of the address
// space not backed by the Go heap. The arena map is structured as a
// two-level array consisting of a "L1" arena map and many "L2" arena
// maps; however, since arenas are large, on many architectures, the
// arena map consists of a single, large L2 map.
//
// The arena map covers the entire possible address space, allowing
// the Go heap to use any part of the address space. The allocator
// attempts to keep arenas contiguous so that large spans (and hence
// large objects) can cross arenas.
/*** 内存分配器** 这最初是基于tcmalloc的，但是有很大的不同。* 参见：http://goog-perftools.sourceforge.net/doc/tcmalloc.html* 翻译：https://blog.csdn.net/DERRANTCM/article/details/105342996** 主分配器在大量页（runs of pages）中工作。* 将较小的分配大小（最大为32 kB，包括32 kB）舍入为大约70个大小类别之一，每个类别都有其自己的大小完全相同的空闲对象集。* 任何空闲的内存页都可以拆分为一个大小类别的对象集，然后使用空闲位图（free bitmap）进行管理。** 分配器的数据结构为：** fixalloc：用于固定大小的堆外对象的空闲列表分配器，用于管理分配器使用的存储。* mheap：malloc堆，以页（8192字节）粒度进行管理。* mspan：由mheap管理的一系列使用中的页面。* mcentral：收集给定大小类的所有跨度。* mcache：具有可用空间的mspans的每P个缓存。* mstats：分配统计信息。** 分配一个小对象沿用了高速缓存的层次结构：** 1.将大小四舍五入为一个较小的类别，然后在此P的mcache中查看相应的mspan。* 扫描mspan的空闲位图以找到空闲位置（slot）。如果有空闲位置，分配它。这都可以在不获取锁的情况下完成。** 2.如果mspan没有可用位置，则从mcentral的具有可用空间的所需size类的mspan列表中获取一个新的mspan。* 获得整个跨度（span）会摊销锁定mcentral的成本。** 3.如果mcentral的mspan列表为空，从mheap获取一系列页以用于mspan。** 4.如果mheap为空或没有足够大的页，则从操作系统中分配一组新的页（至少1MB）。* 分配大量页面将分摊与操作系统进行对话的成本。** 清除mspan并释放对象沿用了类似的层次结构：** 1.如果响应分配而清除了mspan，则将mspan返还到mcache以满足分配。** 2.否则，如果mspan仍有已分配的对象，则将其放在mspan的size类别的mcentral空闲列表上。** 3.否则，如果mspan中的所有对象都是空闲的，则mspan的页面将返回到mheap，并且mspan现在已失效。** 分配和释放大对象直接使用mheap，而绕过mcache和mcentral。** 如果mspan.needzero为false，则mspan中的可用对象位置已被清零。否则，如果needzero为true，* 则在分配对象时将其清零。通过这种方式延迟归零有很多好处：** 1.堆栈帧分配可以完全避免置零。** 2.它具有更好的时间局部性，因为该程序可能即将写入内存。** 3.我们不会将永远不会被重用的页面归零。** 虚拟内存布局** 堆由一组arena组成，这些arena在64位上为64MB，在32位（heapArenaBytes）上为4MB。* 每个arena的起始地址也与arena大小对齐。** 每个arena都有一个关联的heapArena对象，该对象存储该arena的元数据：arena中所有字（word）的堆位图* 和arena中所有页的跨度（span）图。它们本身是堆外分配的。** 由于arena是对齐的，因此可以将地址空间视为一系列arena帧（frame）。arena映射（mheap_.arenas）* 从arena帧号映射到*heapArena，对于不由Go堆支持的部分地址空间，映射为nil。arena映射的结构为两层数组，* 由“L1”arena映射和许多“ L2”arena映射组成；但是，由于arena很大，因此在许多体系结构上，* arena映射都由一个大型L2映射组成。** arena地图覆盖了整个可用的地址空间，从而允许Go堆使用地址空间的任何部分。分配器尝试使arena保持连续，* 以便大跨度（以及大对象）可以跨越arena。**/package runtimeimport ("runtime/internal/atomic""runtime/internal/math""runtime/internal/sys""unsafe"
)const (debugMalloc = falsemaxTinySize   = _TinySizetinySizeClass = _TinySizeClassmaxSmallSize  = _MaxSmallSizepageShift = _PageShiftpageSize  = _PageSizepageMask  = _PageMask// By construction, single page spans of the smallest object class// have the most objects per span.// 通过构造，对象类别最小的单个页面跨度在每个跨度中具有最多的对象。// 每个跨度的最大多对象数maxObjsPerSpan = pageSize / 8concurrentSweep = _ConcurrentSweep_PageSize = 1 << _PageShift_PageMask = _PageSize - 1// _64bit = 1 on 64-bit systems, 0 on 32-bit systems// _64bit = 1在64位系统上，0在32位系统上_64bit = 1 << (^uintptr(0) >> 63) / 2// Tiny allocator parameters, see "Tiny allocator" comment in malloc.go.// Tiny分配器参数，请参阅malloc.go中的“Tiny分配器”注释。在mallocgc方法中_TinySize      = 16_TinySizeClass = int8(2)// FixAlloc的块大小_FixAllocChunk = 16 << 10 // Chunk size for FixAlloc// Per-P, per order stack segment cache size.// 每个P，每个order堆栈段的缓存大小。_StackCacheSize = 32 * 1024// Number of orders that get caching. Order 0 is FixedStack// and each successive order is twice as large.// We want to cache 2KB, 4KB, 8KB, and 16KB stacks. Larger stacks// will be allocated directly.// Since FixedStack is different on different systems, we// must vary NumStackOrders to keep the same maximum cached size.// 获得缓存的order数。Order 0是FixedStack，每个连续的order是其两倍。// 我们要缓存2KB，4KB，8KB和16KB堆栈。较大的堆栈将直接分配。// 由于FixedStack在不同的系统上是不同的，因此我们必须改变NumStackOrders以保持相同的最大缓存大小。// 下面是不同的操作系统对应的FixedStack和NumStackOrders的对应表//   OS               | FixedStack | NumStackOrders//   -----------------+------------+---------------//   linux/darwin/bsd | 2KB        | 4//   windows/32       | 4KB        | 3//   windows/64       | 8KB        | 2//   plan9            | 4KB        | 3_NumStackOrders = 4 - sys.PtrSize/4*sys.GoosWindows - 1*sys.GoosPlan9// heapAddrBits is the number of bits in a heap address. On// amd64, addresses are sign-extended beyond heapAddrBits. On// other arches, they are zero-extended.//// On most 64-bit platforms, we limit this to 48 bits based on a// combination of hardware and OS limitations.//// amd64 hardware limits addresses to 48 bits, sign-extended// to 64 bits. Addresses where the top 16 bits are not either// all 0 or all 1 are "non-canonical" and invalid. Because of// these "negative" addresses, we offset addresses by 1<<47// (arenaBaseOffset) on amd64 before computing indexes into// the heap arenas index. In 2017, amd64 hardware added// support for 57 bit addresses; however, currently only Linux// supports this extension and the kernel will never choose an// address above 1<<47 unless mmap is called with a hint// address above 1<<47 (which we never do).//// arm64 hardware (as of ARMv8) limits user addresses to 48// bits, in the range [0, 1<<48).//// ppc64, mips64, and s390x support arbitrary 64 bit addresses// in hardware. On Linux, Go leans on stricter OS limits. Based// on Linux's processor.h, the user address space is limited as// follows on 64-bit architectures:// heapAddrBits是堆地址中的位数。在amd64上，地址被符号扩展到heapAddrBits之外。在其他架构上，它们是零扩展的。//// 在大多数64位平台上，基于硬件和操作系统限制的组合，我们将其限制为48位。//// amd64硬件将地址限制为48位，符号扩展为64位。前16位不全为0或全为1的地址是“非规范的”且无效。// 由于存在这些“负”地址，因此在计算进入堆竞技场索引的索引之前，在amd64上将地址偏移1 << 47（arenaBaseOffset）。// 2017年，amd64硬件增加了对57位地址的支持;但是，当前只有Linux支持此扩展，内核将永远不会选择大于1 << 47的地址，// 除非调用mmap的提示地址大于1 << 47（我们从未这样做）。//// arm64硬件（自ARMv8起）将用户地址限制为48位，范围为[0，1 << 48）。//// ppc64，mips64和s390x在硬件中支持任意64位地址。在Linux上，Go依靠更严格的OS限制。// 基于Linux的processor.h，在64位体系结构上，用户地址空间受到如下限制：//// Architecture  Name              Maximum Value (exclusive)// ---------------------------------------------------------------------// amd64         TASK_SIZE_MAX     0x007ffffffff000 (47 bit addresses)// arm64         TASK_SIZE_64      0x01000000000000 (48 bit addresses)// ppc64{,le}    TASK_SIZE_USER64  0x00400000000000 (46 bit addresses)// mips64{,le}   TASK_SIZE64       0x00010000000000 (40 bit addresses)// s390x         TASK_SIZE         1<<64 (64 bit addresses)//// These limits may increase over time, but are currently at// most 48 bits except on s390x. On all architectures, Linux// starts placing mmap'd regions at addresses that are// significantly below 48 bits, so even if it's possible to// exceed Go's 48 bit limit, it's extremely unlikely in// practice.//// On 32-bit platforms, we accept the full 32-bit address// space because doing so is cheap.// mips32 only has access to the low 2GB of virtual memory, so// we further limit it to 31 bits.//// On darwin/arm64, although 64-bit pointers are presumably// available, pointers are truncated to 33 bits. Furthermore,// only the top 4 GiB of the address space are actually available// to the application, but we allow the whole 33 bits anyway for// simplicity.// TODO(mknyszek): Consider limiting it to 32 bits and using// arenaBaseOffset to offset into the top 4 GiB.//// WebAssembly currently has a limit of 4GB linear memory.// 这些限制可能会随时间增加，但目前最多为48位，但s390x除外。在所有体系结构上，// Linux都开始将mmap'd区域放置在明显低于48位的地址上，因此，即使有可能超过Go的48位限制，在实践中也极不可能。//// 在32位平台上，我们接受完整的32位地址空间，因为这样做很便宜。 mips32仅可以访问2GB的低虚拟内存，// 因此我们进一步将其限制为31位。//// 在darwin / arm64上，尽管大概可以使用64位指针，但指针会被截断为33位。此外，// 只有地址空间的前4个GiB实际上可供应用程序使用，但是为了简单起见，我们还是允许全部33位。// TODO（mknyszek）：考虑将其限制为32位，并使用arenaBaseOffset偏移到前4个GiB中。//// WebAssembly当前限制为4GB线性内存。// heapAddrBits：堆空间地址位数，间接表示了他可以支持的最大的内存空间heapAddrBits = (_64bit*(1-sys.GoarchWasm)*(1-sys.GoosDarwin*sys.GoarchArm64))*48 + (1-_64bit+sys.GoarchWasm)*(32-(sys.GoarchMips+sys.GoarchMipsle)) + 33*sys.GoosDarwin*sys.GoarchArm64// maxAlloc is the maximum size of an allocation. On 64-bit,// it's theoretically possible to allocate 1<<heapAddrBits bytes. On// 32-bit, however, this is one less than 1<<32 because the// number of bytes in the address space doesn't actually fit// in a uintptr.// maxAlloc是分配的最大大小。在64位上，理论上可以分配1 << heapAddrBits字节。// 但是，在32位上，这比1<<32小1，因为地址空间中的字节数实际上不适合uintptr。maxAlloc = (1 << heapAddrBits) - (1-_64bit)*1// The number of bits in a heap address, the size of heap// arenas, and the L1 and L2 arena map sizes are related by////   (1 << addr bits) = arena size * L1 entries * L2 entries//// Currently, we balance these as follows:// 堆地址中的位数，堆arena的大小以及L1和L2 arena映射的大小与//// 1 << addr位）=arena大小* L1条目* L2条目//// 目前，我们将这些平衡如下：////       Platform  Addr bits  Arena size  L1 entries   L2 entries// --------------  ---------  ----------  ----------  -----------//       */64-bit         48        64MB           1    4M (32MB)// windows/64-bit         48         4MB          64    1M  (8MB)//       */32-bit         32         4MB           1  1024  (4KB)//     */mips(le)         31         4MB           1   512  (2KB)// heapArenaBytes is the size of a heap arena. The heap// consists of mappings of size heapArenaBytes, aligned to// heapArenaBytes. The initial heap mapping is one arena.//// This is currently 64MB on 64-bit non-Windows and 4MB on// 32-bit and on Windows. We use smaller arenas on Windows// because all committed memory is charged to the process,// even if it's not touched. Hence, for processes with small// heaps, the mapped arena space needs to be commensurate.// This is particularly important with the race detector,// since it significantly amplifies the cost of committed// memory.// heapArenaBytes是堆arenas的大小。堆由大小为heapArenaBytes的映射组成，// 并与heapArenaBytes对齐。最初的堆映射是一个arenas。//// 当前在64位非Windows上为64MB，在32位和Windows上为4MB。我们在Windows上使用较小的arenas，// 因为所有已提交的内存都由进程负责，即使未涉及也是如此。因此，对于具有小堆的进程，映射的arenas空间需要相对应。// 这对于竞争检测器尤其重要，因为它会大大增加已提交内存的成本。heapArenaBytes = 1 << logHeapArenaBytes// logHeapArenaBytes is log_2 of heapArenaBytes. For clarity,// prefer using heapArenaBytes where possible (we need the// constant to compute some other constants).// logHeapArenaBytes是heapArenaBytes的log_2。为了清楚起见，// 最好在可能的地方使用heapArenaBytes（我们需要使用常量来计算其他常量）。logHeapArenaBytes = (6+20)*(_64bit*(1-sys.GoosWindows)*(1-sys.GoarchWasm)) + (2+20)*(_64bit*sys.GoosWindows) + (2+20)*(1-_64bit) + (2+20)*sys.GoarchWasm// heapArenaBitmapBytes is the size of each heap arena's bitmap.// heapArenaBitmapBytes是每个堆arena的位图大小。heapArenaBitmapBytes = heapArenaBytes / (sys.PtrSize * 8 / 2)// 每个arena所的页数pagesPerArena = heapArenaBytes / pageSize// arenaL1Bits is the number of bits of the arena number// covered by the first level arena map.//// This number should be small, since the first level arena// map requires PtrSize*(1<<arenaL1Bits) of space in the// binary's BSS. It can be zero, in which case the first level// index is effectively unused. There is a performance benefit// to this, since the generated code can be more efficient,// but comes at the cost of having a large L2 mapping.//// We use the L1 map on 64-bit Windows because the arena size// is small, but the address space is still 48 bits, and// there's a high cost to having a large L2.// arenaL1Bits是第一级arena映射覆盖的arena编号的位数。//// 这个数字应该很小，因为第一级arena映射在二进制文件的BSS中需要PtrSize*(1<<arenaL1Bits)空间。// 它可以为零，在这种情况下，第一级索引实际上未被使用。这会带来性能上的好处，// 因为生成的代码可以更高效，但是以拥有较大的L2映射为代价。//// 我们在64位Windows上使用L1映射，因为arena大小很小，但是地址空间仍然是48位，并且拥有大型L2的成本很高。arenaL1Bits = 6 * (_64bit * sys.GoosWindows)// arenaL2Bits is the number of bits of the arena number// covered by the second level arena index.//// The size of each arena map allocation is proportional to// 1<<arenaL2Bits, so it's important that this not be too// large. 48 bits leads to 32MB arena index allocations, which// is about the practical threshold.// arenaL2Bits是第二级arena索引覆盖的arena编号的位数。//// 每个arena映射分配的大小与1<<arenaL2Bits成正比，因此，不要太大也很重要。// 48位导致32MB arena索引分配，这大约是实际的阈值。arenaL2Bits = heapAddrBits - logHeapArenaBytes - arenaL1Bits// arenaL1Shift is the number of bits to shift an arena frame// number by to compute an index into the first level arena map.// arenaL1Shift是将arena帧号移位以计算进入第一级arena映射的索引的位数。arenaL1Shift = arenaL2Bits// arenaBits is the total bits in a combined arena map index.// This is split between the index into the L1 arena map and// the L2 arena map.// arenaBits是组合的arena映射索引中的总位。这在进入L1 arena映射和L2 arena映射的索引之间进行划分。arenaBits = arenaL1Bits + arenaL2Bits// arenaBaseOffset is the pointer value that corresponds to// index 0 in the heap arena map.//// On amd64, the address space is 48 bits, sign extended to 64// bits. This offset lets us handle "negative" addresses (or// high addresses if viewed as unsigned).//// On aix/ppc64, this offset allows to keep the heapAddrBits to// 48. Otherwize, it would be 60 in order to handle mmap addresses// (in range 0x0a00000000000000 - 0x0afffffffffffff). But in this// case, the memory reserved in (s *pageAlloc).init for chunks// is causing important slowdowns.//// On other platforms, the user address space is contiguous// and starts at 0, so no offset is necessary.// arenaBaseOffset是与堆arena映射中的索引0对应的指针值。//// 在amd64上，地址空间为48位，符号扩展为64位。此偏移量使我们可以处理“负”地址（如果视为无符号，则为高地址）。//// 在aix/ppc64上，此偏移量允许将heapAddrBits保持为48。否则，为了处理mmap地址//（范围为0x0a00000000000000-0x0afffffffffffffff），它将为60。但是在这种情况下，// (s*pageAlloc).init中为块保留的内存会导致严重的速度下降。//// 在其他平台上，用户地址空间是连续的，并且从0开始，因此不需要偏移量。arenaBaseOffset = sys.GoarchAmd64*(1<<47) + (^0x0a00000000000000+1)&uintptrMask*sys.GoosAix// Max number of threads to run garbage collection.// 2, 3, and 4 are all plausible maximums depending// on the hardware details of the machine. The garbage// collector scales well to 32 cpus.// 运行垃圾回收的最大线程数。 2、3和4都是合理的最大值，具体取决于机器的硬件细节。 垃圾收集器可以很好地扩展到32 cpus。_MaxGcproc = 32// minLegalPointer is the smallest possible legal pointer.// This is the smallest possible architectural page size,// since we assume that the first page is never mapped.//// This should agree with minZeroPage in the compiler.//// minLegalPointer是最小的合法指针。 这是可能的最小体系架构页大小，因为我们假设第一页从未映射过。//这应该与编译器中的minZeroPage一致。minLegalPointer uintptr = 4096
)// physPageSize is the size in bytes of the OS's physical pages.
// Mapping and unmapping operations must be done at multiples of
// physPageSize.
//
// This must be set by the OS init code (typically in osinit) before
// mallocinit.
//
// physPageSize是操作系统物理页面的大小（以字节为单位）。
// 映射和取消映射操作必须以physPageSize的倍数完成。
//
// 必须在mallocinit之前通过OS初始化代码（通常在osinit中）进行设置。
var physPageSize uintptr// physHugePageSize is the size in bytes of the OS's default physical huge
// page size whose allocation is opaque to the application. It is assumed
// and verified to be a power of two.
//
// If set, this must be set by the OS init code (typically in osinit) before
// mallocinit. However, setting it at all is optional, and leaving the default
// value is always safe (though potentially less efficient).
//
// Since physHugePageSize is always assumed to be a power of two,
// physHugePageShift is defined as physHugePageSize == 1 << physHugePageShift.
// The purpose of physHugePageShift is to avoid doing divisions in
// performance critical functions.
//
// physHugePageSize是操作系统默认物理大页面大小的大小（以字节为单位），
// 该大小对于应用程序是不透明的。 假定并验证为2的幂。
//
// 如果已设置，则必须在mallocinit之前通过OS初始化代码（通常在osinit中）进行设置。
// 但是，完全设置它是可选的，并且保留默认值始终是安全的（尽管可能会降低效率）。
//
// 由于physHugePageSize始终假定为2的幂，因此physHugePageShift定义为physHugePageSize == 1 << physHugePageShift。
// physHugePageShift的目的是避免对性能至关重要的功能进行划分。
var (physHugePageSize  uintptrphysHugePageShift uint
)// OS memory management abstraction layer
//
// Regions of the address space managed by the runtime may be in one of four
// states at any given time:
// 1) None - Unreserved and unmapped, the default state of any region.
// 2) Reserved - Owned by the runtime, but accessing it would cause a fault.
//               Does not count against the process' memory footprint.
// 3) Prepared - Reserved, intended not to be backed by physical memory (though
//               an OS may implement this lazily). Can transition efficiently to
//               Ready. Accessing memory in such a region is undefined (may
//               fault, may give back unexpected zeroes, etc.).
// 4) Ready - may be accessed safely.
//
// This set of states is more than is strictly necessary to support all the
// currently supported platforms. One could get by with just None, Reserved, and
// Ready. However, the Prepared state gives us flexibility for performance
// purposes. For example, on POSIX-y operating systems, Reserved is usually a
// private anonymous mmap'd region with PROT_NONE set, and to transition
// to Ready would require setting PROT_READ|PROT_WRITE. However the
// underspecification of Prepared lets us use just MADV_FREE to transition from
// Ready to Prepared. Thus with the Prepared state we can set the permission
// bits just once early on, we can efficiently tell the OS that it's free to
// take pages away from us when we don't strictly need them.
//
// For each OS there is a common set of helpers defined that transition
// memory regions between these states. The helpers are as follows:
//
// sysAlloc transitions an OS-chosen region of memory from None to Ready.
// More specifically, it obtains a large chunk of zeroed memory from the
// operating system, typically on the order of a hundred kilobytes
// or a megabyte. This memory is always immediately available for use.
//
// sysFree transitions a memory region from any state to None. Therefore, it
// returns memory unconditionally. It is used if an out-of-memory error has been
// detected midway through an allocation or to carve out an aligned section of
// the address space. It is okay if sysFree is a no-op only if sysReserve always
// returns a memory region aligned to the heap allocator's alignment
// restrictions.
//
// sysReserve transitions a memory region from None to Reserved. It reserves
// address space in such a way that it would cause a fatal fault upon access
// (either via permissions or not committing the memory). Such a reservation is
// thus never backed by physical memory.
// If the pointer passed to it is non-nil, the caller wants the
// reservation there, but sysReserve can still choose another
// location if that one is unavailable.
// NOTE: sysReserve returns OS-aligned memory, but the heap allocator
// may use larger alignment, so the caller must be careful to realign the
// memory obtained by sysReserve.
//
// sysMap transitions a memory region from Reserved to Prepared. It ensures the
// memory region can be efficiently transitioned to Ready.
//
// sysUsed transitions a memory region from Prepared to Ready. It notifies the
// operating system that the memory region is needed and ensures that the region
// may be safely accessed. This is typically a no-op on systems that don't have
// an explicit commit step and hard over-commit limits, but is critical on
// Windows, for example.
//
// sysUnused transitions a memory region from Ready to Prepared. It notifies the
// operating system that the physical pages backing this memory region are no
// longer needed and can be reused for other purposes. The contents of a
// sysUnused memory region are considered forfeit and the region must not be
// accessed again until sysUsed is called.
//
// sysFault transitions a memory region from Ready or Prepared to Reserved. It
// marks a region such that it will always fault if accessed. Used only for
// debugging the runtime.
/*** OS内存管理抽象层** 在任何给定时间，运行时管理的地址空间区域可能处于四种状态之一：* - 1）无（None）——未保留和未映射，这是任何区域的默认状态。* - 2）保留（Reserved）——运行时拥有，但是访问它会导致故障。不计入进程的内存占用。* - 3）已准备（Prepared）——保留，意在不由物理内存支持（尽管OS可能会延迟实现）。*      可以有效过渡到就绪。在这样的区域中访问内存是不确定的（可能会出错，可能会返回意外的零等）。* - 4）就绪（Ready）——可以安全地访问。** 这组状态对于支持所有当前受支持的平台而言绝对不是必需的。只需一个“无”，“保留”和“就绪”就可以解决问题。* 但是，“已准备”状态为我们提供了用于性能目的的灵活性。例如，在POSIX-y操作系统上，“保留”通常是设置了PROT_NONE的私有匿名mmap'd区域，* 要转换到“就绪”状态，需要设置PROT_READ | PROT_WRITE。但是，Prepared的规格不足使我们仅使用MADV_FREE从Ready过渡到Prepared。* 因此，在“准备好”状态下，我们可以提早设置一次权限位，我们可以有效地告诉操作系统，当我们严格不需要它们时，可以自由地将页面从我们手中夺走。** 对于每个操作系统，都有一组通用的帮助程序，这些帮助程序在这些状态之间转换内存区域。帮助程序如下：** sysAlloc* sysAlloc将OS选择的内存区域从“无”转换为“就绪”。更具体地说，它从操作系统中获取大量的零位内存，通常大约为一百千字节或兆字节。* 该内存始终可以立即使用。** sysFree* sysFree将内存区域从任何状态转换为“无（Ready）”。因此，它无条件返回内存。如果在分配过程中检测到内存不足错误，* 或用于划分出地址空间的对齐部分，则使用此方法。仅当sysReserve始终返回与堆分配器的对齐限制对齐的内存区域时，* 如果sysFree是无操作的，这是可以的。** sysReserve* sysReserve将内存区域从“无（None）”转换为“保留（Reserved）”。它以这样一种方式保留地址空间，* 即在访问时（通过权限或未提交内存）会导致致命错误。因此，这种保留永远不会受到物理内存的支持。如果传递给它的指针为非nil，* 则调用者希望在那里保留，但是sysReserve仍然可以选择另一个位置（如果该位置不可用）。** 注意：sysReserve返回OS对齐的内存，但是堆分配器可能使用更大的对齐方式，因此调用者必须小心地重新对齐sysReserve获得的内存。** sysMap* sysMap将内存区域从“保留（Reserved）”状态转换为“已准备（Prepared）”状态。它确保可以将存储区域有效地转换为“就绪（Ready）”。** sysUsed* sysUsed将内存区域从“已准备（Prepared）”过渡到“就绪（Ready）”。它通知操作系统需要内存区域，并确保可以安全地访问该区域。* 在没有明确的提交步骤和严格的过量提交限制的系统上，这通常是不操作的，例如，在Windows上至关重要。** sysUnused* sysUnused将内存区域从“就绪（Ready）”转换为“已准备（Prepared）”。它通知操作系统，不再需要支持该内存区域的物理页，* 并且可以将其重新用于其他目的。 sysUnused内存区域的内容被认为是没用的，在调用sysUsed之前，不得再次访问该区域。** sysFault* sysFault将内存区域从“就绪（Ready）”或“已准备（Prepared）”转换为“保留（Reserved）”。它标记了一个区域，* 以便在访问时总是会发生故障。仅用于调试运行时。**/
func mallocinit() {// 检查_TinySizeClass与_TinySize对应关系if class_to_size[_TinySizeClass] != _TinySize {throw("bad TinySizeClass")}// 确保映射到相同defer大小类别的defer arg大小也映射到相同的malloc大小类别。testdefersizes()// 判断heapArenaBitmapBytes是否是2的指数次方if heapArenaBitmapBytes&(heapArenaBitmapBytes-1) != 0 {// heapBits expects modular arithmetic on bitmap// addresses to work.// heapBits希望对位图地址进行模块化算术运算。throw("heapArenaBitmapBytes not a power of 2")}// Copy class sizes out for statistics table.// 将类别大小拷贝到统计表for i := range class_to_size {memstats.by_size[i].size = uint32(class_to_size[i])}// Check physPageSize.// 检查physPageSize。if physPageSize == 0 {// The OS init code failed to fetch the physical page size.// 操作系统初始化代码无法获取物理页面大小。throw("failed to get system page size")}// 物理页大小比最大物理页还大if physPageSize > maxPhysPageSize {print("system page size (", physPageSize, ") is larger than maximum page size (", maxPhysPageSize, ")\n")throw("bad system page size")}// 物理页大小比最小物理页还小if physPageSize < minPhysPageSize {print("system page size (", physPageSize, ") is smaller than minimum page size (", minPhysPageSize, ")\n")throw("bad system page size")}// 物理页必须是2的幂次方if physPageSize&(physPageSize-1) != 0 {print("system page size (", physPageSize, ") must be a power of 2\n")throw("bad system page size")}// 操作系统默认页大小，物理页必须是2的幂次方if physHugePageSize&(physHugePageSize-1) != 0 {print("system huge page size (", physHugePageSize, ") must be a power of 2\n")throw("bad system huge page size")}// 操作系统默认页大小大于操作系统最大的页大小if physHugePageSize > maxPhysHugePageSize {// physHugePageSize is greater than the maximum supported huge page size.// Don't throw here, like in the other cases, since a system configured// in this way isn't wrong, we just don't have the code to support them.// Instead, silently set the huge page size to zero.// physHugePageSize大于所支持的最大大页面大小。不要像其他情况那样在这里throw错误，// 因为以这种方式配置的系统没有错，所以我们只是没有支持它们的代码。而是将巨大的页面大小静默设置为零。physHugePageSize = 0}if physHugePageSize != 0 {// Since physHugePageSize is a power of 2, it suffices to increase// physHugePageShift until 1<<physHugePageShift == physHugePageSize.// 由于physHugePageSize为2的幂，因此足以将physHugePageShift增大到1<<physHugePageShift == physHugePageSize。for 1<<physHugePageShift != physHugePageSize {physHugePageShift++}}// Initialize the heap.// 初始化堆mheap_.init()_g_ := getg()_g_.m.mcache = allocmcache()// Create initial arena growth hints.// 创建初始arena增长提示。8表示字节数if sys.PtrSize == 8 {// On a 64-bit machine, we pick the following hints// because://// 1. Starting from the middle of the address space// makes it easier to grow out a contiguous range// without running in to some other mapping.//// 2. This makes Go heap addresses more easily// recognizable when debugging.//// 3. Stack scanning in gccgo is still conservative,// so it's important that addresses be distinguishable// from other data.//// Starting at 0x00c0 means that the valid memory addresses// will begin 0x00c0, 0x00c1, ...// In little-endian, that's c0 00, c1 00, ... None of those are valid// UTF-8 sequences, and they are otherwise as far away from// ff (likely a common byte) as possible. If that fails, we try other 0xXXc0// addresses. An earlier attempt to use 0x11f8 caused out of memory errors// on OS X during thread allocations.  0x00c0 causes conflicts with// AddressSanitizer which reserves all memory up to 0x0100.// These choices reduce the odds of a conservative garbage collector// not collecting memory because some non-pointer block of memory// had a bit pattern that matched a memory address.//// However, on arm64, we ignore all this advice above and slam the// allocation at 0x40 << 32 because when using 4k pages with 3-level// translation buffers, the user address space is limited to 39 bits// On darwin/arm64, the address space is even smaller.//// On AIX, mmaps starts at 0x0A00000000000000 for 64-bit.// processes.//// 在64位计算机上，我们选择以下hit因为：//// 1.从地址空间的中间开始，可以轻松扩展到连续范围，而无需运行其他映射。//// 2.这使Go堆地址在调试时更容易识别。//// 3. gccgo中的堆栈扫描仍然很保守，因此将地址与其他数据区分开很重要。//// 从0x00c0开始意味着有效的内存地址将从0x00c0、0x00c1 ... n 小端开始，即c0 00，c1 00，...// 这些都不是有效的UTF-8序列，否则它们是尽可能远离ff（可能是一个公共字节）。// 如果失败，我们尝试其他0xXXc0地址。较早的尝试使用0x11f8导致线程分配期间OS X上的内存不足错误。// 0x00c0导致与AddressSanitizer发生冲突，后者保留了最多0x0100的所有内存。// 这些选择减少了保守的垃圾收集器不收集内存的可能性，因为某些非指针内存块具有与内存地址匹配的位模式。//// 但是，在arm64上，我们忽略了上面的所有建议，并在0x40 << 32处分配，因为当使用具有3级转换缓冲区的4k页面时，// 用户地址空间在darwin/arm64上被限制为39位，地址空间更小。//// 在AIX上，对于64位，mmaps从0x0A00000000000000开始。// 预先进行内存分配，最多分配64次for i := 0x7f; i >= 0; i-- { // i=0b01111111var p uintptrswitch {case GOARCH == "arm64" && GOOS == "darwin":p = uintptr(i)<<40 | uintptrMask&(0x0013<<28)case GOARCH == "arm64":p = uintptr(i)<<40 | uintptrMask&(0x0040<<32)case GOOS == "aix":if i == 0 {// We don't use addresses directly after 0x0A00000000000000// to avoid collisions with others mmaps done by non-go programs.// 我们不会在0x0A00000000000000之后直接使用地址，以免与非执行程序造成的其他mmap冲突。continue}p = uintptr(i)<<40 | uintptrMask&(0xa0<<52)case raceenabled: // 此值已为false// The TSAN runtime requires the heap// to be in the range [0x00c000000000,// 0x00e000000000).p = uintptr(i)<<32 | uintptrMask&(0x00c0<<32)if p >= uintptrMask&0x00e000000000 {continue}default:p = uintptr(i)<<40 | uintptrMask&(0x00c0<<32)}// p是所求的每个hint起始地址// 采用头插法对hint块进行拉链，小端地址在前// 最终所有的地址都分配在了mheap_.arenaHints上hint := (*arenaHint)(mheap_.arenaHintAlloc.alloc())hint.addr = phint.next, mheap_.arenaHints = mheap_.arenaHints, hint}} else {// On a 32-bit machine, we're much more concerned// about keeping the usable heap contiguous.// Hence://// 1. We reserve space for all heapArenas up front so// they don't get interleaved with the heap. They're// ~258MB, so this isn't too bad. (We could reserve a// smaller amount of space up front if this is a// problem.)//// 2. We hint the heap to start right above the end of// the binary so we have the best chance of keeping it// contiguous.//// 3. We try to stake out a reasonably large initial// heap reservation.//// 在32位计算机上，我们更加关注保持可用堆是连续的。// 因此：//// 1.我们为所有的heapArena保留空间，这样它们就不会与heap交错。 它们约为258MB，因此还算不错。// （如果出现问题，我们可以在前面预留较小的空间。）//// 2.我们建议堆从二进制文件的末尾开始，因此我们有最大的机会保持其连续性。//// 3.我们尝试放出一个相当大的初始堆保留。// 计算arena元数据大小const arenaMetaSize = (1 << arenaBits) * unsafe.Sizeof(heapArena{})// 保留内存meta := uintptr(sysReserve(nil, arenaMetaSize))if meta != 0 { // 保留成功，就进行初始化mheap_.heapArenaAlloc.init(meta, arenaMetaSize)}// We want to start the arena low, but if we're linked// against C code, it's possible global constructors// have called malloc and adjusted the process' brk.// Query the brk so we can avoid trying to map the// region over it (which will cause the kernel to put// the region somewhere else, likely at a high// address).// 我们想从低arena地址开始，但是如果我们与C代码链接，则可能全局构造函数调用了malloc并调整了进程的brk。// 查询brk，以便我们避免尝试在其上映射区域（这将导致内核将区域放置在其他地方，可能位于高地址）。// brk和sbrk相关文档// https://blog.csdn.net/yusiguyuan/article/details/39496057// https://blog.csdn.net/Apollon_krj/article/details/54565768procBrk := sbrk0()// If we ask for the end of the data segment but the// operating system requires a little more space// before we can start allocating, it will give out a// slightly higher pointer. Except QEMU, which is// buggy, as usual: it won't adjust the pointer// upward. So adjust it upward a little bit ourselves:// 1/4 MB to get away from the running binary image.// 如果我们要求结束数据段，但是操作系统在开始分配之前需要更多空间，它将给出稍高的指针。// 像往常一样，除了QEMU之外，它还有很多问题：它不会向上调整指针。 因此，我们自己向上调整一点：// 1/4 MB以远离正在运行的二进制映像。p := firstmoduledata.endif p < procBrk {p = procBrk}if mheap_.heapArenaAlloc.next <= p && p < mheap_.heapArenaAlloc.end {p = mheap_.heapArenaAlloc.end}// alignUp(n, a) alignUp将n舍入为a的倍数。 a必须是2的幂。p = alignUp(p+(256<<10), heapArenaBytes)// Because we're worried about fragmentation on// 32-bit, we try to make a large initial reservation.// 因为我们担心32位上的碎片，所以我们尝试进行较大的初始保留。arenaSizes := []uintptr{512 << 20,256 << 20,128 << 20,}// 从在到小尝试分配，首次分配好就结束for _, arenaSize := range arenaSizes {// sysReserveAligned类似于sysReserve，但是返回的指针字节对齐的。// 它可以保留n个或n+align个字节，因此它返回保留的大小。a, size := sysReserveAligned(unsafe.Pointer(p), arenaSize, heapArenaBytes)if a != nil {mheap_.arena.init(uintptr(a), size)p = uintptr(a) + size // For hint belowbreak}}hint := (*arenaHint)(mheap_.arenaHintAlloc.alloc())hint.addr = phint.next, mheap_.arenaHints = mheap_.arenaHints, hint}
}// sysAlloc allocates heap arena space for at least n bytes. The
// returned pointer is always heapArenaBytes-aligned and backed by
// h.arenas metadata. The returned size is always a multiple of
// heapArenaBytes. sysAlloc returns nil on failure.
// There is no corresponding free function.
//
// sysAlloc returns a memory region in the Prepared state. This region must
// be transitioned to Ready before use.
//
// h must be locked.
/*** sysAlloc至少为n个字节分配堆arena空间。返回的指针始终是heapArenaBytes对齐的，* 并由h.arenas元数据支持。返回的大小始终是heapArenaBytes的倍数。 sysAlloc失败时返回nil。* 没有相应的free函数。** sysAlloc返回处于Prepared状态的内存区域。使用前，该区域必须转换为“就绪”。** h必须被锁定。* @param n 待分配的字节数* @return v 地址指针* @return size 分配的字节数**/
func (h *mheap) sysAlloc(n uintptr) (v unsafe.Pointer, size uintptr) {// 进行字节对齐n = alignUp(n, heapArenaBytes)// First, try the arena pre-reservation.// 首先，尝试arena预定。v = h.arena.alloc(n, heapArenaBytes, &memstats.heap_sys)if v != nil {size = ngoto mapped}// Try to grow the heap at a hint address.// 尝试在hint地址处增加堆。for h.arenaHints != nil {hint := h.arenaHintsp := hint.addrif hint.down {p -= n}if p+n < p {// We can't use this, so don't ask.// 我们不能使用它，所以不回应。v = nil} else if arenaIndex(p+n-1) >= 1<<arenaBits {// Outside addressable heap. Can't use.// 外部可寻址堆。无法使用。v = nil} else {v = sysReserve(unsafe.Pointer(p), n)}if p == uintptr(v) {// Success. Update the hint.// 成功。更新hint。if !hint.down {p += n}hint.addr = psize = nbreak}// Failed. Discard this hint and try the next.//// TODO: This would be cleaner if sysReserve could be// told to only return the requested address. In// particular, this is already how Windows behaves, so// it would simplify things there.// 失败了放弃此hint，然后尝试下一个。//// TODO：如果可以告诉sysReserve仅返回所请求的地址，则这样做会更清洁。// 特别是，这已经是Windows的处理方式，因此它将简化那里的事情。if v != nil {sysFree(v, n, nil)}h.arenaHints = hint.nexth.arenaHintAlloc.free(unsafe.Pointer(hint))}if size == 0 {if raceenabled { // 此值已为false// The race detector assumes the heap lives in// [0x00c000000000, 0x00e000000000), but we// just ran out of hints in this region. Give// a nice failure.throw("too many address space collisions for -race mode")}// All of the hints failed, so we'll take any// (sufficiently aligned) address the kernel will give// us.// 所有hint均失败，因此我们将采用内核将提供给我们的任何地址（已充分对齐）。v, size = sysReserveAligned(nil, n, heapArenaBytes)if v == nil {return nil, 0}// Create new hints for extending this region.// 创建用于扩展此区域的新hint。hint := (*arenaHint)(h.arenaHintAlloc.alloc())hint.addr, hint.down = uintptr(v), truehint.next, mheap_.arenaHints = mheap_.arenaHints, hinthint = (*arenaHint)(h.arenaHintAlloc.alloc())hint.addr = uintptr(v) + sizehint.next, mheap_.arenaHints = mheap_.arenaHints, hint}// Check for bad pointers or pointers we can't use.// 检查错误的指针或我们不能使用的指针。{var bad stringp := uintptr(v)if p+size < p {bad = "region exceeds uintptr range"} else if arenaIndex(p) >= 1<<arenaBits {bad = "base outside usable address space"} else if arenaIndex(p+size-1) >= 1<<arenaBits {bad = "end outside usable address space"}if bad != "" {// This should be impossible on most architectures,// but it would be really confusing to debug.// 在大多数体系结构上，这应该是不可能的，但是调试起来确实很混乱。print("runtime: memory allocated by OS [", hex(p), ", ", hex(p+size), ") not in usable address space: ", bad, "\n")throw("memory reservation exceeds address space limit")}}if uintptr(v)&(heapArenaBytes-1) != 0 {throw("misrounded allocation in sysAlloc")}// Transition from Reserved to Prepared.// 转换状态将从预留到已准备。sysMap(v, size, &memstats.heap_sys)mapped:// Create arena metadata.// 创建arena元数据。for ri := arenaIndex(uintptr(v)); ri <= arenaIndex(uintptr(v)+size-1); ri++ {l2 := h.arenas[ri.l1()]if l2 == nil {// Allocate an L2 arena map.// 分配L2arena映射。l2 = (*[1 << arenaL2Bits]*heapArena)(persistentalloc(unsafe.Sizeof(*l2), sys.PtrSize, nil))if l2 == nil {throw("out of memory allocating heap arena map")}atomic.StorepNoWB(unsafe.Pointer(&h.arenas[ri.l1()]), unsafe.Pointer(l2))}if l2[ri.l2()] != nil {throw("arena already initialized")}var r *heapArenar = (*heapArena)(h.heapArenaAlloc.alloc(unsafe.Sizeof(*r), sys.PtrSize, &memstats.gc_sys))if r == nil {r = (*heapArena)(persistentalloc(unsafe.Sizeof(*r), sys.PtrSize, &memstats.gc_sys))if r == nil {throw("out of memory allocating heap arena metadata")}}// Add the arena to the arenas list.// 将arena添加到arena列表中。if len(h.allArenas) == cap(h.allArenas) {size := 2 * uintptr(cap(h.allArenas)) * sys.PtrSizeif size == 0 {size = physPageSize}newArray := (*notInHeap)(persistentalloc(size, sys.PtrSize, &memstats.gc_sys))if newArray == nil {throw("out of memory allocating allArenas")}oldSlice := h.allArenas*(*notInHeapSlice)(unsafe.Pointer(&h.allArenas)) = notInHeapSlice{newArray, len(h.allArenas), int(size / sys.PtrSize)}copy(h.allArenas, oldSlice)// Do not free the old backing array because// there may be concurrent readers. Since we// double the array each time, this can lead// to at most 2x waste.// 不要释放旧的后备阵列，因为可能有并发读取器。// 由于我们每次将阵列加倍，因此最多可能导致2倍的浪费。}h.allArenas = h.allArenas[:len(h.allArenas)+1]h.allArenas[len(h.allArenas)-1] = ri// Store atomically just in case an object from the// new heap arena becomes visible before the heap lock// is released (which shouldn't happen, but there's// little downside to this).// 以原子方式存储，以防新的堆空间中的对象在释放堆锁之前可见（这不应该发生，但这没有什么坏处）。atomic.StorepNoWB(unsafe.Pointer(&l2[ri.l2()]), unsafe.Pointer(r))}// Tell the race detector about the new heap memory.// 告诉竞态检测器新的堆内存。if raceenabled {racemapshadow(v, size)}return
}// sysReserveAligned is like sysReserve, but the returned pointer is
// aligned to align bytes. It may reserve either n or n+align bytes,
// so it returns the size that was reserved.
/*** sysReserveAligned类似于sysReserve，但是返回的指针以字节对齐。* 它可以保留n个或n+align个字节，因此它返回保留的大小。* @param v 地址指针* @param size 待分配的字节数* @param align 对齐的字节数* @return unsafe.Pointer 新的地址指针* @return uintptr 分配的字节数**/
func sysReserveAligned(v unsafe.Pointer, size, align uintptr) (unsafe.Pointer, uintptr) {// Since the alignment is rather large in uses of this// function, we're not likely to get it by chance, so we ask// for a larger region and remove the parts we don't need.// 由于在使用此功能时对齐方式相当大，因此我们不太可能偶然得到它，// 因此我们要求更大的区域并删除不需要的部分。retries := 0
retry:// 先进行未对齐内存保留p := uintptr(sysReserve(v, size+align))switch {case p == 0: // 未分配成功return nil, 0case p&(align-1) == 0:// We got lucky and got an aligned region, so we can// use the whole thing.// 我们很幸运或取到一个对齐区域，所以我们可以使用整个分配的内存。return unsafe.Pointer(p), size + aligncase GOOS == "windows":// On Windows we can't release pieces of a// reservation, so we release the whole thing and// re-reserve the aligned sub-region. This may race,// so we may have to try again.// 在Windows上，我们无法释放部分保留内存，因此我们释放整个内容并重新保留对齐的子区域。// 这可能会生产竞争，所以我们可能必须重试。sysFree(unsafe.Pointer(p), size+align, nil)p = alignUp(p, align)p2 := sysReserve(unsafe.Pointer(p), size)if p != uintptr(p2) {// Must have raced. Try again.// 必定竞争了，需要重试sysFree(p2, size, nil)if retries++; retries == 100 {throw("failed to allocate aligned heap memory; too many retries")}goto retry}// Success.return p2, sizedefault:// Trim off the unaligned parts.// 修剪掉未对齐的部分。pAligned := alignUp(p, align)// 释放[p, pAligned-1]的内存sysFree(unsafe.Pointer(p), pAligned-p, nil)end := pAligned + sizeendLen := (p + size + align) - endif endLen > 0 {// 释放[end, endLen-1]的内存sysFree(unsafe.Pointer(end), endLen, nil)}// 返回分配地址和分配大小return unsafe.Pointer(pAligned), size}
}// base address for all 0-byte allocations
// 所有0字节分配的基地址
var zerobase uintptr// nextFreeFast returns the next free object if one is quickly available.
// Otherwise it returns 0.
/*** 如果一个可用的对象很快可用，则nextFreeFast返回下一个可用的对象。 否则返回0。* @param s 跨度对象* @return gclinkptr是指向gclink的指针，但对垃圾收集器不透明。**/
func nextFreeFast(s *mspan) gclinkptr {theBit := sys.Ctz64(s.allocCache) // Is there a free object in the allocCache?if theBit < 64 {result := s.freeindex + uintptr(theBit)if result < s.nelems {freeidx := result + 1if freeidx%64 == 0 && freeidx != s.nelems {return 0}s.allocCache >>= uint(theBit + 1)s.freeindex = freeidxs.allocCount++return gclinkptr(result*s.elemsize + s.base())}}return 0
}// nextFree returns the next free object from the cached span if one is available.
// Otherwise it refills the cache with a span with an available object and
// returns that object along with a flag indicating that this was a heavy
// weight allocation. If it is a heavy weight allocation the caller must
// determine whether a new GC cycle needs to be started or if the GC is active
// whether this goroutine needs to assist the GC.
//
// Must run in a non-preemptible context since otherwise the owner of
// c could change.
/*** nextFree从缓存范围中返回下一个空闲对象（如果有）。 否则，它将使用可用对象的跨度重新填充高速缓存，* 并返回该对象以及指示这是繁重分配的标志。 如果分配很重，则调用方必须确定是否需要启动新的GC周期，* 或者GC是否处于活动状态，因此该例行程序是否需要协助GC。** 必须在不可抢占的上下文中运行，因为否则c的所有者可能会更改。* @param spc spanClass表示跨度的大小类别和noscan-ness* @return v gclinkptr指针* @return s 跨度指针* @return shouldhelpgc 是否需要gc帮助**/
func (c *mcache) nextFree(spc spanClass) (v gclinkptr, s *mspan, shouldhelpgc bool) {s = c.alloc[spc]shouldhelpgc = falsefreeIndex := s.nextFreeIndex()if freeIndex == s.nelems {// The span is full.// 跨度已满if uintptr(s.allocCount) != s.nelems {println("runtime: s.allocCount=", s.allocCount, "s.nelems=", s.nelems)throw("s.allocCount != s.nelems && freeIndex == s.nelems")}// 重新填充跨度c.refill(spc)// 标记需要gc帮助shouldhelpgc = trues = c.alloc[spc]freeIndex = s.nextFreeIndex()}if freeIndex >= s.nelems {throw("freeIndex is not valid")}v = gclinkptr(freeIndex*s.elemsize + s.base())s.allocCount++if uintptr(s.allocCount) > s.nelems {println("s.allocCount=", s.allocCount, "s.nelems=", s.nelems)throw("s.allocCount > s.nelems")}return
}// Allocate an object of size bytes.
// Small objects are allocated from the per-P cache's free lists.
// Large objects (> 32 kB) are allocated straight from the heap.
/*** 分配一个大小为size字节的对象。* 从每个P缓存的空闲列表中分配小对象。* 从堆直接分配大对象（> 32 kB）。* @param size 分配的内存* @param needzero* @return 分配的地址指针**/
func mallocgc(size uintptr, typ *_type, needzero bool) unsafe.Pointer {if gcphase == _GCmarktermination {// _GCmarktermination: GC标记终止：分配黑色，P帮助GC，写屏障启用throw("mallocgc called with gcphase == _GCmarktermination")}// 空地址if size == 0 {return unsafe.Pointer(&zerobase)}if debug.sbrk != 0 {align := uintptr(16)if typ != nil {// TODO(austin): This should be just//   align = uintptr(typ.align)// but that's only 4 on 32-bit platforms,// even if there's a uint64 field in typ (see #599).// This causes 64-bit atomic accesses to panic.// Hence, we use stricter alignment that matches// the normal allocator better.// TODO(austin)）：这应该只是align = uintptr（typ.align），// 但即使在typ中有uint64字段，它在32位平台上也只有4（请参阅#599）。// 这会导致对64位原子访问出现panic。 因此，我们使用更严格的对齐方式来更好地匹配常规分配器。if size&7 == 0 {align = 8} else if size&3 == 0 {align = 4} else if size&1 == 0 {align = 2} else {align = 1}}return persistentalloc(size, align, &memstats.other_sys)}// assistG is the G to charge for this allocation, or nil if// GC is not currently active.// assistantG是负责此次分配的G，如果GC当前未处于活动状态，则为nil。var assistG *g// 如果允许增变辅助和后台标记工作程序将对象变黑，则gcBlackenEnabled为1。// 仅当gcphase == _GCmark时才可以设置。if gcBlackenEnabled != 0 {// Charge the current user G for this allocation.// 获取负责此次分配的gassistG = getg()if assistG.m.curg != nil {assistG = assistG.m.curg}// Charge the allocation against the G. We'll account// for internal fragmentation at the end of mallocgc.// 我们将在mallocgc末尾统计内部碎片。assistG.gcAssistBytes -= int64(size)if assistG.gcAssistBytes < 0 {// This G is in debt. Assist the GC to correct// this before allocating. This must happen// before disabling preemption.// 这个G负债中。 协助GC进行更正，然后再分配。 这必须在禁用抢占之前发生。gcAssistAlloc(assistG)}}// Set mp.mallocing to keep from being preempted by GC.// Set mp.mallocing to keep from being preempted by GC.mp := acquirem()if mp.mallocing != 0 { // 当前m已经在分配内容了throw("malloc deadlock")}// 处理信号的g和当前的g是同一个if mp.gsignal == getg() {throw("malloc during signal")}mp.mallocing = 1 // 设当前m正在分配内存shouldhelpgc := falsedataSize := sizec := gomcache()var x unsafe.Pointernoscan := typ == nil || typ.ptrdata == 0if size <= maxSmallSize { // 小分配if noscan && size < maxTinySize { // 微小分配// Tiny allocator.//// Tiny allocator combines several tiny allocation requests// into a single memory block. The resulting memory block// is freed when all subobjects are unreachable. The subobjects// must be noscan (don't have pointers), this ensures that// the amount of potentially wasted memory is bounded.//// Size of the memory block used for combining (maxTinySize) is tunable.// Current setting is 16 bytes, which relates to 2x worst case memory// wastage (when all but one subobjects are unreachable).// 8 bytes would result in no wastage at all, but provides less// opportunities for combining.// 32 bytes provides more opportunities for combining,// but can lead to 4x worst case wastage.// The best case winning is 8x regardless of block size.//// Objects obtained from tiny allocator must not be freed explicitly.// So when an object will be freed explicitly, we ensure that// its size >= maxTinySize.//// SetFinalizer has a special case for objects potentially coming// from tiny allocator, it such case it allows to set finalizers// for an inner byte of a memory block.//// The main targets of tiny allocator are small strings and// standalone escaping variables. On a json benchmark// the allocator reduces number of allocations by ~12% and// reduces heap size by ~20%.//// 微小分配器//// 微小分配器 将几个微小的分配请求组合到一个内存块中。当所有子对象均不可访问时，将释放结果存储块。// 子对象必须是noscan（没有指针），以确保限制可能浪费的内存量。//// 用于合并的存储块的大小（maxTinySize）是可调的。当前设置为16个字节，这涉及最坏情况下2倍的内存浪费// （当除了一个子对象之外的所有其他对象均无法访问时）。8字节完全不会浪费，但是合并的机会更少。 3// 2字节提供了更多的合并机会，但可能导致最坏情况下4倍浪费。无论块大小如何，最佳案例获胜都是8倍。//// 不能显式释放从微小分配器获得的对象。因此，当明确释放对象时，我们确保其大小>=maxTinySize。//// SetFinalizer对于可能来自微小分配器的对象具有特殊情况，在这种情况下，它可以为内存块的内部字节设置终结器（finalizers）。//// 微小分配器的主要目标是小字符串和独立的转义变量。在json基准上，分配器将分配数量减少了约12％，并将堆大小减少了约20％。off := c.tinyoffset// Align tiny pointer for required (conservative) alignment.// 对齐微型指针以进行必需的（保守的）对齐。if size&7 == 0 {off = alignUp(off, 8)} else if size&3 == 0 {off = alignUp(off, 4)} else if size&1 == 0 {off = alignUp(off, 2)}if off+size <= maxTinySize && c.tiny != 0 {// The object fits into existing tiny block.// 该对象适合现有的微小块。x = unsafe.Pointer(c.tiny + off)c.tinyoffset = off + sizec.local_tinyallocs++mp.mallocing = 0releasem(mp)return x}// Allocate a new maxTinySize block.// 分配一个新的maxTinySize块。span := c.alloc[tinySpanClass]v := nextFreeFast(span)if v == 0 { // 快速分配没有成功，进行通常规分配v, _, shouldhelpgc = c.nextFree(tinySpanClass)}x = unsafe.Pointer(v)(*[2]uint64)(x)[0] = 0(*[2]uint64)(x)[1] = 0// See if we need to replace the existing tiny block with the new one// based on amount of remaining free space.// 根据剩余的可用空间量，看看是否需要用新的小块替换现有的小块。if size < c.tinyoffset || c.tiny == 0 {c.tiny = uintptr(x)c.tinyoffset = size}size = maxTinySize} else { // 小分配置var sizeclass uint8// 确定大小类别if size <= smallSizeMax-8 {sizeclass = size_to_class8[(size+smallSizeDiv-1)/smallSizeDiv]} else {sizeclass = size_to_class128[(size-smallSizeMax+largeSizeDiv-1)/largeSizeDiv]}size = uintptr(class_to_size[sizeclass])spc := makeSpanClass(sizeclass, noscan)span := c.alloc[spc]v := nextFreeFast(span)if v == 0 {v, span, shouldhelpgc = c.nextFree(spc)}x = unsafe.Pointer(v)if needzero && span.needzero != 0 {memclrNoHeapPointers(unsafe.Pointer(v), size)}}} else { // 大分配var s *mspanshouldhelpgc = true // 需要gc帮助// systemstack在系统堆栈上运行fn。systemstack(func() {s = largeAlloc(size, needzero, noscan)})s.freeindex = 1s.allocCount = 1x = unsafe.Pointer(s.base())size = s.elemsize}var scanSize uintptrif !noscan {// If allocating a defer+arg block, now that we've picked a malloc size// large enough to hold everything, cut the "asked for" size down to// just the defer header, so that the GC bitmap will record the arg block// as containing nothing at all (as if it were unused space at the end of// a malloc block caused by size rounding).// The defer arg areas are scanned as part of scanstack.// 如果分配一个defer+arg块，现在我们已经选择了一个足够大的malloc大小来容纳所有内容，// 请将“asked for”大小减小为defer头部，以便GC位图将arg块记录为不包含任何内容// （好像是由于大小舍入导致的malloc块末尾的未使用空间）。// 延迟arg区域将作为scanstack的一部分进行扫描。if typ == deferType {dataSize = unsafe.Sizeof(_defer{})}// 记录数据类型heapBitsSetType(uintptr(x), size, dataSize, typ)if dataSize > typ.size {// Array allocation. If there are any// pointers, GC has to scan to the last// element.// 数组分配。 如果有任何指针，GC必须扫描到最后一个元素。if typ.ptrdata != 0 {scanSize = dataSize - typ.size + typ.ptrdata}} else {scanSize = typ.ptrdata}c.local_scan += scanSize}// Ensure that the stores above that initialize x to// type-safe memory and set the heap bits occur before// the caller can make x observable to the garbage// collector. Otherwise, on weakly ordered machines,// the garbage collector could follow a pointer to x,// but see uninitialized memory or stale heap bits.// 确保在调用者使x对垃圾收集器可观察之前，将上面的x初始化为类型安全的内存并设置堆存储的位信息。// 否则，在顺序较弱的计算机上，垃圾收集器可能会跟随指向x的指针，但会看到未初始化的内存或陈旧的堆位。publicationBarrier()// Allocate black during GC.// All slots hold nil so no scanning is needed.// This may be racing with GC so do it atomically if there can be// a race marking the bit.// 在GC期间分配黑色。 所有槽位均为nill，因此无需扫描。这可能与GC发生争用，// 因此如果可以进行争夺来标记该位，则可以自动进行。if gcphase != _GCoff { // go未运行// gcmarknewobject将新分配的对象标记为黑色。obj不得包含任何非nil指针。gcmarknewobject(uintptr(x), size, scanSize)}if raceenabled {racemalloc(x, size)}if msanenabled {msanmalloc(x, size)}mp.mallocing = 0 // 标记分配已经完成releasem(mp) // 释放mif debug.allocfreetrace != 0 {tracealloc(x, size, typ)}// 设置内存采样if rate := MemProfileRate; rate > 0 {if rate != 1 && size < c.next_sample {c.next_sample -= size} else {mp := acquirem()profilealloc(mp, x, size)releasem(mp)}}if assistG != nil {// Account for internal fragmentation in the assist// debt now that we know it.// 在assist debt中解决内部碎片化问题assistG.gcAssistBytes -= int64(size - dataSize)}if shouldhelpgc {// gcTrigger是用于启动GC循环的动作（predicate）。 具体来说，它是_GCoff阶段的退出条件。if t := (gcTrigger{kind: gcTriggerHeap}); t.test() {gcStart(t)}}return x
}/*** 大内存分配* @param* @return**/
func largeAlloc(size uintptr, needzero bool, noscan bool) *mspan {// print("largeAlloc size=", size, "\n")// 判断是否内存溢出if size+_PageSize < size {throw("out of memory")}// 计算需要分配的页数npages := size >> _PageShiftif size&_PageMask != 0 {npages++}// Deduct credit for this span allocation and sweep if// necessary. mHeap_Alloc will also sweep npages, so this only// pays the debt down to npage pages.// 扣除此跨度分配的信用，并在必要时进行扫描。 mHeap_Alloc还将清扫npages，因此这只会将债务减少到npage页。deductSweepCredit(npages*_PageSize, npages)// mheap_.alloc从GC的堆中分配新的npage页。s := mheap_.alloc(npages, makeSpanClass(0, noscan), needzero)if s == nil {throw("out of memory")}s.limit = s.base() + size// heapBitsForAddr返回地址addr的heapBits。heapBitsForAddr(s.base()).initSpan(s)return s
}// implementation of new builtin
// compiler (both frontend and SSA backend) knows the signature
// of this function
/*** 内建函数new的实现，编译器（前端和SSA后端）都知道此函数的签名* @param 类型* @return 对象指针**/
func newobject(typ *_type) unsafe.Pointer {return mallocgc(typ.size, typ, true)
}/*** reflect包中的unsafe_New实现* @param* @return**/
//go:linkname reflect_unsafe_New reflect.unsafe_New
func reflect_unsafe_New(typ *_type) unsafe.Pointer {return mallocgc(typ.size, typ, true)
}/*** internal/reflectlite包中的unsafe_New实现* @param* @return**/
//go:linkname reflectlite_unsafe_New internal/reflectlite.unsafe_New
func reflectlite_unsafe_New(typ *_type) unsafe.Pointer {return mallocgc(typ.size, typ, true)
}/*** newarray分配一个类型为typ的n个元素组成的数组。* @param* @return**/
// newarray allocates an array of n elements of type typ.
func newarray(typ *_type, n int) unsafe.Pointer {if n == 1 {return mallocgc(typ.size, typ, true)}// 计算数组创建需要分配的内存mem, overflow := math.MulUintptr(typ.size, uintptr(n))if overflow || mem > maxAlloc || n < 0 {panic(plainError("runtime: allocation size out of range"))}return mallocgc(mem, typ, true)
}/*** reflect包中的unsafe_NewArray实现* @param* @return**/
//go:linkname reflect_unsafe_NewArray reflect.unsafe_NewArray
func reflect_unsafe_NewArray(typ *_type, n int) unsafe.Pointer {return newarray(typ, n)
}func profilealloc(mp *m, x unsafe.Pointer, size uintptr) {mp.mcache.next_sample = nextSample()mProf_Malloc(x, size)
}// nextSample returns the next sampling point for heap profiling. The goal is
// to sample allocations on average every MemProfileRate bytes, but with a
// completely random distribution over the allocation timeline; this
// corresponds to a Poisson process with parameter MemProfileRate. In Poisson
// processes, the distance between two samples follows the exponential
// distribution (exp(MemProfileRate)), so the best return value is a random
// number taken from an exponential distribution whose mean is MemProfileRate.
/*** nextSample返回用于堆分析的下一个采样点。 目标是在每个MemProfileRate字节抽样平均分配，* 但在分配时间轴上具有完全随机的分配； 这对应于带有参数MemProfileRate的泊松过程。* 在泊松过程中，两个样本之间的距离遵循指数分布（exp(MemProfileRate)），* 因此最佳返回值是取自均值为MemProfileRate的指数分布的随机数。* @param* @return**/
func nextSample() uintptr {if GOOS == "plan9" {// Plan 9 doesn't support floating point in note handler.if g := getg(); g == g.m.gsignal {return nextSampleNoFP()}}return uintptr(fastexprand(MemProfileRate))
}// fastexprand returns a random number from an exponential distribution with
// the specified mean.
/*** fastexprand从具有指定平均值的指数分布中返回一个随机数。* @param* @return**/
func fastexprand(mean int) int32 {// Avoid overflow. Maximum possible step is// -ln(1/(1<<randomBitCount)) * mean, approximately 20 * mean.// 避免溢出。 最大可能步长为-ln(1/(1<<randomBitCount))*平均值，大约20*平均值。switch {case mean > 0x7000000:mean = 0x7000000case mean == 0:return 0}// Take a random sample of the exponential distribution exp(-mean*x).// The probability distribution function is mean*exp(-mean*x), so the CDF is// p = 1 - exp(-mean*x), so// q = 1 - p == exp(-mean*x)// log_e(q) = -mean*x// -log_e(q)/mean = x// x = -log_e(q) * mean// x = log_2(q) * (-log_e(2)) * mean    ; Using log_2 for efficiencyconst randomBitCount = 26q := fastrand()%(1<<randomBitCount) + 1qlog := fastlog2(float64(q)) - randomBitCountif qlog > 0 {qlog = 0}const minusLog2 = -0.6931471805599453 // -ln(2)return int32(qlog*(minusLog2*float64(mean))) + 1
}// nextSampleNoFP is similar to nextSample, but uses older,
// simpler code to avoid floating point.
/*** nextSampleNoFP与nextSample相似，但使用更旧，更简单的代码来避免浮点数。* @param* @return**/
func nextSampleNoFP() uintptr {// Set first allocation sample size.rate := MemProfileRateif rate > 0x3fffffff { // make 2*rate not overflowrate = 0x3fffffff}if rate != 0 {return uintptr(fastrand() % uint32(2*rate))}return 0
}/*** 持久化分配结构**/
type persistentAlloc struct {base *notInHeapoff  uintptr
}/*** 全局分配结构**/
var globalAlloc struct {mutexpersistentAlloc
}// persistentChunkSize is the number of bytes we allocate when we grow
// a persistentAlloc.
// persistentChunkSize是我们增长persistentAlloc时分配的字节数。256KB
const persistentChunkSize = 256 << 10// persistentChunks is a list of all the persistent chunks we have
// allocated. The list is maintained through the first word in the
// persistent chunk. This is updated atomically.
// psistenceChunks是我们分配的所有持久性块的列表。
// 该列表通过持久性块中的第一个字（word）进行维护。 这是原子更新的。
var persistentChunks *notInHeap// Wrapper around sysAlloc that can allocate small chunks.
// There is no associated free operation.
// Intended for things like function/type/debug-related persistent data.
// If align is 0, uses default align (currently 8).
// The returned memory will be zeroed.
//
// Consider marking persistentalloc'd types go:notinheap.
/*** 围绕sysAlloc的包装程序，可以分配小内存块。 没有相关的自由操作。* 用于函数/类型/调试相关的持久数据。 如果align为0，则使用默认的align（当前为8）。* 返回的内存将被清零。考虑标记为持久分配的类型go:notinheap。* @param * @return **/
func persistentalloc(size, align uintptr, sysStat *uint64) unsafe.Pointer {var p *notInHeapsystemstack(func() {p = persistentalloc1(size, align, sysStat)})return unsafe.Pointer(p)
}/*** 必须在系统堆栈上运行，因为堆栈增长可以（重新）调用它。 请参阅问题9174。* @param* @return**/
// Must run on system stack because stack growth can (re)invoke it.
// See issue 9174.
//go:systemstack
func persistentalloc1(size, align uintptr, sysStat *uint64) *notInHeap {const (// Windows上的VM预留粒度为64KmaxBlock = 64 << 10 // VM reservation granularity is 64K on windows)if size == 0 {throw("persistentalloc: size == 0")}if align != 0 {if align&(align-1) != 0 {throw("persistentalloc: align is not a power of 2")}if align > _PageSize {throw("persistentalloc: align is too large")}} else {align = 8}if size >= maxBlock {return (*notInHeap)(sysAlloc(size, sysStat))}mp := acquirem()var persistent *persistentAllocif mp != nil && mp.p != 0 {persistent = &mp.p.ptr().palloc} else {lock(&globalAlloc.mutex)persistent = &globalAlloc.persistentAlloc}persistent.off = alignUp(persistent.off, align)if persistent.off+size > persistentChunkSize || persistent.base == nil {persistent.base = (*notInHeap)(sysAlloc(persistentChunkSize, &memstats.other_sys))if persistent.base == nil {if persistent == &globalAlloc.persistentAlloc {unlock(&globalAlloc.mutex)}throw("runtime: cannot allocate memory")}// Add the new chunk to the persistentChunks list.for {chunks := uintptr(unsafe.Pointer(persistentChunks))*(*uintptr)(unsafe.Pointer(persistent.base)) = chunksif atomic.Casuintptr((*uintptr)(unsafe.Pointer(&persistentChunks)), chunks, uintptr(unsafe.Pointer(persistent.base))) {break}}persistent.off = alignUp(sys.PtrSize, align)}p := persistent.base.add(persistent.off)persistent.off += sizereleasem(mp)if persistent == &globalAlloc.persistentAlloc {unlock(&globalAlloc.mutex)}if sysStat != &memstats.other_sys {mSysStatInc(sysStat, size)mSysStatDec(&memstats.other_sys, size)}return p
}/*** inPersistentAlloc报告p是否指向由persistentalloc分配的内存。* 由于它是由cgo检查器代码调用的，而后者由写屏障代码调用，因此必须不能拆分。* @param* @return**/
// inPersistentAlloc reports whether p points to memory allocated by
// persistentalloc. This must be nosplit because it is called by the
// cgo checker code, which is called by the write barrier code.
//go:nosplit
func inPersistentAlloc(p uintptr) bool {chunk := atomic.Loaduintptr((*uintptr)(unsafe.Pointer(&persistentChunks)))for chunk != 0 {if p >= chunk && p < chunk+persistentChunkSize {return true}chunk = *(*uintptr)(unsafe.Pointer(chunk))}return false
}// linearAlloc is a simple linear allocator that pre-reserves a region
// of memory and then maps that region into the Ready state as needed. The
// caller is responsible for locking.
/*** linearAlloc是一个简单的线性分配器，它可以预先保留一个内存区域，* 然后根据需要将该区域映射到“就绪”状态。 调用方负责锁定。**/
type linearAlloc struct {next   uintptr // next free byte // next free bytemapped uintptr // one byte past end of mapped space // 映射空间的末尾一个字节end    uintptr // end of reserved space // 保留空间的结尾
}/*** 初始化线性分配器* @param* @return**/
func (l *linearAlloc) init(base, size uintptr) {l.next, l.mapped = base, basel.end = base + size
}/*** 分配内存* @param* @return**/
func (l *linearAlloc) alloc(size, align uintptr, sysStat *uint64) unsafe.Pointer {p := alignUp(l.next, align)if p+size > l.end {return nil}l.next = p + sizeif pEnd := alignUp(l.next-1, physPageSize); pEnd > l.mapped {// Transition from Reserved to Prepared to Ready.sysMap(unsafe.Pointer(l.mapped), pEnd-l.mapped, sysStat)sysUsed(unsafe.Pointer(l.mapped), pEnd-l.mapped)l.mapped = pEnd}return unsafe.Pointer(p)
}/*** notInHeap是由sysAlloc或persistentAlloc之类的较低级分配器分配的堆外内存。** 通常，最好使用标记为go:notinheap的实类型，但这在不可能的情况下（例如在分配器中）用作通用类型。** TODO：使用它作为sysAlloc，persistentAlloc等的返回类型吗？**/
// notInHeap is off-heap memory allocated by a lower-level allocator
// like sysAlloc or persistentAlloc.
//
// In general, it's better to use real types marked as go:notinheap,
// but this serves as a generic type for situations where that isn't
// possible (like in the allocators).
//
// TODO: Use this as the return type of sysAlloc, persistentAlloc, etc?
//
//go:notinheap
type notInHeap struct{}func (p *notInHeap) add(bytes uintptr) *notInHeap {return (*notInHeap)(unsafe.Pointer(uintptr(unsafe.Pointer(p)) + bytes))
}

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go源码阅读——malloc.go

【博文目录>>>】【项目地址>>>】

内存分配器

分配器的数据结构为：

分配一个小对象沿用了高速缓存的层次结构：

清除mspan并释放对象沿用了类似的层次结构：

虚拟内存布局

OS内存管理抽象层

sysAlloc

sysFree

sysReserve

sysMap

sysUsed

sysUnused

sysFault

状态转换图

hint的选择

64位机器

32位机器

一些常量

源码

go源码阅读——malloc.go相关推荐

最新文章

热门文章

go源码阅读——malloc.go

【博文目录>>>】 【项目地址>>>】

内存分配器

分配器的数据结构为：

分配一个小对象沿用了高速缓存的层次结构：

清除mspan并释放对象沿用了类似的层次结构：

虚拟内存布局

OS内存管理抽象层

sysAlloc

sysFree

sysReserve

sysMap

sysUsed

sysUnused

sysFault

状态转换图

hint的选择

64位机器

32位机器

一些常量

源码

go源码阅读——malloc.go相关推荐

最新文章

热门文章

【博文目录>>>】【项目地址>>>】