Roman Divacky

rdivacky@FreeBSD.org

&tm-attrib.adobe; &tm-attrib.ibm; &tm-attrib.freebsd; &tm-attrib.linux; &tm-attrib.netbsd; &tm-attrib.realnetworks; &tm-attrib.oracle; &tm-attrib.sun; &tm-attrib.general; This masters thesis deals with updating the &linux; emulation layer (the so called Linuxulator). The task was to update the layer to match the functionality of &linux; 2.6. As a reference implementation, the &linux; 2.6.16 kernel was chosen. The concept is loosely based on the NetBSD implementation. Most of the work was done in the summer of 2006 as a part of the Google Summer of Code students program. The focus was on bringing the NPTL (new &posix; thread library) support into the emulation layer, including TLS (thread local storage), futexes (fast user space mutexes), PID mangling, and some other minor things. Many small problems were identified and fixed in the process. My work was integrated into the main &os; source repository and will be shipped in the upcoming 7.0R release. We, the emulation development team, are working on making the &linux; 2.6 emulation the default emulation layer in &os;. In the last few years the open source &unix; based operating systems started to be widely deployed on server and client machines. Among these operating systems I would like to point out two: &os;, for its BSD heritage, time proven code base and many interesting features and &linux; for its wide user base, enthusiastic open developer community and support from large companies. &os; tends to be used on server class machines serving heavy duty networking tasks with less usage on desktop class machines for ordinary users. While &linux; has the same usage on servers, but it is used much more by home based users. This leads to a situation where there are many binary only programs available for &linux; that lack support for &os;. Naturally, a need for the ability to run &linux; binaries on a &os; system arises and this is what this thesis deals with: the emulation of the &linux; kernel in the &os; operating system. During the Summer of 2006 Google Inc. sponsored a project which focused on extending the &linux; emulation layer (the so called Linuxulator) in &os; to include &linux; 2.6 facilities. This thesis is written as a part of this project. In this section we are going to describe every operating system in question. How they deal with syscalls, trapframes etc. all the low-level stuff. We also describe the way they understand common &unix; primitives like what a PID is, what a thread is, etc. In the third subsection we talk about how &unix; on &unix; emulation could be done in general. &unix; is an operating system with a long history that has influenced almost every other operating system currently in use. Starting in the 1960s, its development continues to this day (although in different projects). &unix; development soon forked into two main ways: the BSDs and System III/V families. They mutually influenced themselves by growing a common &unix; standard. Among the contributions originated in BSD we can name virtual memory, TCP/IP networking, FFS, and many others. The System V branch contributed to SysV interprocess communication primitives, copy-on-write, etc. &unix; itself does not exist any more but its ideas have been used by many other operating systems world wide thus forming the so called &unix;-like operating systems. These days the most influential ones are &linux;, Solaris, and possibly (to some extent) &os;. There are in-company &unix; derivatives (AIX, HP-UX etc.), but these have been more and more migrated to the aforementioned systems. Let us summarize typical &unix; characteristics. Every running program constitutes a process that represents a state of the computation. Running process is divided between kernel-space and user-space. Some operations can be done only from kernel space (dealing with hardware etc.), but the process should spend most of its lifetime in the user space. The kernel is where the management of the processes, hardware, and low-level details take place. The kernel provides a standard unified &unix; API to the user space. The most important ones are covered below. Common &unix; API defines a syscall as a way to issue commands from a user space process to the kernel. The most common implementation is either by using an interrupt or specialized instruction (think of SYSENTER/SYSCALL instructions for ia32). Syscalls are defined by a number. For example in &os;, the syscall number 85 is the &man.swapon.2; syscall and the syscall number 132 is &man.mkfifo.2;. Some syscalls need parameters, which are passed from the user-space to the kernel-space in various ways (implementation dependant). Syscalls are synchronous. Another possible way to communicate is by using a trap. Traps occur asynchronously after some event occurs (division by zero, page fault etc.). A trap can be transparent for a process (page fault) or can result in a reaction like sending a signal (division by zero). There are other APIs (System V IPC, shared memory etc.) but the single most important API is signal. Signals are sent by processes or by the kernel and received by processes. Some signals can be ignored or handled by a user supplied routine, some result in a predefined action that cannot be altered or ignored. Kernel instances are processed first in the system (so called init). Every running process can create its identical copy using the &man.fork.2; syscall. Some slightly modified versions of this syscall were introduced but the basic semantic is the same. Every running process can morph into some other process using the &man.exec.3; syscall. Some modifications of this syscall were introduced but all serve the same basic purpose. Processes end their lives by calling the &man.exit.2; syscall. Every process is identified by a unique number called PID. Every process has a defined parent (identified by its PID). Traditional &unix; does not define any API nor implementation for threading, while &posix; defines its threading API but the implementation is undefined. Traditionally there were two ways of implementing threads. Handling them as separate processes (1:1 threading) or envelope the whole thread group in one process and managing the threading in userspace (1:N threading). Comparing main features of each approach: 1:1 threading - heavyweight threads - the scheduling cannot be altered by the user (slightly mitigated by the &posix; API) + no syscall wrapping necessary + can utilize multiple CPUs 1:N threading + lightweight threads + scheduling can be easily altered by the user - syscalls must be wrapped - cannot utilize more than one CPU The &os; project is one of the oldest open source operating systems currently available for daily use. It is a direct descendant of the genuine &unix; so it could be claimed that it is a true &unix; although licensing issues do not permit that. The start of the project dates back to the early 1990's when a crew of fellow BSD users patched the 386BSD operating system. Based on this patchkit a new operating system arose named &os; for its liberal license. Another group created the NetBSD operating system with different goals in mind. We will focus on &os;. &os; is a modern &unix;-based operating system with all the features of &unix;. Preemptive multitasking, multiuser facilities, TCP/IP networking, memory protection, symmetric multiprocessing support, virtual memory with merged VM and buffer cache, they are all there. One of the interesting and extremely useful features is the ability to emulate other &unix;-like operating systems. As of December 2006 and 7-CURRENT development, the following emulation functionalities are supported: &os;/i386 emulation on &os;/amd64 &os;/i386 emulation on &os;/ia64 &linux;-emulation of &linux; operating system on &os; NDIS-emulation of Windows networking drivers interface NetBSD-emulation of NetBSD operating system PECoff-support for PECoff &os; executables SVR4-emulation of System V revision 4 &unix; Actively developed emulations are the &linux; layer and various &os;-on-&os; layers. Others are not supposed to work properly nor be usable these days. &os; development happens in a central CVS repository where only a selected team of so called commiters can write. This repository posseses several branches; the most interesting are the HEAD branch, in &os; nomenclature called -CURRENT, and RELENG_X branches, where X stands for a number indicating a major version of &os;. As of December 2006, there are development branches for 6.X development (RELENG_6) and for the 5.X development (RELENG_5). Other branches are closed and not actively maintained or only fed with security patches by the Security Officer of the &os; project. Historically the active development was done in the HEAD branch so it was considered extremely unstable and supposed to happen to break at any time. This is not true any more as the Perforce (commercial version control system) repository was introduced so that active development happen there. There are many branches in Perforce where development of certain parts of the system happens and these branches are from time to time merged back to the main CVS repository thus effectively putting the given feature to the &os; operating system. Tha same happened with the rdivacky_linuxolator branch where development of this thesis code was going on. More info about the &os; operating system can be found at [2]. &os; is traditional flavor of &unix; in the sense of dividing the run of processes into two halves: kernel space and user space run. There are two types of process entry to the kernel: a syscall and a trap. There is only one way to return. In the subsequent sections we will describe the three gates to/from the kernel. The whole description applies to the i386 architecture as the Linuxulator only exists there but the concept is similar on other architectures. The information was taken from [1] and the source code. &os; has an abstraction called an execution class loader, which is a wedge into the &man.execve.2; syscall. This employs a structure sysentvec, which describes an executable ABI. It contains things like errno translation table, signal translation table, various functions to serve syscall needs (stack fixup, coredumping, etc.). Every ABI the &os; kernel wants to support must define this structure, as it is used later in the syscall processing code and at some other places. System entries are handled by trap handlers, where we can access both the kernel-space and the user-space at once. Syscalls on &os; are issued by executing interrupt 0x80 with register %eax set to a desired syscall number with arguments passed on the stack. When a process issues an interrupt 0x80, the int0x80 syscall trap handler is issued (defined in sys/i386/i386/exception.s), which prepares arguments (i.e. copies them on to the stack) for a call to a C function &man.syscall.2; (defined in sys/i386/i386/trap.c), which processes the passed in trapframe. The processing consists of preparing the syscall (depending on the sysvec entry), determining if the syscall is 32-bit or 64-bit one (changes size of the parameters), then the parameters are copied, including the syscall. Next, the actual syscall function is executed with processing of the return code (special cases for ERESTART and EJUSTRETURN errors). Finally an userret() is scheduled, switching the process back to the users-pace. The parameters to the actual syscall handler are passed in the form of struct thread *td, struct syscall args * arguments where the second parameter is a pointer to the copied in structure of parameters. Handling of traps in &os; is similar to the handling of syscalls. Whenever a trap occurs, an assembler handler is called. It is chosen between alltraps, alltraps with regs pushed or calltrap depending on the type of the trap. This handler prepares arguments for a call to a C function trap() (defined in sys/i386/i386/trap.c), which then processes the occurred trap. After the processing it might send a signal to the process and/or exit to userland using userret(). Exits from kernel to userspace happen using the assembler routine doreti regardless of whether the kernel was entered via a trap or via a syscall. This restores the program status from the stack and returns to the userspace. &os; operating system adheres to the traditional &unix; scheme, where every process has a unique identification number, the so called PID (Process ID). PID numbers are allocated either linearly or randomly ranging from 0 to PID_MAX. The allocation of PID numbers is done using linear searching of PID space. Every thread in a process receives the same PID number as result of the &man.getpid.2; call. There are currently two ways to implement threading in &os;. The first way is M:N threading followed by the 1:1 threading model. The default library used is M:N threading (libpthread) and you can switch at runtime to 1:1 threading (libthr). The plan is to switch to 1:1 library by default soon. Although those two libraries use the same kernel primitives, they are accessed through different API(es). The M:N library uses the kse_* family of syscalls while the 1:1 library uses the thr_* family of syscalls. Because of this, there is no general concept of thread ID shared between kernel and userspace. Of course, both threading libraries implement the pthread thread ID API. Every kernel thread (as described by struct thread) has td tid identifier but this is not directly accessible from userland and solely serves the kernel's needs. It is also used for 1:1 threading library as pthread's thread ID but handling of this is internal to the library and cannot be relied on. As stated previously there are two implementations of threading in &os;. The M:N library divides the work between kernel space and userspace. Thread is an entity that gets scheduled in the kernel but it can represent various number of userspace threads. M userspace threads get mapped to N kernel threads thus saving resources while keeping the ability to exploit multiprocessor parallelism. Further information about the implementation can be obtained from the man page or [1]. The 1:1 library directly maps a userland thread to a kernel thread thus greatly simplifying the scheme. None of these designs implement a fairness mechanism (such a mechanism was implemented but it was removed recently because it caused serious slowdown and made the code more difficult to deal with). &linux; is a &unix;-like kernel originally developed by Linus Torvalds, and now being contributed to by a massive crowd of programmers all around the world. From its mere beginnings to todays, with wide support from companies such as IBM or Google, &linux; is being associated with its fast development pace, full hardware support and benevolent dictator model of organization. &linux; development started in 1991 as a hobbyist project at University of Helsinki in Finland. Since then it has obtained all the features of a modern &unix;-like OS: multiprocessing, multiuser support, virtual memory, networking, basically everything is there. There are also highly advanced features like virtualization etc. As of 2006 &linux; seems to be the most widely used open source operating system with support from independent software vendors like Oracle, RealNetworks, Adobe, etc. Most of the commercial software distributed for &linux; can only be obtained in a binary form so recompilation for other operating systems is impossible. Most of the &linux; development happens in a Git version control system. Git is a distributed system so there is no central source of the &linux; code, but some branches are considered prominent and official. The version number scheme implemented by &linux; consists of four numbers A.B.C.D. Currently development happens in 2.6.C.D, where C represents major version, where new features are added or changed while D is a minor version for bugfixes only. More information can be obtained from [4]. &linux; follows the traditional &unix; scheme of dividing the run of a process in two halves: the kernel and user space. The kernel can be entered in two ways: via a trap or via a syscall. The return is handled only in one way. The further description applies to &linux; 2.6 on the &i386; architecture. This information was taken from [3]. Syscalls in &linux; are performed (in userspace) using syscallX macros where X substitutes a number representing the number of parameters of the given syscall. This macro translates to a code that loads %eax register with a number of the syscall and executes interrupt 0x80. After this syscall return is called, which translates negative return values to positive errno values and sets res to -1 in case of an error. Whenever the interrupt 0x80 is called the process enters the kernel in system call trap handler. This routine saves all registers on the stack and calls the selected syscall entry. Note that the &linux; calling convention expects parameters to the syscall to be passed via registers as shown here: parameter -> %ebx parameter -> %ecx parameter -> %edx parameter -> %esi parameter -> %edi parameter -> %ebp There are some exceptions to this, where &linux; uses different calling convention (most notably the clone syscall). Zhe trap handlers are introduced in arch/i386/kernel/traps.c and most of these handlers live in arch/i386/kernel/entry.S, where handling of the traps happens. Return from the syscall is managed by syscall &man.exit.3;, which checks for the process having unfinished work, then checks whether we used user-supplied selectors. If this happens stack fixing is applied and finally the registers are restored from the stack and the process returns to the userspace. In the 2.6 version, the &linux; operating system redefined some of the traditional &unix; primitives, notably PID, TID and thread. PID is defined not to be unique for every process, so for some processes (threads) &man.getppid.2; returns the same value. Unique identification of process is provided by TID. This is because NPTL (New &posix; Thread Library) defines threads to be normal processes (so called 1:1 threading). Spawning a new process in &linux; 2.6 happens using the clone syscall (fork variants are reimplemented using it). This clone syscall defines a set of flags that affect behaviour of the cloning process regarding thread implementation. The semantic is a bit fuzzy as there is no single flag telling the syscall to create a thread. Implemented clone flags are: CLONE_VM - processes share their memory space CLONE_FS - share umask, cwd and namespace CLONE_FILES - share open files CLONE_SIGHAND - share signal handlers and blocked signals CLONE_PARENT - share parent CLONE_THREAD - be thread (further explanation below) CLONE_NEWNS - new namespace CLONE_SYSVSEM - share SysV undo structures CLONE_SETTLS - setup TLS at supplied address CLONE_PARENT_SETTID - set TID in the parent CLONE_CHILD_CLEARTID - clear TID in the child CLONE_CHILD_SETTID - set TID in the child CLONE_PARENT sets the real parent to the parent of the caller. This is useful for threads because if thread A creates thread B we want thread B to be parented to the parent of the whole thread group. CLONE_THREAD does exactly the same thing as CLONE_PARENT, CLONE_VM and CLONE_SIGHAND, rewrites PID to be the same as PID of the caller, sets exit signal to be none and enters the thread group. CLONE_SETTLS sets up GDT entries for TLS handling. The CLONE_*_*TID set of flags sets/clears user supplied address to TID or 0. As you can see the CLONE_THREAD does most of the work and does not seem to fit the scheme very well. The original intention is unclear (even for authors, according to comments in the code) but I think originally there was one threading flag, which was then parcelled among many other flags but this separation was never fully finished. It is also unclear what this partition is good for as glibc does not use that so only hand-written use of the clone permits a programmer to access this features. For non-threaded programs the PID and TID are the same. For threaded programs the first thread PID and TID are the same and every created thread shares the same PID and gets assigned a unique TID (because CLONE_THREAD is passed in) also parent is shared for all processes forming this threaded program. The code that implements &man.pthread.create.3; in NPTL defines the clone flags like this: int clone_flags = (CLONE_VM | CLONE_FS | CLONE_FILES | CLONE_SIGNAL | CLONE_SETTLS | CLONE_PARENT_SETTID | CLONE_CHILD_CLEARTID | CLONE_SYSVSEM #if __ASSUME_NO_CLONE_DETACHED == 0 | CLONE_DETACHED #endif | 0); The CLONE_SIGNAL is defined like #define CLONE_SIGNAL (CLONE_SIGHAND | CLONE_THREAD) the last 0 means no signal is sent when any of the threads exits. According to a dictionary definition, emulation is the ability of a program or device to imitate another program or device. This is achieved by providing the same reaction to a given stimulus as the emulated object. In practice, the software world mostly sees three types of emulation - a program used to emulate a machine (QEMU, various game console emulators etc.), software emulation of a hardware facility (OpenGL emulators, floating point units emulation etc.) and operating system emulation (either in kernel of the operating system or as a userspace program). Emulation is usually used in a place, where using the original component is not feasible nor possible at all. For example someone might want to use a program developed for a different operating system than he uses. Then emulation comes in handy. Sometimes there is no other way but to use emulation - e.g. when the hardware device you try to use does not exist (yet/anymore) then there is no other way but emulation. This happens often when porting an operating system to a new (non-existent) platform. Sometimes it is just cheaper to emulate. Looking from an implementation point of view, there are two main approaches to the implementation of emulation. You can either emulate the whole thing - accepting possible inputs of the original object, maintaining inner state and emitting correct output based on the state and/or input. This kind of emulation does not require any special conditions and basically can be implemented anywhere for any device/program. The drawback is that implementing such emulation is quite difficult, time-consuming and error-prone. In some cases we can use a simpler approach. Imagine you want to emulate a printer that prints from left to right on a printer that prints from right to left. It is obvious that there is no need for a complex emulation layer but simply reversing of the printed text is sufficient. Sometimes the emulating environment is very similar to the emulated one so just a thin layer of some translation is necessary to provide fully working emulation! As you can see this is much less demanding to implement, so less time-consuming and error-prone than the previous approach. But the necessary condition is that the two environments must be similar enough. The third approach combines the two previous. Most of the time the objects do not provide the same capabilities so in a case of emulating the more powerful one on the less powerful we have to emulate the missing features with full emulation described above. This master thesis deals with emulation of &unix; on &unix;, which is exactly the case, where only a thin layer of translation is sufficient to provide full emulation. The &unix; API consists of a set of syscalls, which are usually self contained and do not affect some global kernel state. There are a few syscalls that affect inner state but this can be dealt with by providing some structures that maintain the extra state. No emulation is perfect and emulations tend to lack some parts but this usually does not cause any serious drawbacks. Imagine a game console emulator that emulates everything but music output. No doubt that the games are playable and one can use the emulator. It might not be that comfortable as the original game console but its an acceptable compromise between price and comfort. The same goes with the &unix; API. Most programs can live with a very limited set of syscalls working. Those syscalls tend to be the oldest ones (&man.read.2;/&man.write.2;, &man.fork.2; family, &man.signal.3; handling, &man.exit.3;, &man.socket.2; API) hence it is easy to emulate because their semantics is shared among all &unix;es, which exist todays. As stated earlier, &os; supports running binaries from several other &unix;es. This works because &os; has an abstraction called the execution class loader. This wedges into the &man.execve.2; syscall, so when &man.execve.2; is about to execute a binary it examines its type. There are basically two types of binaries in &os;. Shell-like text scripts which are identified by #! as their first two characters and normal (typically ELF) binaries, which are a representation of a compiled executable object. The vast majority (one could say all of them) of binaries in &os; are from type ELF. ELF files contain a header, which specifies the OS ABI for this ELF file. By reading this information, the operating system can accurately determine what type of binary the given file is. Every OS ABI must be registered in the &os; kernel. This applies to the &os; native OS ABI, as well. So when &man.execve.2; executes a binary it iterates through the list of registered APIs and when it finds the right one it starts to use the information contained in the OS ABI description (its syscall table, errno translation table, etc.). So every time the process calls a syscall, it uses its own set of syscalls instead of some global one. This effectively provides a very elegant and easy way of supporting execution of various binary formats. The nature of emulation of different OSes (and also some other subsystems) led developers to invite a handler event mechanism. There are various places in the kernel, where a list of event handlers are called. Every subsystem can register an event handler and they are called accordingly. For example, when a process exits there is a handler called that possibly cleans up whatever the subsystem needs to be cleaned. Those simple facilities provide basically everything that is needed for the emulation infrastructure and in fact these are basically the only things necessary to implement the &linux; emulation layer. Emulation layers need some support from the operating system. I am going to describe some of the supported primitives in the &os; operating system. Contributed by: &a.attilio; The &os; synchronization primitive set is based on the idea to supply a rather huge number of different primitives in a way that the better one can be used for every particular, appropriate situation. To a high level point of view you can consider three kinds of synchronization primitives in the &os; kernel: atomic operations and memory barriers locks scheduling barriers Below there are descriptions for the 3 families. For every lock, you should really check the linked manpage (where possible) for more detailed explanations. Atomic operations are implemented through a set of functions performing simple aritmetics on memory operands in an atomic way with respect to external events (interrupts, preemption, etc.). Atomic operations can guarantee atomicity just on small data types (in the magnitude order of the .long. architecture C data type), so should be rarely used directly in the end-level code, if not only for very simple operations (like flag setting in a bitmap, for example). In fact, it is rather simple and common to write down a wrong semantic based on just atomic operations (usually referred as lock-less). The &os; kernel offers a way to perform atomic operations in conjunction with a memory barrier. The memory barriers will guarantee that an atomic operation will happen following some specified ordering with respect to other memory accesses. For example, if we need that an atomic operation happen just after all other pending writes (in terms of instructions reordering buffers activities) are completed, we need to explicitly use a memory barrier in conjunction to this atomic operation. So it is simple to understand why memory barriers play a key role for higher-level locks building (just as refcounts, mutexes, etc.). For a detailed explanatory on atomic operations, please refer to &man.atomic.9;. It is far, however, noting that atomic operations (and memory barriers as well) should ideally only be used for building front-ending locks (as mutexes). Refcounts are interfaces for handling reference counters. They are implemented through atomic operations and are intended to be used just for cases, where the reference counter is the only one thing to be protected, so even something like a spin-mutex is deprecated. Using the refcount interface for structures, where a mutex is already used is often wrong since we should probably close the reference counter in some already protected paths. A manpage discussing refcount does not exist currently, just check sys/refcount.h for an overview of the existing API. &os; kernel has huge classes of locks. Every lock is defined by some peculiar properties, but probably the most important is the event linked to contesting holders (or in other terms, the behaviour of threads unable to acquire the lock). &os;'s locking scheme presents three different behaviours for contenders: spinning blocking sleeping numbers are not casual Spin locks let waiters to spin until they cannot acquire the lock. An important matter do deal with is when a thread contests on a spin lock if it is not descheduled. Since the &os; kernel is preemptive, this exposes spin lock at the risk of deadlocks that can be solved just disabling interrupts while they are acquired. For this and other reasons (like lack of priority propagation support, poorness in load balancing schemes between CPUs, etc.), spin locks are intended to protect very small paths of code, or ideally not to be used at all if not explicitly requested (explained later). Block locks let waiters to be descheduled and blocked until the lock owner does not drop it and wakes up one or more contenders. In order to avoid starvation issues, blocking locks do priority propagation from the waiters to the owner. Block locks must be implemented through the turnstile interface and are intended to be the most used kind of locks in the kernel, if no particular conditions are met. Sleep locks let waiters to be descheduled and fall asleep until the lock holder does not drop it and wakes up one or more waiters. Since sleep locks are intended to protect large paths of code and to cater asynchronous events, they do not do any form of priority propagation. They must be implemented through the &man.sleepqueue.9; interface. The order used to acquire locks is very important, not only for the possibility to deadlock due at lock order reversals, but even because lock acquisition should follow specific rules linked to locks natures. If you give a look at the table above, the practical rule is that if a thread holds a lock of level n (where the level is the number listed close to the kind of lock) it is not allowed to acquire a lock of superior levels, since this would break the specified semantic for a path. For example, if a thread holds a block lock (level 2), it is allowed to acquire a spin lock (level 1) but not a sleep lock (level 3), since block locks are intended to protect smaller paths than sleep lock (these rules are not about atomic operations or scheduling barriers, however). This is a list of lock with their respective behaviours: spin mutex - spinning - &man.mutex.9; sleep mutex - blocking - &man.mutex.9; pool mutex - blocking - &man.mtx.pool.9; sleep family - sleeping - &man.sleep.9; pause tsleep msleep msleep spin msleep rw msleep sx condvar - sleeping - &man.condvar.9; rwlock - blocking - &man.rwlock.9; sxlock - sleeping - &man.sx.9; lockmgr - sleeping - &man.lockmgr.9; semaphores - sleeping - &man.sema.9; Among these locks only mutexes, sxlocks, rwlocks and lockmgrs are intended to handle recursion, but currently recursion is only supported by mutexes and lockmgrs. Scheduling barriers are intended to be used in order to drive scheduling of threading. They consist mainly of three different stubs: critical sections (and preemption) sched_bind sched_pin Generally, these should be used only in a particular context and even if they can often replace locks, they should be avoided because they do not let the diagnose of simple eventual problems with locking debugging tools (as &man.witness.4;). The &os; kernel has been made preemptive basically to deal with interrupt threads. In fact, in order to avoid high interrupt latency, time-sharing priority threads can be preempted by interrupt threads (in this way, they do not need to wait to be scheduled as the normal path previews). Preemption, however, introduces new racing points that need to be handled, as well. Often, in order to deal with preemption, the simplest thing to do is to completely disable it. A critical section defines a piece of code (borderlined by the pair of functions &man.critical.enter.9; and &man.critical.exit.9;, where preemption is guaranteed to not happen (until the protected code is fully executed). This can often replace a lock effectively but should be used carefully in order to not lose the whole advantage that preemption brings. Another way to deal with preemption is the sched_pin() interface. If a piece of code is closed in the sched_pin() and sched_unpin() pair of functions it is guaranteed that the respective thread, even if it can be preempted, it will always be executed on the same CPU. Pinning is very effective in the particular case when we have to access at per-cpu datas and we assume other threads will not change those data. The latter condition will determine a critical section as a too strong condition for our code. sched_bind is an API used in order to bind a thread to a particular CPU for all the time it executes the code, until a sched_unbind function call does not unbind it. This feature has a key role in situations where you cannot trust the current state of CPUs (for example, at very early stages of boot), as you want to avoid your thread to migrate on inactive CPUs. Since sched_bind and sched_unbind manipulate internal scheduler structures, they need to be enclosed in sched_lock acquisition/releasing when used. Various emulation layers sometimes require some additional per-process data. It can manage separate structures (a list, a tree etc.) containing these data for every process but this tends to be slow and memory consuming. To solve this problem the &os; proc structure contains p_emuldata, which is a void pointer to some emulation layer specific data. This proc entry is protected by the proc mutex. The &os; proc structure contains a p_sysent entry that identifies, which ABI this process is running. In fact, it is a pointer to the sysentvec described above. So by comparing this pointer to the address where the sysentvec structure for the given ABI is stored we can effectively determine whether the process belongs to our emulation layer. The code typically looks like: if (__predict_true(p->p_sysent != &elf_&linux;_sysvec)) return; As you can see, we effectively use the __predict_true modifier to collapse the most common case (&os; process) to a simple return operation thus preserving high performance. This code should be turned into a macro because currently it is not very flexible, i.e. we do not support &linux;64 emulation nor A.OUT &linux; processes on i386. The &os; VFS subsystem is very complex but the &linux; emulation layer uses just a small subset via a well defined API. It can either operate on vnodes or file handlers. Vnode represents a virtual vnode, i.e. representation of a node in VFS. Another representation is a file handler, which represents an opened file from the perspective of a process. A file handler can represent a socket or an ordinary file. A file handler contains a pointer to its vnode. More then one file handler can point to the same vnode. The &man.namei.9; routine is a central entry point to pathname lookup and translation. It traverses the path point by point from the starting point to the end point using lookup function, which is internal to VFS. The &man.namei.9; syscall can cope with symlinks, absolute and relative paths. When a path is looked up using &man.namei.9; it is inputed to the name cache. This behaviour can be supressed. This routine is used all over the kernel and its performance is very critical. The &man.vn.fullpath.9; function takes the best effort to traverse VFS name cache and returns a path for a given (locked) vnode. This process is unreliable but works just fine for the most common cases. The unreliability is because it relies on VFS cache (it does not traverse the on medium structures), it does not work with hardlinks, etc. This routine is used in several places in the Linuxulator. fgetvp - given a thread and a file descripton number it returns the associated vnode &man.vn.lock.9; - locks a vnode vn_unlock - unlocks a vnode &man.VOP.READDIR.9; - reads a directory referenced by a vnode &man.VOP.GETATTR.9; - gets attributes of a file or a directory referenced by a vnode &man.VOP.LOOKUP.9; - looks up a path to a given directory &man.VOP.OPEN.9; - opens a file referenced by a vnode &man.VOP.CLOSE.9; - closes a file referenced by a vnode &man.vput.9; - decrements the use count for a vnode and unlocks it &man.vrele.9; - decrements the use count for a vnode &man.vref.9; - increments the use count for a vnode fget - given a thread and a file descriptor number it returns associated file handler and references it fdrop - drops a reference to a file handler fhold - references a file handler This section deals with implementation of &linux; emulation layer in &os; operating system. It first describes the machine dependent part talking about how and where interaction between userland and kernel is implemented. It talks about syscalls, signals, ptrace, traps, stack fixup. This part discusses i386 but it is written generally so other architectures should not differ very much. The next part is the machine independent part of the Linuxulator. This section only covers i386 and ELF handling. A.OUT is obsolete and untested. Syscall handling is mostly written in linux_sysvec.c, which covers most of the routines pointed out in the sysentvec structure. When a &linux; process running on &os; issues a syscall, the general syscall routine calls linux prepsyscall routine for the &linux; ABI. &linux; passes arguments to syscalls via registers (that is why it is limited to 6 parameters on i386) while &os; uses the stack. The &linux; prepsyscall routine must copy parameters from registers to the stack. The order of the registers is: %ebx, %ecx, %edx, %esi, %edi, %ebp. The catch is that this is true for only most of the syscalls. Some (most notably clone) uses a different order but it is luckily easy to fix by inserting a dummy parameter in the linux_clone prototype. Every syscall implemented in the Linuxulator must have its prototype with various flags in syscalls.master. The form of the file is: ... AUE_FORK STD { int linux_fork(void); } ... AUE_CLOSE NOPROTO { int close(int fd); } ... The first column represents the syscall number. The second column is for auditing support. The third column represents the syscall type. It is either STD, OBSOL, NOPROTO and UNIMPL. STD is a standard syscall with full prototype and implementation. OBSOL is obsolete and defines just the prototype. NOPROTO means that the syscall is implemented elsewhere so do not prepend ABI prefix, etc. UNIMPL means that the syscall will be substituted with the nosys syscall (a syscall just printing out a message about the syscall not being implemented and returning ENOSYS). From syscalls.master a script generates three files: linux_syscall.h, linux_proto.h and linux_sysent.c. The linux_syscall.h contains definitions of syscall names and their numerical value, e.g.: ... #define LINUX_SYS_linux_fork 2 ... #define LINUX_SYS_close 6 ... The linux_proto.h contains structure definitions of arguments to every syscall, e.g.: struct linux_fork_args { register_t dummy; }; And finally, linux_sysent.c contains structure describing the system entry table, used to actually dispatch a syscall, e.g.: { 0, (sy_call_t *)linux_fork, AUE_FORK, NULL, 0, 0 }, /* 2 = linux_fork */ { AS(close_args), (sy_call_t *)close, AUE_CLOSE, NULL, 0, 0 }, /* 6 = close */ As you can see linux_fork is implemented in Linuxulator itself so the definition is of STD type and has no argument, which is exhibited by the dummy argument structure. On the other hand close is just an alias for real &os; &man.close.2; so it has no linux arguments structure associated and in the system entry table it is not prefixed with linux as it calls the real &man.close.2; in the kernel. The &linux; emulation layer is not complete, as some syscalls are not implemented properly and some are not implemented at all. The emulation layer employs a facility to mark unimplemented syscalls with the DUMMY macro. These dummy definitions reside in linux_dummy.c in a form of DUMMY(syscall);, which is then translated to various syscall auxiliary files and the implementation consists of printing a message saying that this syscall is not implemented. The UNIMPL prototype is not used because we want to be able to identify the name of the syscall that was called in order to know what syscalls are more important to implement. Signal handling is done generally in the &os; kernel for all binary compatibilities with a call to a compat-dependent layer. &linux; compatibility layer defines linux_sendsig routine for this purpose. This routine first checks whether the signal has been installed with a SA_SIGINFO in which case it calls linux_rt_sendsig routine instead. Furthermore, it allocates (or reuses an already existing) signal handle context, then it builds a list of arguments for the signal handler. It translates the signal number based on the signal translation table, assigns a handler, translates sigset. Then it saves context for the sigreturn routine (various registers, translated trap number and signal mask). Finally, it copies out the signal context to the userspace and prepares context for the actual signal handler to run. This routine is similar to linux_sendsig just the signal context preparation is different. It adds siginfo, ucontext, and some &posix; parts. It might be worth considering whether those two functions could not be merged with a benefit of less code duplication and possibly even faster execution. This syscall is used for return from the signal handler. It does some security checks and restores the original process context. It also unmasks the signal in process signal mask. Many &unix; derivates implement the &man.ptrace.2; syscall in order to allow various tracking and debugging features. This facility enables the tracing process to obtain various information about the traced process, like register dumps, any memory from the process address space, etc. and also to trace the process like in stepping an instruction or between system entries (syscalls and traps). &man.ptrace.2; also lets you set various information in the traced process (registers etc.). &man.ptrace.2; is a &unix;-wide standard implemented in most &unix;es around the world. &linux; emulation in &os; implements the &man.ptrace.2; facility in linux_ptrace.c. The routines for converting registers between &linux; and &os; and the actual &man.ptrace.2; syscall emulation syscall. The syscall is a long switch block that implements its counterpart in &os; for every &man.ptrace.2; command. The &man.ptrace.2; commands are mostly equal between &linux; and &os; so usually just a small modification is needed. For example, PT_GETREGS in &linux; operates on direct data while &os; uses a pointer to the data so after performing a (native) &man.ptrace.2; syscall, a copyout must be done to preserve &linux; semantics. The &man.ptrace.2; implementation in Linuxulator has some known weaknesses. There have been panics seen when using strace (which is a &man.ptrace.2; consumer) in the Linuxulator environment. Also PT_SYSCALL is not implemented. Whenever a &linux; process running in the emulation layer traps the trap itself is handled transparently with the only exception of the trap translation. &linux; and &os; differs in opinion on what a trap is so this is dealt with here. The code is actually very short: static int translate_traps(int signal, int trap_code) { if (signal != SIGBUS) return signal; switch (trap_code) { case T_PROTFLT: case T_TSSFLT: case T_DOUBLEFLT: case T_PAGEFLT: return SIGSEGV; default: return signal; } } The RTLD run-time link-editor expects so called AUX tags on stack during an execve so a fixup must be done to ensure this. Of course, every RTLD system is different so the emulation layer must provide its own stack fixup routine to do this. So does Linuxulator. The elf_linux_fixup simply copies out AUX tags to the stack and adjusts the stack of the user space process to point right after those tags. So RTLD works in a smart way. The &linux; emulation layer on i386 also supports &linux; A.OUT binaries. Pretty much everything described in the previous sections must be implemented for A.OUT support (beside traps translation and signals sending). The support for A.OUT binaries is no longer maintained, especially the 2.6 emulation does not work with it but this does not cause any problem, as the linux-base in ports probably do not support A.OUT binaries at all. This support will probably be removed in future. Most of the stuff necessary for loading &linux; A.OUT binaries is in imgact_linux.c file. This section talks about machine independent part of the Linuxulator. It covers the emulation infrastructure needed for &linux; 2.6 emulation, the thread local storage (TLS) implementation (on i386) and futexes. Then we talk briefly about some syscalls. One of the major areas of progress in development of &linux; 2.6 was threading. Prior to 2.6, the &linux; threading support was implemented in the linuxthreads library. The library was a partial implementation of &posix; threading. The threading was implemented using separate processes for each thread using the clone syscall to let them share the address space (and other things). The main weaknesses of this approach was that every thread had a different PID, signal handling was broken (from the pthreads perspective), etc. Also the performance was not very good (use of SIGUSR signals for threads synchronization, kernel resource consumption, etc.) so to overcome these problems a new threading system was developed and named NPTL. The NPTL library focused on two things but a third thing came along so it is usually considered a part of NPTL. Those two things were embedding of threads into a process structure and futexes. The additional third thing was TLS, which is not directly required by NPTL but the whole NPTL userland library depends on it. Those improvements yielded in much improved performance and standards conformance. NPTL is a standard threading library in &linux; systems these days. The &os; Linuxulator implementation approaches the NPTL in three main areas. The TLS, futexes and PID mangling, which is meant to simulate the &linux; threads. Further sections describe each of these areas. These sections deal with the way &linux; threads are managed and how we simulate that in &os;. The &linux; emulation layer in &os; supports runtime setting of the emulated version. This is done via &man.sysctl.8;, namely compat.linux.osrelease, which is set to 2.4.2 by default (as of April 2007) and with all &linux; versions up to 2.6 it just determined what &man.uname.1 outputs. It is different with 2.6 emulation where setting this &man.sysctl.8; affects runtime behaviour of the emulation layer. When set to 2.6.x it sets the value of linux_use_linux26 while setting to something else keeps it unset. This variable (plus per-prison variables of the very same kind) determines whether 2.6 infrastructure (mainly PID mangling) is used in the code or not. The version setting is done system-wide and this affects all &linux; processes. The &man.sysctl.8; should not be changed when running any &linux; binary as it might harm things. The semantics of &linux; threading are a little confusing and uses entirely different nomenclature to &os;. A process in &linux; consists of a struct task embedding two identifier fields - PID and TGID. PID is not a process ID but it is a thread ID. The TGID identifies a thread group in other words a process. For single-threaded process the PID equals the TGID. The thread in NPTL is just an ordinary process that happens to have TGID not equal to PID and have a group leader not equal to itself (and shared VM etc. of course). Everything else happens in the same way as to an ordinary process. There is no separation of a shared status to some external structure like in &os;. This creates some duplication of information and possible data inconsistency. The &linux; kernel seems to use task -> group information in some places and task information elsewhere and it is really not very consistent and looks error-prone. Every NPTL thread is created by a call to the clone syscall with a specific set of flags (more in the next subsection). The NPTL implements strict 1:1 threading. In &os; we emulate NPTL threads with ordinary &os; processes that share VM space, etc. and the PID gymnastic is just mimiced in the emulation specific structure attached to the process. The structure attached to the process looks like: struct linux_emuldata { pid_t pid; int *child_set_tid; /* in clone(): Child.s TID to set on clone */ int *child_clear_tid;/* in clone(): Child.s TID to clear on exit */ struct linux_emuldata_shared *shared; int pdeath_signal; /* parent death signal */ LIST_ENTRY(linux_emuldata) threads; /* list of linux threads */ }; The PID is used to identify the &os; process that attaches this structure. The child_se_tid and child_clear_tid are used for TID address copyout when a process exits and is created. The shared pointer points to a structure shared among threads. The pdeath_signal variable identifies the parent death signal and the threads pointer is used to link this structure to the list of threads. The linux_emuldata_shared structure looks like: struct linux_emuldata_shared { int refs; pid_t group_pid; LIST_HEAD(, linux_emuldata) threads; /* head of list of linux threads */ }; The refs is a reference counter being used to determine when we can free the structure to avoid memory leaks. The group_pid is to identify PID ( = TGID) of the whole process ( = thread group). The threads pointer is the head of the list of threads in the process. The linux_emuldata structure can be obtained from the process using em_find. The prototype of the function is: struct linux_emuldata *em_find(struct proc *, int locked); Here, proc is the process we want the emuldata structure from and the locked parameter determines whether we want to lock or not. The accepted values are EMUL_DOLOCK and EMUL_DOUNLOCK. More about locking later. Because of the described different view knowing what a process ID and thread ID is between &os; and &linux; we have to translate the view somehow. We do it by PID mangling. This means that we fake what a PID (=TGID) and TID (=PID) is between kernel and userland. The rule of thumb is that in kernel (in Linuxulator) PID = PID and TGID = shared -> group pid and to userland we present PID = shared -> group_pid and TID = proc -> p_pid. The PID member of linux_emuldata structure is a &os; PID. The above affects mainly getpid, getppid, gettid syscalls. Where we use PID/TGID respectively. In copyout of TIDs in child_clear_tid and child_set_tid we copy out &os; PID. The clone syscall is the way threads are created in &linux;. The syscall prototype looks like this: int linux_clone(l_int flags, void *stack, void *parent_tidptr, int dummy, void * child_tidptr); The flags parameter tells the syscall how exactly the processes should be cloned. As described above, &linux; can create processes sharing various things independently, for example two processes can share file descriptors but not VM, etc. Last byte of the flags parameter is the exit signal of the newly created process. The stack parameter if non-NULL tells, where the thread stack is and if it is NULL we are supposed to copy-on-write the calling process stack (i.e. do what normal &man.fork.2; routine does). The parent_tidptr parameter is used as an address for copying out process PID (i.e. thread id) once the process is sufficiently instantiated but is not runnable yet. The dummy parameter is here because of the very strange calling convention of this syscall on i386. It uses the registers directly and does not let the compiler do it what results in the need of a dummy syscall. The child_tidptr parameter is used as an address for copying out PID once the process has finished forking and when the process exits. The syscall itself proceeds by setting corresponding flags depending on the flags passed in. For example, CLONE_VM maps to RFMEM (sharing of VM), etc. The only nit here is CLONE_FS and CLONE_FILES because &os; does not allow setting this separately so we fake it by not setting RFFDG (copying of fd table and other fs information) if either of these is defined. This does not cause any problems, because those flags are always set together. After setting the flags the process is forked using the internal fork1 routine, the process is instrumented not to be put on a run queue, i.e. not to be set runnable. After the forking is done we possibly reparent the newly created process to emulate CLONE_PARENT semantics. Next part is creating the emulation data. Threads in &linux; does not signal their parents so we set exit signal to be 0 to disable this. After that setting of child_set_tid and child_clear_tid is performed enabling the functionality later in the code. At this point we copy out the PID to the address specified by parent_tidptr. The setting of process stack is done by simply rewriting thread frame %esp register (%rsp on amd64). Next part is setting up TLS for the newly created process. After this &man.vfork.2; semantics might be emulated and finally the newly created process is put on a run queue and copying out its PID to the parent process via clone return value is done. The clone syscall is able and in fact is used for emulating classic &man.fork.2; and &man.vfork.2; syscalls. Newer glibc in a case of 2.6 kernel uses clone to implement &man.fork.2; and &man.vfork.2; syscalls. The locking is implemented to be per-subsystem because we do not expect a lot of contention on these. There are two locks: emul_lock used to protect manipulating of linux_emuldata and emul_shared_lock used to manipulate linux_emuldata_shared. The emul_lock is a nonsleepable blocking mutex while emul_shared_lock is a sleepable blocking sx_lock. Because of the per-subsystem locking we can coalesce some locks and that is why the em find offers the non-locking access. This section deals with TLS also known as thread local storage. Threads in computer science are entities within a process that can be scheduled independently from each other. The threads in the process share process wide data (file descriptors, etc.) but also have their own stack for their own data. Sometimes there is a need for process-wide data specific to a given thread. Imagine a name of the thread in execution or something like that. The traditional &unix; threading API, pthreads provides a way to do it via &man.pthread.key.create.3;, &man.pthread.setspecific.3; and &man.pthread.getspecific.3; where a thread can create a key to the thread local data and using &man.pthread.getspecific.3; or &man.pthread.getspecific.3; to manipulate those data. You can easily see that this is not the most comfortable way this could be accomplished. So various producers of C/C++ compilers introduced a better way. They defined a new modifier keyword thread that specifies that a variable is thread specific. A new method of accessing such variables was developed as well (at least on i386). The pthreads method tends to be implemented in userspace as a trivial lookup table. The performance of such a solution is not very good. So the new method uses (on i386) segment registers to address a segment, where TLS area is stored so the actual accessing of a thread variable is just appending the segment register to the address thus addressing via it. The segment registers are usually %gs and %fs acting like segment selectors. Every thread has its own area where the thread local data are stored and the segment must be loaded on every context switch. This method is very fast and used almost exclusively in the whole i386 &unix; world. Both &os; and &linux; implement this approach and it yields very good results. The only drawback is the need to reload the segment on every context switch which can slowdown context switches. &os; tries to avoid this overhead by using only 1 segment descriptor for this while &linux; uses 3. Interesting thing is that almost nothing uses more than 1 descriptor (only Wine seems to use 2) so &linux; pays this unnecessary price for context switches. The i386 architecture implements the so called segments. A segment is a description of an area of memory. The base address (bottom) of the memory area, the end of it (ceiling), type, protection, etc. The memory described by a segment can be accessed using segment selector registers (%cs, %ds, %ss, %es, %fs, %gs). For example let us suppose we have a segment which base address is 0x1234 and length and this code: mov %edx,%gs:0x10 This will load the content of the %edx register into memory location 0x1244. Some segment registers have a special use, for example %cs is used for code segment and %ss is used for stack segment but %fs and %gs are generally unused. Segments are either stored in a global GDT table or in a local LDT table. LDT is accessed via an entry in the GDT. The LDT can store more types of segments. LDT can be per process. Both tables define upto 8191 entries. There are two main ways of setting up TLS in &linux;. It can be set when cloning a process using the clone syscall or it can call set_thread_area. When a process passes CLONE_SETTLS flag to clone, the kernel expects the memory pointed to by the %esi register a &linux; user space representation of a segment, which gets translated to the machine representation of a segment and loaded into a GDT slot. The GDT slot can be specified with a number or -1 can be used meaning that the system itself should choose the first free slot. In practice, the vast majority of programs use only one TLS entry and does not care about the number of the entry. We exploit this in the emulation and in fact depend on it. Loading of TLS for the current thread happens by calling set_thread_area while loading TLS for a second process in clone is done in the separate block in clone. Those two functions are very similar. The only difference being the actual loading of the GDT segment, which happens on the next context switch for the newly created process while set_thread_area must load this directly. The code basically does this. It copies the &linux; form segment descriptor from the userland. The code checks for the number of the descriptor but because this differs between &os; and &linux; we fake it a little. We only support indexes of 6, 3 and -1. The 6 is genuine &linux; number, 3 is genuine &os; one and -1 means autoselection. Then we set the descriptor number to constant 3 and copy out this to the userspace. We rely on the userspace process using the number from the descriptor but this works most of the time (have never seen a case where this did not work) as the userspace process typically passes in 1. Then we convert the descriptor from the &linux; form to a machine dependant form (i.e. operating system independent form) and copy this to the &os; defined segment descriptor. Finally we can load it. We assign the descriptor to threads PCB (process control block) and load the %gs segment using load_gs. This loading must be done in a critical section so that nothing can interrupt us. The CLONE_SETTLS case works exactly like this just the loading using load_gs is not performed. The segment used for this (segment number 3) is shared for this use between &os; processes and &linux; processes so the &linux; emulation layer does not add any overhead over plain &os;. The amd64 implementation is similar to the i386 one but there was initially no 32bit segment descriptor used for this purpose (hence not even native 32bit TLS users worked) so we had to add such a segment and implement its loading on every context switch (when a flag signaling use of 32bit is set). Apart from this the TLS loading is exactly the same just the segment numbers are different and the descriptor format and the loading differs slightly. Threads need some kind of synchronization and &posix; provides some of them: mutexes for mutual exclusion, read-write locks for mutual exclusion with biased ratio of reads and writes and condition variables for signaling a status change. It is interesting to note that &posix; threading API lacks support for semaphores. Those synchronization routines implementations are heavily dependant on the type threading support we have. In pure 1:M (userspace) model the implementation can be solely done in userspace and thus be very fast (the condition variables will probably end up being implemented using signals, i.e. not fast) and simple. In 1:1 model, the situation is also quite clear - the threads must be synchronized using kernel facilites (which is very slow because a syscall must be performed). The mixed M:N scenario just combines the first and second approach or rely solely on kernel. Threads synchronization is a vital part of thread-enabled programming and its performance can affect resulting program a lot. Recent benchmarks on &os; operating system showed that an improved sx_lock implementation yielded 40% speedup in ZFS (a heavy sx user), this is in-kernel stuff but it shows clearly how important the performance of synchronization primitives is. Threaded programs should be written with as little contention on locks as possible. Otherwise, instead of doing useful work the thread just waits on a lock. Because of this, the most well written threaded programs show little locks contention. &linux; implements 1:1 threading, i.e. it has to use in-kernel synchronization primitives. As stated earlier, well written threaded programs have little lock contention. So a typical sequence could be performed as two atomic increase/decrease mutex reference counter, which is very fast, as presented by the following example: pthread_mutex_lock(&mutex); .... pthread_mutex_unlock(&mutex); 1:1 threading forces us to perform two syscalls for those mutex calls, which is very slow. The solution &linux; 2.6 implements is called futexes. Futexes implement the check for contention in userspace and call kernel primitives only in a case of contention. Thus the typical case takes place without any kernel intervention. This yields reasonably fast and flexible synchronization primitives implementation. The futex syscall looks like this: int futex(void *uaddr, int op, int val, struct timespec *timeout, void *uaddr2, int val3); In this example uaddr is an address of the mutex in userspace, op is an operation we are about to perform and the other parameters have per-operation meaning. Futexes implement the following operations: FUTEX_WAIT FUTEX_WAKE FUTEX_FD FUTEX_REQUEUE FUTEX_CMP_REQUEUE FUTEX_WAKE_OP This operation verifies that on address uaddr the value val is written. If not, EWOULDBLOCK is returned, otherwise the thread is queued on the futex and gets suspended. If the argument timeout is non-zero it specifies the maximum time for the sleeping, otherwise the sleeping is infinite. This operation takes a futex at uaddr and wakes up val first futexes queued on this futex. This operations associates a file descriptor with a given futex. This operation takes val threads queued on futex at uaddr, wakes them up, and takes val2 next threads and requeues them on futex at uaddr2. This operation does the same as FUTEX_REQUEUE but it checks that val3 equals to val first. This operation performs an atomic operation on val3 (which contains coded some other value) and uaddr. Then it wakes up val threads on futex at uaddr and if the atomic operation returned a positive number it wakes up val2 threads on futex at uaddr2. The operations implemented in FUTEX_WAKE_OP: FUTEX_OP_SET FUTEX_OP_ADD FUTEX_OP_OR FUTEX_OP_AND FUTEX_OP_XOR There is no val2 parameter in the futex prototype. The val2 is taken from the struct timespec *timeout parameter for operations FUTEX_REQUEUE, FUTEX_CMP_REQUEUE and FUTEX_WAKE_OP. The futex emulation in &os; is taken from NetBSD and further extended by us. It is placed in linux_futex.c and linux_futex.h files. The futex structure looks like: struct futex { void *f_uaddr; int f_refcount; LIST_ENTRY(futex) f_list; TAILQ_HEAD(lf_waiting_paroc, waiting_proc) f_waiting_proc; }; And the structure waiting_proc is: struct waiting_proc { struct thread *wp_t; struct futex *wp_new_futex; TAILQ_ENTRY(waiting_proc) wp_list; }; A futex is obtained using the futex_get function, which searches a linear list of futexes and returns the found one or creates a new futex. When releasing a futex from the use we call the futex_put function, which decreases a reference counter of the futex and if the refcount reaches zero it is released. When a futex queues a thread for sleeping it creates a working_proc structure and puts this structure to the list inside the futex structure then it just performs a &man.tsleep.9; to suspend the thread. The sleep can be timed out. After &man.tsleep.9; returns (the thread was woken up or it timed out) the working_proc structure is removed from the list and is destroyed. All this is done in the futex_sleep function. If we got woken up from futex_wake we have wp_new_futex set so we sleep on it. This way the actual requeueing is done in this function. Waking up a thread sleeping on a futex is performed in the futex_wake function. First in this function we mimic the strange &linux; behaviour, where it wakes up N threads for all operations, the only exception is that the REQUEUE operations are performed on N+1 threads. But this usually does not make any difference as we are waking up all threads. Next in the function in the loop we wake up n threads, after this we check if there is a new futex for requeueing. If so, we requeue up to n2 threads on the new futex. This cooperates with futex_sleep. The FUTEX_WAKE_OP operation is quite complicated. First we obtain two futexes at addresses uaddr and uaddr2 then we perform the atomic operation using val3 and uaddr2. Then val waiters on the first futex is woken up and if the atomic operation condition holds we wake up val2 (i.e. timeout) waiter on the second futex. The atomic operation takes two parameters encoded_op and uaddr. The encoded operation encodes the operation itself, comparing value, operation argument, and comparing argument. The pseudocode for the operation is like this one: oldval = *uaddr2 *uaddr2 = oldval OP oparg And this is done atomically. First a copying in of the number at uaddr is performed and the operation is done. The code handles page faults and if no page fault occurs oldval is compared to cmparg argument with cmp comparator. Futex implementation uses two lock lists protecting sx_lock and global locks (either Giant or another sx_lock). Every operation is performed locked from the start to the very end. In this section I am going to describe some smaller syscalls that are worth mentioning because their implementation is not obvious or those syscalls are interesting from other point of view. During development of &linux; 2.6.16 kernel, the *at syscalls were added. Those syscalls (openat for example) work exactly like their at-less counterparts with the slight exception of the dirfd parameter. This parameter changes where the given file, on which the syscall is to be performed, is. When the filename parameter is absolute dirfd is ignored but when the path to the file is relative, it comes to the play. The dirfd paramtere is a directory relative to which the relative pathname is checked. The dirfd parameter is a file descriptor of some directory or AT_FDCWD. So for example the openat syscall can be like this: file descriptor 123 = /tmp/foo/, current working directory = /tmp/ openat(123, /tmp/bah\, flags, mode) /* opens /tmp/bah */ openat(123, bah\, flags, mode) /* opens /tmp/foo/bah */ openat(AT_FDWCWD, bah\, flags, mode) /* opens /tmp/bah */ openat(stdio, bah\, flags, mode) /* returns error because stdio is not a directory */ This infrastructure is necessary to avoid races when opening files outside the working directory. Imagine that a process consists of two threads, thread A and thread B. Thread A issues open(./tmp/foo/bah., flags, mode) and before returning it gets preempted and thread B runs. Thread B does not care about the needs of thread A and renames or removes /tmp/foo/. We got a race. To avoid this we can open /tmp/foo and use it as dirfd for openat syscall. This also enables user to implement per-thread working directories. &linux; family of *at syscalls contains: linux_openat, linux_mkdirat, linux_mknodat, linux_fchownat, linux_futimesat, linux_fstatat64, linux_unlinkat, linux_renameat, linux_linkat, linux_symlinkat, linux_readlinkat, linux_fchmodat and linux_faccessat. All these are implemented using the modified &man.namei.9; routine and simple wrapping layer. The implementation is done by altering the &man.namei.9; routine (described above) to take additional parameter dirfd in its nameidata structure, which specifies the starting point of the pathname lookup instead of using the current working directory every time. The resolution of dirfd from file descriptor number to a vnode is done in native *at syscalls. When dirfd is AT_FDCWD the dvp entry in nameidata structure is NULL but when dirfd is a different number we obtain a file for this file descriptor, check whether this file is valid and if there is vnode attached to it then we get a vnode. Then we check this vnode for being a directory. In the actual &man.namei.9; routine we simply substitute the dvp vnode for dp variable in the &man.namei.9; function, which determines the starting point. The &man.namei.9; is not used directly but via a trace of different functions on various levels. For example the openat goes like this: openat() --> kern_openat() --> vn_open() -> namei() For this reason kern_open and vn_open must be altered to incorporate the additional dirfd parameter. No compat layer is created for those because there are not many users of this and the users can be easily converted. This general implementation enables &os; to implement their own *at syscalls. This is being discussed right now. The ioctl interface is quite fragile due to its generality. We have to bear in mind that devices differ between &linux; and &os; so some care must be applied to do ioctl emulation work right. The ioctl handling is implemented in linux_ioctl.c, where linux_ioctl function is defined. This function simply iterates over sets of ioctl handlers to find a handler that implements a given command. The ioctl syscall has three parameters, the file descriptor, command and an argument. The command is a 16-bit number, which in theory is divided into high 8 bits determining class of the ioctl command and low 8 bits, which are the actual command within the given set. The emulation takes advantage of this division. We implement handlers for each set, like sound_handler or disk_handler. Each handler has a maximum command and a minimum command defined, which is used for determining what handler is used. There are slight problems with this approach because &linux; does not use the set division consistently so sometimes ioctls for a different set are inside a set they should not belong to (SCSI generic ioctls inside cdrom set, etc.). &os; currently does not implement many &linux; ioctls (compared to NetBSD, for example) but the plan is to port those from NetBSD. The trend is to use &linux; ioctls even in the native &os; drivers because of the easy porting of applications. Every syscall should be debuggable. For this purpose we introduce a small infrastructure. We have the ldebug facility, which tells whether a given syscall should be debugged (settable via a sysctl). For printing we have LMSG and ARGS macros. Those are used for altering a printable string for uniform debuging messages. As of April 2007 the &linux; emulation layer is capable of emulating the &linux; 2.6.16 kernel quite well. The remaining problems concern futexes, unfinished *at family of syscalls, problematic signals delivery, missing epoll and inotify and probably some bugs we have not discovered yet. Despite this we are capable of running basically all the &linux; programs included in &os; Ports Collection with Fedora Core 4 at 2.6.16 and there are some rudimentary reports of success with Fedora Core 6 at 2.6.16. The Fedora Core 6 linux_base was recently commited enabling some further testing of the emulation layer and giving us some more hints where we should put our effort in implementing missing stuff. We are able to run the most used applications like www/linux-firefox, www/linux-opera, net-im/skype and some games from the Ports Collection. Some of the programs exhibit bad behaviour under 2.6 emulation but this is currently under investigation and hopefully will be fixed soon. The only big application that is known not to work is the &linux; &java; Development Kit and this is because of the requirement of epoll facility which is not directly related to the &linux; kernel 2.6. We hope to enable 2.6.16 emulation by default some time after &os; 7.0 is released at least to expose the 2.6 emulation parts for some wider testing. Once this is done we can switch to Fedora Core 6 linux_base, which is the ultimate plan. Future work should focus on fixing the remaining issues with futexes, implement the rest of the *at family of syscalls, fix the signal delivery and possibly implement the epoll and inotify facilities. We hope to be able to run the most important programs flawlessly soon, so we will be able to switch to the 2.6 emulation by default and make the Fedora Core 6 the default linux_base because our currently used Fedora Core 4 is not supported any more. The other possible goal is to share our code with NetBSD and DragonflyBSD. NetBSD has some support for 2.6 emulation but its far from finished and not really tested. DragonflyBSD has expressed some interest in porting the 2.6 improvements. Generally, as &linux; develops we would like to keep up with their development, implementing newly added syscalls. Splice comes to mind first. Some already implemented syscalls are also heavily crippled, for example mremap and others. Some performance improvements can also be made, finer grained locking and others. I cooperated on this project with (in alphabetical order): &a.jhb; &a.kib; Emmanuel Dreyfus Scot Hetzel &a.jkim; &a.netchild; &a.ssouhlal; Li Xiao &a.davidxu; I would like to thank all those people for their advices, code reviews and general support. Marshall Kirk McKusick - George V. Nevile-Neil. Design and Implementation of the &os; operating system. Addison-Wesley, 2005.