Handed out Wednesday, October 10, 2012
Part A due Thursday, October 18, 2012
Part B due Thursday, October 25, 2012
Part C due Thursday, November 1, 2012
In this lab you will implement preemptive multitasking among multiple simultaneously active user-mode environments.
In part A you will add multiprocessor support to JOS, implement round-robin scheduling, and add basic environment management system calls (calls that create and destroy environments, and allocate/map memory).
In part B, you will implement a Unix-like
which allows a user-mode environment to create copies of
Finally, in part C you will add support for inter-process communication (IPC), allowing different user-mode environments to communicate and synchronize with each other explicitly. You will also add support for hardware clock interrupts and preemption.
Use Git to commit your Lab 3 source, fetch the latest version of the course repository, and then create a local branch called lab4 based on our lab4 branch, origin/lab4:
athena% cd ~/6.828/lab athena% add git athena% git pull Already up-to-date. athena% git checkout -b lab4 origin/lab4 Branch lab4 set up to track remote branch refs/remotes/origin/lab4. Switched to a new branch "lab4" athena% git merge lab3 Merge made by recursive. ... athena%Lab 4 contains a number of new source files, some of which you should browse before you start:
|kern/cpu.h||Kernel-private definitions for multiprocessor support|
|kern/mpconfig.c||Code to read the multiprocessor configuration|
|kern/lapic.c||Kernel code driving the local APIC unit in each processor|
|kern/mpentry.S||Assembly-language entry code for non-boot CPUs|
|kern/spinlock.h||Kernel-private definitions for spin locks, including the big kernel lock|
|kern/spinlock.c||Kernel code implementing spin locks|
|kern/sched.c||Code skeleton of the scheduler that you are about to implement|
This lab is divided into three parts, A, B, and C. We have allocated one week in the schedule for each part.
As before, you will need to do all of the regular exercises described in the lab and at least one challenge problem. (You do not need to do one challenge problem per part, just one for the whole lab.) Additionally, you will need to write up a brief description of the challenge problem that you implemented. If you implement more than one challenge problem, you only need to describe one of them in the write-up, though of course you are welcome to do more. Place the write-up in a file called answers-lab4.txt in the top level of your lab directory before handing in your work.
In the first part of this lab, you will first extend JOS to run on a multiprocessor system, and then implement some new JOS kernel system calls to allow user-level environments to create additional new environments. You will also implement cooperative round-robin scheduling, allowing the kernel to switch from one environment to another when the current environment voluntarily relinquishes the CPU (or exits). Later in part C you will implement preemptive scheduling, which allows the kernel to re-take control of the CPU from an environment after a certain time has passed even if the environment does not cooperate.
We are going to make JOS support "symmetric multiprocessing" (SMP), a multiprocessor model in which all CPUs have equivalent access to system resources such as memory and I/O buses. While all CPUs are functionally identical in SMP, during the boot process they can be classified into two types: the bootstrap processor (BSP) is responsible for initializing the system and for booting the operating system; and the application processors (APs) are activated by the BSP only after the operating system is up and running. Which processor is the BSP is determined by the hardware and the BIOS. Up to this point, all your existing JOS code has been running on the BSP.
In an SMP system, each CPU has an accompanying local APIC (LAPIC) unit. The LAPIC units are responsible for delivering interrupts throughout the system. The LAPIC also provides its connected CPU with a unique identifier. In this lab, we make use of the following basic functionality of the LAPIC unit (in kern/lapic.c):
STARTUPinterprocessor interrupt (IPI) from the BSP to the APs to bring up other CPUs (see
A processor accesses its LAPIC using memory-mapped I/O (MMIO). In MMIO, a portion of physical memory is hardwired to the registers of some I/O devices, so the same load/store instructions typically used to access memory can be used to access device registers. You've already seen one IO hole at physical address 0xA0000 (we use this to write to the CGA display buffer). The LAPIC lives in a hole starting at physical address 0xFE000000 (32MB short of 4GB), so it's too high for us to access using our usual direct map at KERNBASE. The JOS virtual memory map leaves a 4MB gap at MMIOBASE so we have a place to map devices like this. Since later labs introduce more MMIO regions, you'll write a simple function to allocate space from this region and map device memory to it.
mmio_map_region in kern/pmap.c. To
see how this is used, look at the beginning of
lapic_init in kern/lapic.c. You'll have to do
the next exercise, too, before the tests for
mmio_map_region will run.
Before booting up APs, the BSP should first collect information
about the multiprocessor system, such as the total number of
CPUs, their APIC IDs and the MMIO address of the LAPIC unit.
mp_init() function in kern/mpconfig.c
retrieves this information by reading the MP configuration
table that resides in the BIOS's region of memory.
boot_aps() function (in kern/init.c) drives
the AP bootstrap process. APs start in real mode, much like how the
bootloader started in boot/boot.S, so
copies the AP entry code (kern/mpentry.S) to a memory
location that is addressable in the real mode. Unlike with the
bootloader, we have some control over where the AP will start
executing code; we copy the entry code to 0x7000
MPENTRY_PADDR), but any unused, page-aligned
physical address below 640KB would work.
boot_aps() activates APs one after another, by
STARTUP IPIs to the LAPIC unit of the corresponding
AP, along with an initial
CS:IP address at which the AP
should start running its entry code (
MPENTRY_PADDR in our
case). The entry code in kern/mpentry.S is quite similar to
that of boot/boot.S. After some brief setup, it puts the AP
into protected mode with paging enabled, and then calls the C setup
mp_main() (also in kern/init.c).
boot_aps() waits for the AP to signal a
CPU_STARTED flag in
cpu_status field of
struct CpuInfo before going on to wake up the next one.
kern/init.c, and the assembly code in
kern/mpentry.S. Make sure you understand the control flow
transfer during the bootstrap of APs. Then modify your implementation
page_init() in kern/pmap.c to avoid adding
the page at
MPENTRY_PADDR to the free list, so that we
can safely copy and run AP bootstrap code at that physical address.
Your code should pass the updated
test (but might fail the updated
test, which we will fix soon).
KERNBASEjust like everything else in the kernel, what is the purpose of macro
MPBOOTPHYS? Why is it necessary in kern/mpentry.S but not in boot/boot.S? In other words, what could go wrong if it were omitted in kern/mpentry.S?
When writing a multiprocessor OS, it is important to distinguish
between per-CPU state that is private to each processor, and global
state that the whole system shares. kern/cpu.h defines most
of the per-CPU state, including
struct CpuInfo, which stores
cpunum() always returns the ID of the
CPU that calls it, which can be used as an index into arrays like
cpus. Alternatively, the macro
shorthand for the current CPU's
Here is the per-CPU state you should be aware of:
Per-CPU kernel stack.
Because multiple CPUs can trap into the kernel simultaneously, we need a separate kernel stack for each processor to prevent them from interfering with each other's execution. The array
percpu_kstacks[NCPU][KSTKSIZE] reserves space for NCPU's
worth of kernel stacks.
In Lab 2, you mapped the physical memory that
refers to as the BSP's kernel stack just below
Similarly, in this lab, you will map each CPU's kernel stack into this
region with guard pages acting as a buffer between them. CPU 0's
stack will still grow down from
KSTACKTOP; CPU 1's stack
KSTKGAP bytes below the bottom of CPU 0's
stack, and so on. inc/memlayout.h shows the mapping layout.
Per-CPU TSS and TSS descriptor.
A per-CPU task state segment (TSS) is also needed in order to specify where each CPU's kernel stack lives. The TSS for CPU i is stored in
cpus[i].cpu_ts, and the corresponding TSS descriptor is
defined in the GDT entry
gdt[(GD_TSS0 >> 3) + i]. The
ts variable defined in kern/trap.c will
no longer be useful.
Per-CPU current environment pointer.
Since each CPU can run different user process simultaneously, we redefined the symbol
curenv to refer to
points to the environment currently executing on the
current CPU (the CPU on which the code is running).
Per-CPU system registers.
All registers, including system registers, are private to a CPU. Therefore, instructions that initialize these registers, such as
lidt(), etc., must
be executed once on each CPU. Functions
trap_init_percpu() are defined for this purpose.
mem_init_mp() (in kern/pmap.c) to map
per-CPU stacks starting
KSTACKTOP, as shown in
inc/memlayout.h. The size of each stack is
KSTKSIZE bytes plus
KSTKGAP bytes of
unmapped guard pages. Your code should pass the new check in
The code in
initializes the TSS and
TSS descriptor for the BSP. It worked in Lab 3, but is incorrect
when running on other CPUs. Change the code so that it can work
on all CPUs. (Note: your new code should not use the global
ts variable any more.)
When you finish the above exercises, run JOS in QEMU with 4 CPUs using make qemu CPUS=4 (or make qemu-nox CPUS=4), you should see output like this:
... Physical memory: 66556K available, base = 640K, extended = 65532K check_page_alloc() succeeded! check_page() succeeded! check_kern_pgdir() succeeded! check_page_installed_pgdir() succeeded! SMP: CPU 0 found 4 CPU(s) enabled interrupts: 1 2 SMP: CPU 1 starting SMP: CPU 2 starting SMP: CPU 3 starting
Our current code spins after initializing the AP in
mp_main(). Before letting the AP get any further, we need
to first address race conditions when multiple CPUs run kernel code
simultaneously. The simplest way to achieve this is to use a big
The big kernel lock is a single global lock that is held whenever an
environment enters kernel mode, and is released when the environment
returns to user mode. In this model, environments in user mode can run
concurrently on any available CPUs, but no more than one environment can
run in kernel mode; any other environments that try to enter kernel mode
are forced to wait.
kern/spinlock.h declares the big kernel lock, namely
kernel_lock. It also provides
unlock_kernel(), shortcuts to acquire and
release the lock. You should apply the big kernel lock at four locations:
i386_init(), acquire the lock before the BSP wakes up the other CPUs.
mp_main(), acquire the lock after initializing the AP, and then call
sched_yield()to start running environments on this AP.
trap(), acquire the lock when trapped from user mode. To determine whether a trap happened in user mode or in kernel mode, check the low bits of the
env_run(), release the lock right before switching to user mode. Do not do that too early or too late, otherwise you will experience races or deadlocks.
Apply the big kernel lock as described above, by calling
the proper locations.
How to test if your locking is correct? You can't at this moment! But you will be able to after you implement the scheduler in the next exercise.
Challenge! The big kernel lock is simple and easy to use. Nevertheless, it eliminates all concurrency in kernel mode. Most modern operating systems use different locks to protect different parts of their shared state, an approach called fine-grained locking. Fine-grained locking can increase performance significantly, but is more difficult to implement and error-prone. If you are brave enough, drop the big kernel lock and embrace concurrency in JOS!
It is up to you to decide the locking granularity (the amount of data that a lock protects). As a hint, you may consider using spin locks to ensure exclusive access to these shared components in the JOS kernel:
Your next task in this lab is to change the JOS kernel so that it can alternate between multiple environments in "round-robin" fashion. Round-robin scheduling in JOS works as follows:
sched_yield()in the new kern/sched.c is responsible for selecting a new environment to run. It searches sequentially through the
envsarray in circular fashion, starting just after the previously running environment (or at the beginning of the array if there was no previously running environment), picks the first environment it finds with a status of
ENV_RUNNABLE(see inc/env.h), and calls
env_run()to jump into that environment.
sched_yield()must never run the same environment on two CPUs at the same time. It can tell that an environment is currently running on some CPU (possibly the current CPU) because that environment's status will be
sys_yield(), which user environments can call to invoke the kernel's
sched_yield()function and thereby voluntarily give up the CPU to a different environment.
Implement round-robin scheduling in
as described above. Don't forget to modify
syscall() to dispatch
Modify kern/init.c to create three (or more!) environments that all run the program user/yield.c. You should see the environments switch back and forth between each other five times before terminating, like this:
... Hello, I am environment 00001000. Hello, I am environment 00001001. Hello, I am environment 00001002. Back in environment 00001000, iteration 0. Back in environment 00001001, iteration 0. Back in environment 00001002, iteration 0. Back in environment 00001000, iteration 1. Back in environment 00001001, iteration 1. Back in environment 00001002, iteration 1. ...
After the yield programs exit, there will be no runnable environment in the system, the scheduler should invoke the JOS kernel monitor. If any of this does not happen, then fix your code before proceeding.
env_run()you should have called
lcr3(). Before and after the call to
lcr3(), your code makes references (at least it should) to the variable
e, the argument to
env_run. Upon loading the
%cr3register, the addressing context used by the MMU is instantly changed. But a virtual address (namely
e) has meaning relative to a given address context--the address context specifies the physical address to which the virtual address maps. Why can the pointer
ebe dereferenced both before and after the addressing switch?
Challenge! Add a less trivial scheduling policy to the kernel, such as a fixed-priority scheduler that allows each environment to be assigned a priority and ensures that higher-priority environments are always chosen in preference to lower-priority environments. If you're feeling really adventurous, try implementing a Unix-style adjustable-priority scheduler or even a lottery or stride scheduler. (Look up "lottery scheduling" and "stride scheduling" in Google.)
Write a test program or two
that verifies that your scheduling algorithm is working correctly
(i.e., the right environments get run in the right order).
It may be easier to write these test programs
once you have implemented
fork() and IPC
in parts B and C of this lab.
The JOS kernel currently does not allow applications
to use the x86 processor's x87 floating-point unit (FPU),
MMX instructions, or Streaming SIMD Extensions (SSE).
to provide a save area for the processor's floating point state,
and extend the context switching code
to save and restore this state properly
when switching from one environment to another.
FXRSTOR instructions may be useful,
but note that these are not in the old i386 user's manual
because they were introduced in more recent processors.
Write a user-level test program
that does something cool with floating-point.
Although your kernel is now capable of running and switching between multiple user-level environments, it is still limited to running environments that the kernel initially set up. You will now implement the necessary JOS system calls to allow user environments to create and start other new user environments.
Unix provides the
fork() system call
as its process creation primitive.
the entire address space of calling process (the parent)
to create a new process (the child).
The only differences between the two observable from user space
are their process IDs and parent process IDs
(as returned by
In the parent,
fork() returns the child's process ID,
while in the child,
fork() returns 0.
By default, each process gets its own private address space, and
neither process's modifications to memory are visible to the other.
You will provide a different, more primitive
set of JOS system calls
for creating new user-mode environments.
With these system calls you will be able to implement
fork() entirely in user space,
in addition to other styles of environment creation.
The new system calls you will write for JOS are as follows:
sys_exoforkcall. In the parent,
sys_exoforkwill return the
envid_tof the newly created environment (or a negative error code if the environment allocation failed). In the child, however, it will return 0. (Since the child starts out marked as not runnable,
sys_exoforkwill not actually return in the child until the parent has explicitly allowed this by marking the child runnable using....)
ENV_NOT_RUNNABLE. This system call is typically used to mark a new environment ready to run, once its address space and register state has been fully initialized.
For all of the system calls above that accept environment IDs,
the JOS kernel supports the convention
that a value of 0 means "the current environment."
This convention is implemented by
We have provided a very primitive implementation
of a Unix-like
in the test program user/dumbfork.c.
This test program uses the above system calls
to create and run a child environment
with a copy of its own address space.
The two environments
then switch back and forth using
as in the previous exercise.
The parent exits after 10 iterations,
whereas the child exits after 20.
Implement the system calls described above
You will need to use various functions
in kern/pmap.c and kern/env.c,
For now, whenever you call
pass 1 in the
Be sure you check for any invalid system call arguments,
-E_INVAL in that case.
Test your JOS kernel with user/dumbfork
and make sure it works before proceeding.
Add the additional system calls necessary
to read all of the vital state of an existing environment
as well as set it up.
Then implement a user mode program that forks off a child environment,
runs it for a while (e.g., a few iterations of
then takes a complete snapshot or checkpoint
of the child environment,
runs the child for a while longer,
and finally restores the child environment to the state it was in
at the checkpoint
and continues it from there.
Thus, you are effectively "replaying"
the execution of the child environment from an intermediate state.
Make the child environment perform some interaction with the user
so that the user can view and mutate its internal state,
and verify that with your checkpoint/restart
you can give the child environment a case of selective amnesia,
making it "forget" everything that happened beyond a certain point.
This completes Part A of the lab; check it using make grade and hand it in using make handin as usual. If you are trying to figure out why a particular test case is failing, run ./grade-lab4 -v, which will show you the output of the kernel builds and QEMU runs for each test, until a test fails. When a test fails, the script will stop, and then you can inspect jos.out to see what the kernel actually printed.
As mentioned earlier,
Unix provides the
fork() system call
as its primary process creation primitive.
fork() system call
copies the address space of the calling process (the parent)
to create a new process (the child).
xv6 Unix implements
fork() by copying all data from the
parent's pages into new pages allocated for the child.
This is essentially the same approach
The copying of the parent's address space into the child is
the most expensive part of the
However, a call to
is frequently followed almost immediately
by a call to
exec() in the child process,
which replaces the child's memory with a new program.
This is what the the shell typically does, for example.
In this case,
the time spent copying the parent's address space is largely wasted,
because the child process will use
very little of its memory before calling
For this reason,
later versions of Unix took advantage
of virtual memory hardware
to allow the parent and child to share
the memory mapped into their respective address spaces
until one of the processes actually modifies it.
This technique is known as copy-on-write.
To do this,
fork() the kernel would
copy the address space mappings
from the parent to the child
instead of the contents of the mapped pages,
and at the same time mark the now-shared pages read-only.
When one of the two processes tries to write to one of these shared pages,
the process takes a page fault.
At this point, the Unix kernel realizes that the page
was really a "virtual" or "copy-on-write" copy,
and so it makes a new, private, writable copy of the page for the
In this way, the contents of individual pages aren't actually copied
until they are actually written to.
This optimization makes a
fork() followed by
exec() in the child much cheaper:
the child will probably only need to copy one page
(the current page of its stack)
before it calls
In the next piece of this lab, you will implement a "proper"
fork() with copy-on-write,
as a user space library routine.
fork() and copy-on-write support in user space
has the benefit that the kernel remains much simpler
and thus more likely to be correct.
It also lets individual user-mode programs
define their own semantics for
A program that wants a slightly different implementation
(for example, the expensive always-copy version like
or one in which the parent and child actually share memory afterward)
can easily provide its own.
A user-level copy-on-write
fork() needs to know about
page faults on write-protected pages, so that's what you'll
Copy-on-write is only one of many possible uses
for user-level page fault handling.
It's common to set up an address space so that page faults indicate when some action needs to take place. For example, most Unix kernels initially map only a single page in a new process's stack region, and allocate and map additional stack pages later "on demand" as the process's stack consumption increases and causes page faults on stack addresses that are not yet mapped. A typical Unix kernel must keep track of what action to take when a page fault occurs in each region of a process's space. For example, a fault in the stack region will typically allocate and map new page of physical memory. A fault in the program's BSS region will typically allocate a new page, fill it with zeroes, and map it. In systems with demand-paged executables, a fault in the text region will read the corresponding page of the binary off of disk and then map it.
This is a lot of information for the kernel to keep track of. Instead of taking the traditional Unix approach, you will decide what to do about each page fault in user space, where bugs are less damaging. This design has the added benefit of allowing programs great flexibility in defining their memory regions; you'll use user-level page fault handling later for mapping and accessing files on a disk-based file system.
In order to handle its own page faults,
a user environment will need to register
a page fault handler entrypoint with the JOS kernel.
The user environment registers its page fault entrypoint
via the new
sys_env_set_pgfault_upcall system call.
We have added a new member to the
to record this information.
sys_env_set_pgfault_upcall system call.
Be sure to enable permission checking
when looking up the environment ID of the target environment,
since this is a "dangerous" system call.
During normal execution,
a user environment in JOS
will run on the normal user stack:
its ESP register starts out pointing at
and the stack data it pushes resides on the page
When a page fault occurs in user mode,
the kernel will restart the user environment
running a designated user-level page fault handler
on a different stack,
namely the user exception stack.
In essence, we will make the JOS kernel
implement automatic "stack switching"
on behalf of the user environment,
in much the same way that the x86 processor
already implements stack switching on behalf of JOS
when transferring from user mode to kernel mode!
The JOS user exception stack is also one page in size,
and its top is defined to be at virtual address
so the valid bytes of the user exception stack
While running on this exception stack,
the user-level page fault handler
can use JOS's regular system calls to map new pages or adjust mappings
so as to fix whatever problem originally caused the page fault.
Then the user-level page fault handler returns,
via an assembly language stub,
to the faulting code on the original stack.
Each user environment that wants to support user-level page fault handling
will need to allocate memory for its own exception stack,
sys_page_alloc() system call introduced in part A.
You will now need to change the page fault handling code in kern/trap.c to handle page faults from user mode as follows. We will call the state of the user environment at the time of the fault the trap-time state.
If there is no page fault handler registered,
the JOS kernel destroys the user environment with a message as before.
the kernel sets up a trap frame on the exception stack that looks like
struct UTrapframe from inc/trap.h:
<-- UXSTACKTOP trap-time esp trap-time eflags trap-time eip trap-time eax start of struct PushRegs trap-time ecx trap-time edx trap-time ebx trap-time esp trap-time ebp trap-time esi trap-time edi end of struct PushRegs tf_err (error code) fault_va <-- %esp when handler is run
The kernel then arranges for the user environment to resume execution with the page fault handler running on the exception stack with this stack frame; you must figure out how to make this happen. The fault_va is the virtual address that caused the page fault.
If the user environment is already running on the user exception stack
when an exception occurs,
then the page fault handler itself has faulted.
In this case,
you should start the new stack frame just under the current
tf->tf_esp rather than at
You should first push an empty 32-bit word, then a
To test whether
tf->tf_esp is already on the user
exception stack, check whether it is in the range
Implement the code in
required to dispatch page faults to the user-mode handler.
Be sure to take appropriate precautions
when writing into the exception stack.
(What happens if the user environment runs out of space
on the exception stack?)
Next, you need to implement the assembly routine that will
take care of calling the C page fault handler and resume
execution at the original faulting instruction.
This assembly routine is the handler that will be registered
with the kernel using
The interesting part is returning to the original point in
the user code that caused the page fault.
You'll return directly there, without going back through
The hard part is simultaneously switching stacks and
re-loading the EIP.
Finally, you need to implement the C user library side of the user-level page fault handling mechanism.
Run user/faultread. You should see:
...  new env 00001000  user fault va 00000000 ip 0080003a TRAP frame ...  free env 00001000
Run user/faultdie. You should see:
...  new env 00001000 i faulted at va deadbeef, err 6  exiting gracefully  free env 00001000
Run user/faultalloc. You should see:
...  new env 00001000 fault deadbeef this string was faulted in at deadbeef fault cafebffe fault cafec000 this string was faulted in at cafebffe  exiting gracefully  free env 00001000
If you see only the first "this string" line, it means you are not handling recursive page faults properly.
Run user/faultallocbad. You should see:
...  new env 00001000  user_mem_check assertion failure for va deadbeef  free env 00001000
Make sure you understand why user/faultalloc and user/faultallocbad behave differently.
Challenge! Extend your kernel so that not only page faults, but all types of processor exceptions that code running in user space can generate, can be redirected to a user-mode exception handler. Write user-mode test programs to test user-mode handling of various exceptions such as divide-by-zero, general protection fault, and illegal opcode.
You now have the kernel facilities
to implement copy-on-write
entirely in user space.
We have provided a skeleton for your
fork() should create a new environment,
then scan through the parent environment's entire address space
and set up corresponding page mappings in the child.
The key difference is that,
dumbfork() copied pages,
fork() will initially only copy page mappings.
copy each page only when one of the environments tries to write it.
The basic control flow for
fork() is as follows:
pgfault()as the C-level page fault handler, using the
set_pgfault_handler()function you implemented above.
sys_exofork()to create a child environment.
duppage, which should map the page copy-on-write into the address space of the child and then remap the page copy-on-write in its own address space.
duppagesets both PTEs so that the page is not writeable, and to contain
PTE_COWin the "avail" field to distinguish copy-on-write pages from genuine read-only pages.
The exception stack is not remapped this way, however. Instead you need to allocate a fresh page in the child for the exception stack. Since the page fault handler will be doing the actual copying and the page fault handler runs on the exception stack, the exception stack cannot be made copy-on-write: who would copy it?
fork() also needs to handle pages that are
present, but not writable or copy-on-write.
Each time one of the environments writes a copy-on-write page that it hasn't yet written, it will take a page fault. Here's the control flow for the user page fault handler:
_pgfault_upcall, which calls
pgfault()checks that the fault is a write (check for
FEC_WRin the error code) and that the PTE for the page is marked
PTE_COW. If not, panic.
pgfault()allocates a new page mapped at a temporary location and copies the contents of the faulting page contents into it. Then the fault handler maps the new page at the appropriate address with read/write permissions, in place of the old read-only mapping.
pgfault in lib/fork.c.
Test your code with the forktree program. It should produce the following messages, with interspersed 'new env', 'free env', and 'exiting gracefully' messages. The messages may not appear in this order, and the environment IDs may be different.
1000: I am '' 1001: I am '0' 2000: I am '00' 2001: I am '000' 1002: I am '1' 3000: I am '11' 3001: I am '10' 4000: I am '100' 1003: I am '01' 5000: I am '010' 4001: I am '011' 2002: I am '110' 1004: I am '001' 1005: I am '111' 1006: I am '101'
Implement a shared-memory
sfork(). This version should have the parent
and child share all their memory pages
(so writes in one environment appear in the other)
except for pages in the stack area,
which should be treated in the usual copy-on-write manner.
sfork() instead of regular
Also, once you have finished implementing IPC in part C,
sfork() to run user/pingpongs.
You will have to find a new way to provide the functionality
of the global
Your implementation of
makes a huge number of system calls. On the x86, switching into
the kernel using interrupts has non-trivial cost. Augment the
system call interface
so that it is possible to send a batch of system calls at once.
fork to use this interface.
How much faster is your new
You can answer this (roughly) by using analytical
arguments to estimate how much of an improvement batching
system calls will make to the performance of your
fork: How expensive is an
instruction? How many times do you execute
fork? Is accessing the TSS stack
switch also expensive? And so on...
Alternatively, you can boot your kernel on real hardware
and really benchmark your code. See the
(read time-stamp counter) instruction, defined in the IA32
manual, which counts the number of clock cycles that have
elapsed since the last processor reset. QEMU doesn't emulate
this instruction faithfully (it can either count the number of
virtual instructions executed or use the host TSC, neither of
which reflects the number of cycles a real CPU would
This ends part B. As usual, you can grade your submission with make grade and hand it in with make handin.
In the final part of lab 4 you will modify the kernel to preempt uncooperative environments and to allow environments to pass messages to each other explicitly.
Run the user/spin test program. This test program forks off a child environment, which simply spins forever in a tight loop once it receives control of the CPU. Neither the parent environment nor the kernel ever regains the CPU. This is obviously not an ideal situation in terms of protecting the system from bugs or malicious code in user-mode environments, because any user-mode environment can bring the whole system to a halt simply by getting into an infinite loop and never giving back the CPU. In order to allow the kernel to preempt a running environment, forcefully retaking control of the CPU from it, we must extend the JOS kernel to support external hardware interrupts from the clock hardware.
External interrupts (i.e., device interrupts) are referred to as IRQs.
There are 16 possible IRQs, numbered 0 through 15.
The mapping from IRQ number to IDT entry is not fixed.
pic_init in picirq.c maps IRQs 0-15
to IDT entries
IRQ_OFFSET is defined to be decimal 32.
Thus the IDT entries 32-47 correspond to the IRQs 0-15.
For example, the clock interrupt is IRQ 0.
Thus, IDT[IRQ_OFFSET+0] (i.e., IDT) contains the address of
the clock's interrupt handler routine in the kernel.
IRQ_OFFSET is chosen so that the device interrupts
do not overlap with the processor exceptions,
which could obviously cause confusion.
(In fact, in the early days of PCs running MS-DOS,
IRQ_OFFSET effectively was zero,
which indeed caused massive confusion between handling hardware interrupts
and handling processor exceptions!)
In JOS, we make a key simplification compared to xv6 Unix.
External device interrupts are always disabled
when in the kernel (and, like xv6, enabled when in user space).
External interrupts are controlled by the
FL_IF flag bit
When this bit is set, external interrupts are enabled.
While the bit can be modified in several ways,
because of our simplification, we will handle it solely
through the process of saving and restoring
as we enter and leave user mode.
You will have to ensure that the
FL_IF flag is set in
user environments when they run so that when an interrupt arrives, it
gets passed through to the processor and handled by your interrupt code.
Otherwise, interrupts are masked,
or ignored until interrupts are re-enabled.
We masked interrupts with the very first instruction of the bootloader,
and so far we have never gotten around to re-enabling them.
Modify kern/trapentry.S and kern/trap.c to
initialize the appropriate entries in the IDT and provide
handlers for IRQs 0 through 15. Then modify the code
env_alloc() in kern/env.c to ensure
that user environments are always run with interrupts enabled.
The processor never pushes an error code or checks the Descriptor Privilege Level (DPL) of the IDT entry when invoking a hardware interrupt handler. You might want to re-read section 9.2 of the 80386 Reference Manual, or section 5.8 of the IA-32 Intel Architecture Software Developer's Manual, Volume 3, at this time.
After doing this exercise, if you run your kernel with any test program that runs for a non-trivial length of time (e.g., spin), you should see the kernel print trap frames for hardware interrupts. While interrupts are now enabled in the processor, JOS isn't yet handling them, so you should see it misattribute each interrupt to the currently running user environment and destroy it. Eventually it should run out of environments to destroy and drop into the monitor.
In the user/spin program, after the child environment was first run, it just spun in a loop, and the kernel never got control back. We need to program the hardware to generate clock interrupts periodically, which will force control back to the kernel where we can switch control to a different user environment.
The calls to
i386_init in init.c),
which we have written for you,
set up the clock and the interrupt controller to generate interrupts.
You now need to write the code to handle these interrupts.
Modify the kernel's
so that it calls
to find and run a different environment
whenever a clock interrupt takes place.
You should now be able to get the user/spin test to work:
the parent environment should fork off the child,
sys_yield() to it a couple times
but in each case regain control of the CPU after one time slice,
and finally kill the child environment and terminate gracefully.
This is a great time to do some regression testing. Make sure that you haven't broken any earlier part of that lab that used to work (e.g. forktree) by enabling interrupts. Also, try running with multiple CPUs using make CPUS=2 target. You should also be able to pass stresssched now. Run make grade to see for sure. You should now get a total score of 65/75 points on this lab.
(Technically in JOS this is "inter-environment communication" or "IEC", but everyone else calls it IPC, so we'll use the standard term.)
We've been focusing on the isolation aspects of the operating system, the ways it provides the illusion that each program has a machine all to itself. Another important service of an operating system is to allow programs to communicate with each other when they want to. It can be quite powerful to let programs interact with other programs. The Unix pipe model is the canonical example.
There are many models for interprocess communication. Even today there are still debates about which models are best. We won't get into that debate. Instead, we'll implement a simple IPC mechanism and then try it out.
You will implement a few additional JOS kernel system calls
that collectively provide a simple interprocess communication mechanism.
You will implement two
Then you will implement two library wrappers
The "messages" that user environments can send to each other using JOS's IPC mechanism consist of two components: a single 32-bit value, and optionally a single page mapping. Allowing environments to pass page mappings in messages provides an efficient way to transfer more data than will fit into a single 32-bit integer, and also allows environments to set up shared memory arrangements easily.
To receive a message, an environment calls
This system call de-schedules the current
environment and does not run it again until a message has
When an environment is waiting to receive a message,
any other environment can send it a message -
not just a particular environment,
and not just environments that have a parent/child arrangement
with the receiving environment.
In other words, the permission checking that you implemented in Part A
will not apply to IPC,
because the IPC system calls are carefully designed so as to be "safe":
an environment cannot cause another environment to malfunction
simply by sending it messages
(unless the target environment is also buggy).
To try to send a value, an environment calls
sys_ipc_try_send with both the receiver's
environment id and the value to be sent. If the named
environment is actually receiving (it has called
sys_ipc_recv and not gotten a value yet),
then the send delivers the message and returns 0. Otherwise
the send returns
-E_IPC_NOT_RECV to indicate
that the target environment is not currently expecting
to receive a value.
A library function
ipc_recv in user space will take care
sys_ipc_recv and then looking up
the information about the received values in the current
Similarly, a library function
take care of repeatedly calling
until the send succeeds.
When an environment calls
with a valid
dstva parameter (below
the environment is stating that it is willing to receive a page mapping.
If the sender sends a page,
then that page should be mapped at
in the receiver's address space.
If the receiver already had a page mapped at
then that previous page is unmapped.
When an environment calls
with a valid
it means the sender wants to send the page
currently mapped at
srcva to the receiver,
After a successful IPC,
the sender keeps its original mapping
for the page at
srcva in its address space,
but the receiver also obtains a mapping for this same physical page
dstva originally specified by the receiver,
in the receiver's address space.
As a result this page becomes shared between the sender and receiver.
If either the sender or the receiver does not indicate
that a page should be transferred,
then no page is transferred.
After any IPC
the kernel sets the new field
in the receiver's
to the permissions of the page received,
or zero if no page was received.
sys_ipc_try_send in kern/syscall.c.
Read the comments on both before implementing them, since they
have to work together.
When you call
envid2env in these routines, you should
checkperm flag to 0,
meaning that any environment is allowed to send
IPC messages to any other environment,
and the kernel does no special permission checking
other than verifying that the target envid is valid.
Use the user/pingpong and user/primes functions to test your IPC mechanism. You might find it interesting to read user/primes.c to see all the forking and IPC going on behind the scenes.
have to loop? Change the system call interface so it
doesn't have to. Make sure you can handle multiple
environments trying to send to one environment at the
Challenge! The prime sieve is only one neat use of message passing between a large number of concurrent programs. Read C. A. R. Hoare, ``Communicating Sequential Processes,'' Communications of the ACM 21(8) (August 1978), 666-667, and implement the matrix multiplication example.
Challenge! One of the most impressive examples of the power of message passing is Doug McIlroy's power series calculator, described in M. Douglas McIlroy, ``Squinting at Power Series,'' Software--Practice and Experience, 20(7) (July 1990), 661-683. Implement his power series calculator and compute the power series for sin(x+x^3).
Challenge! Make JOS's IPC mechanism more efficient by applying some of the techniques from Liedtke's paper, "Improving IPC by Kernel Design", or any other tricks you may think of. Feel free to modify the kernel's system call API for this purpose, as long as your code is backwards compatible with what our grading scripts expect.
This ends part C. Make sure you pass all of the make grade tests and don't forget to write up your answers to the questions and a description of your challenge exercise solution in answers-lab4.txt.
Before handing in, use git status and git diff to examine your changes and don't forget to git add answers-lab4.txt. When you're ready, commit your changes with git commit -am 'my solutions to lab 4', then make handin and follow the directions.