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 itself.
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,
$ cd ~/cs422/lab $ git checkout --no-track -b lab4 jos/lab4 Switched to a new branch "lab4" $ git push --set-upstream origin lab4 Total 0 (delta 0), reused 0 (delta 0) To /c/cs422/SUBMIT/lab/netid.git * [new branch] lab4 -> lab4 Branch lab4 set up to track remote branch refs/remotes/origin/lab4. $ git merge lab3 Merge made by recursive. ... $
Lab 4 contains a number of new source files, some of which you should browse before you start:
||Kernel-private definitions for multiprocessor support|
||Code to read the multiprocessor configuration|
||Kernel code driving the local APIC unit in each processor|
||Assembly-language entry code for non-boot CPUs|
||Kernel-private definitions for spin locks, including the big kernel lock|
||Kernel code implementing spin locks|
||Code skeleton of the scheduler that you are about to implement|
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
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
(we use this to write to the CGA display buffer).
The LAPIC lives in a hole starting at physical address
(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
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.
kern/pmap.c. To see how this is used, look at the beginning of
kern/lapic.c. You'll have to do the next exercise, too, before the tests for
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_aps() 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
but any unused, page-aligned physical address below 640KB would work.
boot_aps() activates APs one after another,
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
After some brief setup, it puts the AP into protected mode with paging enabled,
and then calls the C setup routine
mp_main() (also in
boot_aps() waits for the AP to signal a
cpu_status field of its
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 of
kern/pmap.cto avoid adding the page at
MPENTRY_PADDRto the free list, so that we can safely copy and run AP bootstrap code at that physical address. Your code should pass the updated
check_page_free_list()test (but might fail the updated
check_kern_pgdir()test, which we will fix soon).
kern/mpentry.Sside by side with
boot/boot.S. Bearing in mind that
kern/mpentry.Sis compiled and linked to run above
KERNBASEjust like everything else in the kernel, what is the purpose of macro
MPBOOTPHYS? Why is it necessary in
kern/mpentry.Sbut not in
boot/boot.S? In other words, what could go wrong if it were omitted in
kern/mpentry.S? \ Hint: recall the differences between the link address and the load address that we have discussed in Lab 1.
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,
struct CpuInfo, which stores per-CPU variables.
cpunum() always returns the ID of the CPU that calls it,
which can be used as an index into arrays like
Alternatively, the macro
thiscpu is shorthand
for the current CPU's
Here is the per-CPU state you should be aware of:
percpu_kstacks[NCPU][KSTKSIZE]reserves space for NCPU's worth of kernel stacks.
In Lab 2, you mapped the physical memory that
bootstack 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
CPU 1's stack will start
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
and the corresponding TSS descriptor is defined
in the GDT entry
gdt[(GD_TSS0 >> 3) + i].
ts variable defined in
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
which 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,
must be executed once on each CPU. Functions
trap_init_percpu() are defined for this purpose.
kern/pmap.c) to map per-CPU stacks starting at
KSTACKTOP, as shown in
inc/memlayout.h. The size of each stack is
KSTKGAPbytes of unmapped guard pages. Your code should pass the new check in
The code in
kern/trap.c) 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
tsvariable 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
Note that until you set up IRQ handlers in part C, running code with more CPUs than processes may cause a general protection fault, since any idle CPUs will eventually receive a timer interrupt and fail to find the handler, which causes a GPF.
Our current code spins after initializing the AP in
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 kernel lock.
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
It also provides
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
unlock_kernel()at 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.
- It seems that using the big kernel lock guarantees that only one CPU can run the kernel code at a time. Why do we still need separate kernel stacks for each CPU? Describe a scenario in which using a shared kernel stack will go wrong, even with the protection of the big kernel lock.
Challenge 1 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.cis 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
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
sched_yield()as described above. Don't forget to modify
kern/init.cto 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. ...
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.
In your implementation of
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?
Whenever the kernel switches from one environment to another, it must ensure the old environment's registers are saved so they can be restored properly later. Why? Where does this happen?
Challenge 2 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.
Challenge 3 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). Extend the Env structure 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. The FXSAVE and 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.
fork() copies 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_exofork: This system call creates a new environment
with an almost blank slate:
nothing is mapped in the user portion of its address space,
and it is not runnable.
The new environment will have the same register state
as the parent environment at the time of the
In the parent,
sys_exofork will return
envid_t of 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_exofork will not actually return in the child
until the parent has explicitly allowed this
by marking the child runnable using....)
sys_env_set_status: Sets the status of a specified environment
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.
sys_page_alloc: Allocates a page of physical memory and maps it
at a given virtual address in a given environment's address space.
sys_page_map: Copy a page mapping (not the contents of a page!)
from one environment's address space to another,
leaving a memory sharing arrangement in place
so that the new and the old mappings
both refer to the same page of physical memory.
sys_page_unmap: Unmap a page mapped at a given virtual address
in a given environment.
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
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 in
kern/syscall.c. You will need to use various functions in
envid2env(). For now, whenever you call
envid2env(), pass 1 in the
checkpermparameter. Be sure you check for any invalid system call arguments, returning
-E_INVALin that case. Test your JOS kernel with
user/dumbforkand make sure it works before proceeding.
Challenge 4 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 sysyield()), 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 using syscgetc() or readline() 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.
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 that
The copying of the parent's address space into the child
is the most expensive part of the
However, a call to
fork() 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, on
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 faulting process.
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 an
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" Unix-like
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 implement first.
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
env_pgfault_upcall, to record this information.
sys_env_set_pgfault_upcallsystem 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:
ESP register starts out pointing at
and the stack data it pushes resides on the page
When a page fault occurs in user mode,
however, 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 are
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
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.
Otherwise, the kernel sets up a trap frame
on the exception stack that looks like a
struct UTrapframe from
<-- 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.
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
kern/trap.crequired 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
lib/pfentry.S. 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 kernel. 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.
user/faultread. You should see:
...  new env 00001000  user fault va 00000000 ip 0080003a TRAP frame ...  free env 00001000
user/faultdie. You should see:
...  new env 00001000 i faulted at va deadbeef, err 6  exiting gracefully  free env 00001000
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.
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/faultallocbad behave differently.
Challenge 5 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
fork() 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, while
dumbfork() copied pages,
fork() will initially only copy page mappings.
fork() will 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.
For each writable or copy-on-write page in its address space below UTOP,
the parent calls
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.
duppage sets both PTEs so that the page is not writeable,
and to contain
PTE_COW in 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.
The parent sets the user page fault entrypoint for the child to look like its own.
The child is now ready to run, so the parent marks it runnable.
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.
Test your code with the
forktreeprogram. 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'
Challenge 6 Implement a shared-memory fork() called 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. Modify user/forktree.c to use sfork() instead of regular fork(). Also, once you have finished implementing IPC in part C, use your sfork() to run user/pingpongs. You will have to find a new way to provide the functionality of the global thisenv pointer.
Challenge 7 Your implementation of fork 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. Then change fork to use this interface.
How much faster is your new fork?
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 int 0x30 instruction? How many times do you execute int 0x30 in your 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 RDTSC (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 require).
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.
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.
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
%eflags register (see
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
%eflags register 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.
kern/trap.cto initialize the appropriate entries in the IDT and provide handlers for IRQs 0 through 15. Then modify the code in
kern/env.cto 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.
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
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
trap_dispatch()function so that it calls
sched_yield()to find and run a different environment whenever a clock interrupt takes place.
You should now be able to get the
user/spintest 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
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 system calls,
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 been received.
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
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 of calling
and then looking up the information
about the received values in the current environment's
Similarly, a library function
ipc_send will take care of
sys_ipc_try_send 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
dstva 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
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.
Exercise 15 Implement
kern/syscall.c. Read the comments on both before implementing them, since they have to work together. When you call
envid2envin these routines, you should set the
checkpermflag 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.
Then implement the
user/primesfunctions to test your IPC mechanism. You might find it interesting to read
user/primes.cto see all the forking and IPC going on behind the scenes.
Challenge 8 Why does ipc_send 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 same time.
Challenge 9 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 10 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 11 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 in
Before handing in,
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',