I have a very specific need:
to partially replace the content of a flash and to move MTD partition boundaries.
Current map is:
u-boot 0x000000 0x040000
u-boot-env 0x040000 0x010000
kernel 0x050000 0x230000
initrd 0x280000 0x170000
scripts 0x3f0000 0x010000
filesystem 0x400000 0xbf0000
firmware 0xff0000 0x010000
While desired output is:
u-boot 0x000000 0x040000
u-boot-env 0x040000 0x010000
kernel 0x050000 0x230000
filesystem 0x280000 0xd70000
firmware 0xff0000 0x010000
This means to collapse initrd, scripts and filesystem into a single area while leaving the others alone.
Problem is this should be achieved from the running system (booted with the "old" configuration") and I should rewrite kernel and "new" filesystem before rebooting.
The system is an embedded, so I have little space for maneuver (I have a SD card, though).
Of course the rewritten kernel will have "new" configuration written in its DTB.
Problem is transition.
Note: I have seen this Question, but it is very old and it has drawback to need kernel patches, which I would like to avoid.
NOTE2: this question has been flagged for deletion because "not about programming". I beg to disagree: I need to perform said operation on ~14k devices, most of them already sold to customers, so any workable solution should involve, at the very least, scripting.
NOTE3: if absolutely necessary I can even consider (small) kernel modifications (YES, I have means to update kernel remotely).
I will leave the Accepted answer as-is, but, for anyone who happens to come here to find a solution, I want to point out that:
Recent (<4 years old) mtd-utils, coupled with 4.0+ kernel support:
Definition of a "master" device (MTD device representing the full, unpartitioned Flash). This is a kernel option.
mtd-utils has a specific mtd-part utility that can add/delete MTD partitions dynamically. NOTE: this utility woks IF (and only if) the above is defined in Kernel.
With the above utility it's possible to build multiple, possibly overlapping partitions; use with care!
I have three ideas/suggestions:
Instead of moving the partitions, can you just split the "new" filesystem image into chunks and write them to the corresponding "old" MTD partitions? This way you don't really need to change MTD partition map. After booting into the new kernel, it will see the new contiguous root filesystem. For JFFS2 filesystem, it should be fairly straightforward to do using split or dd, flash_erase and nandwrite. Something like:
# WARNING: this script assumes that it runs from tmpfs and the old root filesystem is already unmounted.
# Also, it assumes that your shell has arithmetic evaluation, which handles hex (my busybox 1.29 ash does this).
# assuming newrootfs.img is the image of new rootfs
new_rootfs_img="newrootfs.img"
mtd_initrd="/dev/mtd3"
mtd_initrd_size=0x170000
mtd_scripts="/dev/mtd4"
mtd_scripts_size=0x010000
mtd_filesystem="/dev/mtd5"
mtd_filesystem_size=0xbf0000
# prepare chunks of new filesystem image
bs="0x1000"
# note: using arithmetic evaluation $(()) to convert from hex and do the math.
# dd doesn't handle hex numbers ("dd: invalid number '0x1000'") -- $(()) works this around
dd if="${new_rootfs_img}" of="rootfs_initrd" bs=$(( bs )) count=$(( mtd_initrd_size / bs ))
dd if="${new_rootfs_img}" of="rootfs_scripts" bs=$(( bs )) count=$(( mtd_scripts_size / bs )) skip=$(( mtd_initrd_size / bs ))
dd if="${new_rootfs_img}" of="rootfs_filesystem" bs=$(( bs )) count=$(( mtd_filesystem_size / bs )) skip=$(( ( mtd_initrd_size + mtd_scripts_size ) / bs ))
# there's no going back after this point
flash_eraseall -j "${mtd_initrd}"
flash_eraseall -j "${mtd_scripts}"
flash_eraseall -j "${mtd_filesystem}"
nandwrite -p "${mtd_initrd}" rootfs_initrd
nandwrite -p "${mtd_scripts}" rootfs_scripts
nandwrite -p "${mtd_filesystem}" rootfs_filesystem
# don't forget to update the kernel too
There is kernel support for concatenating MTD devices (which is exactly what you're trying to do). I don't see an easy way to use it, but you could create a kernel module, which concatenates the desired partitions for you into a contiguous MTD device.
In order to combine the 3 MTD partitions into one to write the new filesystem, you could create a dm-linear mapping over the 3 mtdblocks, and then turn it back into an MTD device using block2mtd. (i.e. mtdblock + device mapper linear + block2mtd) But it looks very awkward and I don't know if it'll work well (for say, OOB data).
EDIT1: added a comment explaining use of $(( bs )) -- to convert from hex as dd doesn't handle hex numbers directly (neither coreutils, nor busybox dd).
AFAIK, #andrey 's answer suggestion 1 is wrong.
an mtd partition is made of a sequence of blocks, any of which could be bad or go bad anytime. this is why the simple mtd char abstraction exists: an mtd char device (not the mtdblock one) is read sequentially and skips bad blocks. nandwrite also writes sequentially and skips bad blocks.
an mtd char device sort of acts like:
a single file into which you cannot random access, from which you can only read sequentially from the beginning to the end (or to where you get bored).
a single file into which you cannot random access, to which you can only write sequentially from the beginning (or from an erase block where you previously stopped reading) all the way to the end. (that is, you can truncate and append, but you cannot write mid-file.) to write you need to previously erase all erase blocks from where you start writing to the end of the partition.
this means that the partition size is the maximum theoretical capacity, but typically the capacity will be less due to bad blocks, and can be effectively reduced every time you rewrite the partition. you can never expect to write the full size of an mtd partition.
this is were #andrey 's suggestion 1 is wrong: it breaks up the file to be written into max-sized pieces before writing each piece. but you never know beforehand how much data will fit into an mtd partition without actually writing that data.
instead, you typically need to write some data, and you pray there will be enough good blocks to fit it. if at some point there are not, the write fails and the device reached end-of-life. needless to say, the larger the fraction of a partition you need, the higher the likelihood that the write will fail (and when that happens, it typically means that the device is toast).
to actually implement something akin to suggestion 1, you need to start writing into a partition (skipping bad blocks), and when you run out of erase blocks, you continue writing into the next partition, and so on. the point being: you cannot know where the data boundaries will lay until you actually write the data and fill each partition; there is no other way.
Related
I want to copy files from one place to another and the problem is I deal with a lot of sparse files.
Is there any (easy) way of copying sparse files without becoming huge at the destination?
My basic code:
out, err := os.Create(bricks[0] + "/" + fileName)
in, err := os.Open(event.Name)
io.Copy(out, in)
Some background theory
Note that io.Copy() pipes raw bytes – which is sort of understandable once you consider that it pipes data from an io.Reader to an io.Writer which provide Read([]byte) and Write([]byte), correspondingly.
As such, io.Copy() is able to deal with absolutely any source providing
bytes and absolutely any sink consuming them.
On the other hand, the location of the holes in a file is a "side-channel" information which "classic" syscalls such as read(2) hide from their users.
io.Copy() is not able to convey such side-channel information in any way.
IOW, initially, file sparseness was an idea to just have efficient storage of the data behind the user's back.
So, no, there's no way io.Copy() could deal with sparse files in itself.
What to do about it
You'd need to go one level deeper and implement all this using the syscall package and some manual tinkering.
To work with holes, you should use the SEEK_HOLE and SEEK_DATA special values for the lseek(2) syscall which are, while formally non-standard, are supported by all major platforms.
Unfortunately, the support for those "whence" positions is not present
neither in the stock syscall package (as of Go 1.8.1)
nor in the golang.org/x/sys tree.
But fear not, there are two easy steps:
First, the stock syscall.Seek() is actually mapped to lseek(2)
on the relevant platforms.
Next, you'd need to figure out the correct values for SEEK_HOLE and
SEEK_DATA for the platforms you need to support.
Note that they are free to be different between different platforms!
Say, on my Linux system I can do simple
$ grep -E 'SEEK_(HOLE|DATA)' </usr/include/unistd.h
# define SEEK_DATA 3 /* Seek to next data. */
# define SEEK_HOLE 4 /* Seek to next hole. */
…to figure out the values for these symbols.
Now, say, you create a Linux-specific file in your package
containing something like
// +build linux
const (
SEEK_DATA = 3
SEEK_HOLE = 4
)
and then use these values with the syscall.Seek().
The file descriptor to pass to syscall.Seek() and friends
can be obtained from an opened file using the Fd() method
of os.File values.
The pattern to use when reading is to detect regions containing data, and read the data from them – see this for one example.
Note that this deals with reading sparse files; but if you'd want to actually transfer them as sparse – that is, with keeping this property of them, – the situation is more complicated: it appears to be even less portable, so some research and experimentation is due.
On Linux, it appears you could try to use fallocate(2) with
FALLOC_FL_PUNCH_HOLE | FALLOC_FL_KEEP_SIZE to try to punch a hole at the
end of the file you're writing to; if that legitimately fails
(with syscall.EOPNOTSUPP), you just shovel as many zeroed blocks to the destination file as covered by the hole you're reading – in the hope
the OS will do the right thing and will convert them to a hole by itself.
Note that some filesystems do not support holes at all – as a concept.
One example is the filesystems in the FAT family.
What I'm leading you to is that inability of creating a sparse file might
actually be a property of the target filesystem in your case.
You might find Go issue #13548 "archive/tar: add support for writing tar containing sparse files" to be of interest.
One more note: you might also consider checking whether the destination directory to copy a source file resides in the same filesystem as the source file, and if this holds true, use the syscall.Rename() (on POSIX systems)
or os.Rename() to just move the file across different directories w/o
actually copying its data.
You don't need to resort to syscalls.
package main
import "os"
func main() {
f, _ := os.Create("/tmp/sparse.dat")
f.Write([]byte("start"))
f.Seek(1024*1024*10, 0)
f.Write([]byte("end"))
}
Then you'll see:
$ ls -l /tmp/sparse.dat
-rw-rw-r-- 1 soren soren 10485763 Jun 25 14:29 /tmp/sparse.dat
$ du /tmp/sparse.dat
8 /tmp/sparse.dat
It's true you can't use io.Copy as is. Instead you need to implement an alternative to io.Copy which reads a chunk from the src, checks if it's all '\0'. If it is, just dst.Seek(len(chunk), os.SEEK_CUR) to skip past that part in dst. That particular implementation is left as an exercise to the reader :)
I've been trying to read the implementation of a kernel module, and I'm stumbling on this piece of code.
unsigned long addr = (unsigned long) buf;
if (!IS_ALIGNED(addr, 1 << 9)) {
DMCRIT("#%s in %s is not sector-aligned. I/O buffer must be sector-aligned.", name, caller);
BUG();
}
The IS_ALIGNED macro is defined in the kernel source as follows:
#define IS_ALIGNED(x, a) (((x) & ((typeof(x))(a) - 1)) == 0)
I understand that data has to be aligned along the size of a datatype to work, but I still don't understand what the code does.
It left-shifts 1 by 9, then subtracts by 1, which gives 111111111. Then 111111111 does bitwise-and with x.
Why does this code work? How is this checking for byte alignment?
In systems programming it is common to need a memory address to be aligned to a certain number of bytes -- that is, several lowest-order bits are zero.
Basically, !IS_ALIGNED(addr, 1 << 9) checks whether addr is on a 512-byte (2^9) boundary (the last 9 bits are zero). This is a common requirement when erasing flash locations because flash memory is split into large blocks which must be erased or written as a single unit.
Another application of this I ran into. I was working with a certain DMA controller which has a modulo feature. Basically, that means you can allow it to change only the last several bits of an address (destination address in this case). This is useful for protecting memory from mistakes in the way you use a DMA controller. Problem it, I initially forgot to tell the compiler to align the DMA destination buffer to the modulo value. This caused some incredibly interesting bugs (random variables that have nothing to do with the thing using the DMA controller being overwritten... sometimes).
As far as "how does the macro code work?", if you subtract 1 from a number that ends with all zeroes, you will get a number that ends with all ones. For example, 0b00010000 - 0b1 = 0b00001111. This is a way of creating a binary mask from the integer number of required-alignment bytes. This mask has ones only in the bits we are interested in checking for zero-value. After we AND the address with the mask containing ones in the lowest-order bits we get a 0 if any only if the lowest 9 (in this case) bits are zero.
"Why does it need to be aligned?": This comes down to the internal makeup of flash memory. Erasing and writing flash is a much less straightforward process then reading it, and typically it requires higher-than-logic-level voltages to be supplied to the memory cells. The circuitry required to make write and erase operations possible with a one-byte granularity would waste a great deal of silicon real estate only to be used rarely. Basically, designing a flash chip is a statistics and tradeoff game (like anything else in engineering) and the statistics work out such that writing and erasing in groups gives the best bang for the buck.
At no extra charge, I will tell you that you will be seeing a lot of this type of this type of thing if you are reading driver and kernel code. It may be helpful to familiarize yourself with the contents of this article (or at least keep it around as a reference): https://graphics.stanford.edu/~seander/bithacks.html
There are product_uuid and product_serial files in dir /sys/class/dmi/id/.
How it are generated? What is the difference?
Can I change this files?
Is it save a value after reinstall operation system?
How it are generated?
Those values are generated in kernel code. You can find them pretty easily using git grep command (with keywords you are interested in) in your kernel source directory:
$ git grep --all-match -n -e '\bdmi\b' -e product_uuid -e product_serial
So, product_uuid and product_serial sysfs nodes are created in drivers/firmware/dmi-id.c:
DEFINE_DMI_ATTR_WITH_SHOW(product_serial, 0400, DMI_PRODUCT_SERIAL);
DEFINE_DMI_ATTR_WITH_SHOW(product_uuid, 0400, DMI_PRODUCT_UUID);
From DEFINE_DMI_ATTR_WITH_SHOW definition you can see that both attributes are accessed via sys_dmi_field_show() function, which in turn calls dmi_get_system_info(), which just returns corresponding element from dmi_ident array. This table is populated in dmi_decode() routine:
dmi_save_ident(dm, DMI_PRODUCT_SERIAL, 7);
dmi_save_uuid(dm, DMI_PRODUCT_UUID, 8);
So product_uuid is generated in dmi_save_uuid() function. Just read its code to understand how it's done.
product_serial is generated in dmi_save_ident() function. It boils down to code like this:
(struct dmi_header *)(dmi_base)[7];
where dmi_base is address (remapped to virtual memory obviously) of DMI table, and 7 corresponds to DMI_PRODUCT_SERIAL constant.
To better understand this please see SMBIOS specification, specifically Table 9 – System Information (Type 1) Structure, which corresponds to this command:
# dmidecode --type 1
What is the difference?
As for product_uuid -- look at SMBIOS specification, section 7.2.1 System - UUID. It has description and also table with explanation for each part of this number. Using that table you can decode your UUID and extract some information from it, like timestamp, etc.
As for product_serial -- I believe it's self-explanatory, it's just a serial number for your device. You can usually find it printed on some sticker on your computer. For example, for my laptop it's on the bottom. It's the same string that I see in /sys/class/dmi/id/product_serial.
Can I change this files?
Those files are actually not real files but just an interface to kernel functions. Read about sysfs for details. So in order to "change" those files you need to edit mentioned kernel files accordingly, then rebuild the whole kernel and boot it (instead of one provided by your distribution).
Also, as #ChristopheVu-Brugier mentioned in comment, you can change those values in DMI table (in some tricky way though). But I wouldn't recommend it. Those values definitely have some meaning and may be useful in some cases (if not for you, then for some software in your PC).
Is it save a value after reinstall operation system?
Those values are actually obtained from DMI table, which is hardcoded along with BIOS to permanent memory (flash chip with BIOS on your motherboard) and you just read those values from this DMI table using kernel functions by reading those files.
I'm trying to create a large empty file on a VFAT partition by using the `dd' command in an embedded linux box:
dd if=/dev/zero of=/mnt/flash/file bs=1M count=1 seek=1023
The intention was to skip the first 1023 blocks and write only 1 block at the end of the file, which should be very quick on a native EXT3 partition, and it indeed is. However, this operation turned out to be quite slow on a VFAT partition, along with the following message:
lowmem_shrink:: nr_to_scan=128, gfp_mask=d0, other_free=6971, min_adj=16
// ... more `lowmem_shrink' messages
Another attempt was to fopen() a file on the VFAT partition and then fseek() to the end to write the data, which has also proved slow, along with the same messages from the kernel.
So basically, is there a quick way to create the file on the VFAT partition (without traversing the first 1023 blocks)?
Thanks.
Why are VFAT "skipping" writes so slow ?
Unless the VFAT filesystem driver were made to "cheat" in this respect, creating large files on FAT-type filesystems will always take a long time. The driver, to comply with FAT specification, will have to allocate all data blocks and zero-initialize them, even if you "skip" the writes. That's because of the "cluster chaining" FAT does.
The reason for that behaviour is FAT's inability to support either:
UN*X-style "holes" in files (aka "sparse files")
that's what you're creating on ext3 with your testcase - a file with no data blocks allocated to the first 1GB-1MB of it, and a single 1MB chunk of actually committed, zero-initialized blocks) at the end.
NTFS-style "valid data length" information.
On NTFS, a file can have uninitialized blocks allocated to it, but the file's metadata will keep two size fields - one for the total size of the file, another for the number of bytes actually written to it (from the beginning of the file).
Without a specification supporting either technique, the filesystem would always have to allocate and zerofill all "intermediate" data blocks if you skip a range.
Also remember that on ext3, the technique you used does not actually allocate blocks to the file (apart from the last 1MB). If you require the blocks preallocated (not just the size of the file set large), you'll have to perform a full write there as well.
How could the VFAT driver be modified to deal with this ?
At the moment, the driver uses the Linux kernel function cont_write_begin() to start even an asynchronous write to a file; this function looks like:
/*
* For moronic filesystems that do not allow holes in file.
* We may have to extend the file.
*/
int cont_write_begin(struct file *file, struct address_space *mapping,
loff_t pos, unsigned len, unsigned flags,
struct page **pagep, void **fsdata,
get_block_t *get_block, loff_t *bytes)
{
struct inode *inode = mapping->host;
unsigned blocksize = 1 << inode->i_blkbits;
unsigned zerofrom;
int err;
err = cont_expand_zero(file, mapping, pos, bytes);
if (err)
return err;
zerofrom = *bytes & ~PAGE_CACHE_MASK;
if (pos+len > *bytes && zerofrom & (blocksize-1)) {
*bytes |= (blocksize-1);
(*bytes)++;
}
return block_write_begin(mapping, pos, len, flags, pagep, get_block);
}
That is a simple strategy but also a pagecache trasher (your log messages are a consequence of the call to cont_expand_zero() which does all the work, and is not asynchronous). If the filesystem were to split the two operations - one task to do the "real" write, and another one to do the zero filling, it'd appear snappier.
The way this could be achieved while still using the default linux filesystem utility interfaces were by internally creating two "virtual" files - one for the to-be-zerofilled area, and another for the actually-to-be-written data. The real file's directory entry and FAT cluster chain would only be updated once the background task is actually complete, by linking its last cluster with the first one of the "zerofill file" and the last cluster of that one with the first one of the "actual write file". One would also want to go for a directio write to do the zerofilling, in order to avoid trashing the pagecache.
Note: While all this is technically possible for sure, the question is how worthwhile would it be to do such a change ? Who needs this operation all the time ? What would side effects be ?
The existing (simple) code is perfectly acceptable for smaller skipping writes, you won't really notice its presence if you create a 1MB file and write a single byte at the end. It'll bite you only if you go for filesizes on the order of the limits of what the FAT filesystem allows you to do.
Other options ...
In some situations, the task at hand involves two (or more) steps:
freshly format (e.g.) a SD card with FAT
put one or more big files onto it to "pre-fill" the card
(app-dependent, optional)
pre-populate the files, or
put a loopback filesystem image into them
One of the cases I've worked on we've folded the first two - i.e. modified mkdosfs to pre-allocate/ pre-create files when making the (FAT32) filesystem. That's pretty simple, when writing the FAT tables just create allocated cluster chains instead of clusters filled with the "free" marker. It's also got the advantage that the data blocks are guaranteed to be contiguous, in case your app benefits from this. And you can decide to make mkdosfs not clear the previous contents of the data blocks. If you know, for example, that one of your preparation steps involves writing the entire data anyway or doing ext3-in-file-on-FAT (pretty common thing - linux appliance, sd card for data exchange with windows app/gui), then there's no need to zero out anything / double-write (once with zeroes, once with whatever-else). If your usecase fits this (i.e. formatting the card is a useful / normal step of the "initialize it for use" process anyway) then try it out; a suitably-modified mkdosfs is part of TomTom's dosfsutils sources, see mkdosfs.c search for the -N command line option handling.
When talking about preallocation, as mentioned, there's also posix_fallocate(). Currently on Linux when using FAT, this will do essentially the same as a manual dd ..., i.e. wait for the zerofill. But the specification of the function doesn't mandate it being synchronous. The block allocation (FAT cluster chain generation) would have to be done synchronously, but the VFAT on-disk dirent size update and the data block zerofills could be backgrounded / delayed (i.e. either done at low-prio in background or only done if explicitly requested via fdsync() / sync() so that the app can e.g. alloc blocks, write the contents with non-zeroes itself ...). That's technique / design though; I'm not aware of anyone having done that kernel modification yet, if only for experimenting.
I have a script that creates file system in a file on a linux machine. I see that to create the file system, it uses 'dd' with bs=x option, reads from /dev/zero and writes to a file. I think usually specifying ibs/obs/bs is useful to read from real hardware devices as one has specific block size constraints. In this case however, as it is reading from virtual device and writing to a file, I don't see any point behind using 'bs=x bytes' option. Is my understanding wrong here?
(Just in case if it helps, this file system is later on used to boot a qemu vm)
To understand block sizes, you have to be familiar with tape drives. If you're not interested in tape drives - for example, you don't think you're ever going to use one - then you can go back to sleep now.
Remember the tape drives from films in the 60s, 70s, maybe even 80s? The ones where the reel went spinning around, and so on? Not your Exabyte or even QIC - quarter-inch cartridge - tapes; your good old fashioned reel-to-reel half-inch tape drives? On those, block size mattered.
The data on a tape was written in blocks. Each block was separated from the next by an inter-record gap.
----+-------+-----+-------+-----+----
... | block | IRG | block | IRG | ...
----+-------+-----+-------+-----+----
Depending on the tape drive hardware and software, there were a variety of problems that could happen. For example, if the tape was written with a block size of 5120 bytes and you read the tape with a block size of 512 bytes, then the tape drive might read the first block, return you 512 bytes of it, and then discard the remaining data; the next read would start on the next block. Conversely, if the tape was written with a block size of 512 bytes and you requested blocks of 5120 bytes, you would get short reads; each read would return just 512 bytes, and if your software wasn't paying attention, you'd be reading garbage. There was also the issue that the tape drive had to get up to speed to read the block, and then slow down. The ASCII art suggests that the IRG was smaller than the data blocks; that was not necessarily the case. And it took time to read one block, overshoot the IRG, rewind backwards to get to the next block, and start forwards again. And if the tape drive didn't have the memory to buffer data - the cheaper ones did not - then you could seriously affect your tape drive performance.
War story: work prepared on newer machine with a slightly more modern tape drive. I wrote a tape using tar without a sensible block size (so it defaulted to 512 bytes). It was a large bit of software - must have been, oh, less than 100 MB in total (a long time ago, in other words). The tape wrote nicely because the machine was modern enough, and it took just a few seconds to do so. But, I had to get the material off the tape on a machine with an older tape drive, one that did not have any on-board buffer. So, it read the material, 512 bytes at a time, and the reel rocked forward, reading one block, and then rocked back all but maybe half an inch, and then read forwards to get to the next block, and then rocked back, and ... well, you could see it doing this, and since it took appreciable portions of a second to read each 512 byte block, the total time taken was horrendous. My flight was due to leave...and I needed to get that data across too. (It was long enough ago, and in a land far enough away, that last minute changes to flights weren't much of an option either.) To cut a long story short, it did get read - but if I'd used a sensible block size (such as 5120 bytes instead of the default of 512), I would have been done much, much quicker and with much less danger of missing the plane (but I did actually catch the plane, with maybe 20 minutes to spare, despite a taxi ride across Paris in the rush hour).
With more modern tape drives, there was enough memory on the drive to do buffering and getting a tape drive to stream - write continuously without reversing - was feasible. It used to be that I'd use a block size like 256 KB to get QIC tapes to stream. I've not done much with tape drives recently - let's see, not this millennium and not much for a few years before that, either; certainly not much since CD and DVD became the software distribution mechanisms of choice (when electronic download wasn't used).
But the block size really did matter in the old days. And dd provided good support for it. You could even transfer data from a tape drive that was written with, say, 4 KB block to another that you wanted to write with, say, 16 KB blocks, by specifying the ibs (input block size) separately from the obs (output block size). Darned useful!
Also, the count parameter is in terms of the (input) block size. It was useful to say 'dd bs=1024 count=1024 if=/dev/zero of=/my/file/of/zeroes' to copy 1 MB of zeroes around. Or to copy 1 MB of a file.
The importance of dd is vastly diminished; it was an essential part of the armoury for anybody who worked with tape drives a decade or more ago.
The block size is the number of bytes that are read and written at a time. Presumably there is a count= option, and that is specified in units of the block size. If there is a skip= or seek= option, those will also be in block size units. However if you are reading and writing a regular file, and there are no disk errors, then the block size doesn't really matter as long as you can scale those parameters accordingly and they are still integers. However certain sizes may be more efficient than others.
For reading from /dev/zero, it doesn't matter. ibs/obs/bs specify how many bytes will be read at a time. It's helpful to choose a number based on the way bytes are read/written in the operating system. For instance, Linux usually reads from a hard drive in 4096 byte chunks. If you have at least some idea about how the underlying hardware reads/writes, it might be a good idea to specify ibs/obs/bs. By the way, if you specify bs, it will override whatever you specify for ibs and obs.
In addition to the great answer by Jonathan Leffler, keep in mind that the bs= option isn't always a substitute for using both ibs= and obs=, in particular for the old [ugly] days of tape drives.
The bs= option reserves the right for dd to write the data as soon as it's read. This can cause you to no longer have identically sized blocks on the output. Here is GNU's take on this, but the behavior dates back as far as I can remember (80's):
(bs=) Set both input and output block sizes to bytes. This makes dd read and write bytes per block, overriding any ‘ibs’ and ‘obs’ settings. In addition, if no data-transforming conv option is specified, input is copied to the output as soon as it’s read, even if it is smaller than the block size.
For instance, back in the QIC days on an old Sun system, if you did this:
tar cvf /dev/rst0c /bla
It would work, but cause an enormous amount of back and forth thrashing while the drive wrote a small block, and tried to backup and read to reposition itself properly for the next write.
If you swapped this with:
tar cvf - /bla | dd ibs=16K obs=16K of=/dev/rst0c
You'd get the QIC drive writing much larger chunks and not thrashing quite so much.
However, if you made the mistake of this:
tar cvf - /bla | dd bs=16K of=/dev/rst0c
You'd run the risk of having precisely the same thrashing you had before depending upon how much data was available at the time of each read.
Specifying both ibs= and obs= precludes this from happening.