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NAME

       tc-hfcs - Hierarchical Fair Service Curve

HISTORY & INTRODUCTION

       HFSC  (Hierarchical  Fair Service Curve) is a network packet scheduling algorithm that was
       first presented at SIGCOMM'97. Developed as  a  part  of  ALTQ  (ALTernative  Queuing)  on
       NetBSD,  found  its way quickly to other BSD systems, and then a few years ago became part
       of the linux kernel. Still, it's not the most popular scheduling algorithm - especially if
       compared  to HTB - and it's not well documented for the enduser. This introduction aims to
       explain how HFSC works without using  too  much  math  (although  some  math  it  will  be
       inevitable).

       In short HFSC aims to:

           1)  guarantee  precise  bandwidth  and delay allocation for all leaf classes (realtime
               criterion)

           2)  allocate excess bandwidth fairly as specified  by  class  hierarchy  (linkshare  &
               upperlimit criterion)

           3)  minimize  any  discrepancy  between  the  service  curve  and the actual amount of
               service provided during linksharing

       The main "selling" point of HFSC is feature (1), which  is  achieved  by  using  nonlinear
       service curves (more about what it actually is later). This is particularly useful in VoIP
       or games, where not only a guarantee  of  consistent  bandwidth  is  important,  but  also
       limiting  the  initial  delay of a data stream. Note that it matters only for leaf classes
       (where the actual queues are) - thus class hierarchy is ignored in the realtime case.

       Feature (2) is well, obvious - any algorithm featuring class hierarchy  (such  as  HTB  or
       CBQ)  strives  to  achieve  that. HFSC does that well, although you might end with unusual
       situations, if you define service  curves  carelessly  -  see  section  CORNER  CASES  for
       examples.

       Feature  (3)  is mentioned due to the nature of the problem. There may be situations where
       it's either not possible to guarantee service of all curves at the same time, and/or  it's
       impossible to do so fairly. Both will be explained later. Note that this is mainly related
       to interior (aka aggregate) classes, as the leafs are already handled by (1). Still,  it's
       perfectly possible to create a leaf class without realtime service, and in such a case the
       caveats will naturally extend to leaf classes as well.

ABBREVIATIONS

       For the remaining part of the document, we'll use following shortcuts:

           RT - realtime
           LS - linkshare
           UL - upperlimit
           SC - service curve

BASICS OF HFSC

       To understand how HFSC works, we must first introduce a service curve.   Overall,  it's  a
       nondecreasing  function  of some time unit, returning the amount of service (an allowed or
       allocated amount of bandwidth) at some specific point in time. The purpose of it should be
       subconsciously  obvious:  if  a  class  was  allowed  to transfer not less than the amount
       specified by its service curve, then the service curve is not violated.

       Still, we need more elaborate criterion than just the above (although in the most  generic
       case it can be reduced to it). The criterion has to take two things into account:

           ·   idling periods

           ·   the  ability  to "look back", so if during current active period the service curve
               is violated, maybe it isn't if we count excess bandwidth received  during  earlier
               active period(s)

       Let's define the criterion as follows:

           (1) For each t1, there must exist t0 in set B, so S(t1-t0) <= w(t0,t1)

       Here 'w' denotes the amount of service received during some time period between t0 and t1.
       B is a set of all times, where a session  becomes  active  after  idling  period  (further
       denoted as 'becoming backlogged'). For a clearer picture, imagine two situations:

           a)  our session was active during two periods, with a small time gap between them

           b)  as in (a), but with a larger gap

       Consider (a): if the service received during both periods meets (1), then all is well. But
       what if it doesn't do so during the 2nd period? If the amount of service  received  during
       the  1st  period  is  larger  than the service curve, then it might compensate for smaller
       service during the 2nd period and the gap - if the gap is small enough.

       If the gap is larger (b) - then it's less likely to happen (unless  the  excess  bandwidth
       allocated  during  the  1st  part  was really large). Still, the larger the gap - the less
       interesting is what happened in the past (e.g. 10 minutes  ago)  -  what  matters  is  the
       current traffic that just started.

       From HFSC's perspective, more interesting is answering the following question: when should
       we start transferring packets, so  a  service  curve  of  a  class  is  not  violated.  Or
       rephrasing  it:  How much X() amount of service should a session receive by time t, so the
       service curve is not violated. Function X() defined as below is the basic  building  block
       of  HFSC, used in: eligible, deadline, virtual-time and fit-time curves. Of course, X() is
       based on equation (1) and is defined recursively:

           ·   At the 1st backlogged period  beginning  function  X  is  initialized  to  generic
               service curve assigned to a class

           ·   At any subsequent backlogged period, X() is:
               min(X() from previous period ; w(t0)+S(t-t0) for t>=t0),
               ... where t0 denotes the beginning of the current backlogged period.

       HFSC  uses  either  linear,  or  two-piece  linear  service  curves.  In case of linear or
       two-piece linear convex functions (first slope < second slope), min()  in  X's  definition
       reduces  to the 2nd argument. But in case of two-piece concave functions, the 1st argument
       might quickly become lesser for some t>=t0. Note, that for some backlogged period, X()  is
       defined only from that period's beginning. We also define X^(-1)(w) as smallest t>=t0, for
       which X(t) = w. We have to define it this way, as X() is usually not an injection.

       The above generic X() can be one of the following:

           E() In realtime criterion, selects packets eligible for sending. If none are eligible,
               HFSC will use linkshare criterion. Eligible time 'et' is calculated with reference
               to packets' heads ( et = E^(-1)(w) ). It's based on RT service curve, but in  case
               of a convex curve, uses its 2nd slope only.

           D() In  realtime  criterion,  selects the most suitable packet from the ones chosen by
               E(). Deadline time 'dt' corresponds to packets' tails (dt = D^(-1)(w+l), where 'l'
               is packet's length). Based on RT service curve.

           V() In  linkshare  criterion,  arbitrates  which packet to send next. Note that V() is
               function of a virtual time - see LINKSHARE CRITERION section for details.  Virtual
               time  'vt'  corresponds  to  packets'  heads (vt = V^(-1)(w)). Based on LS service
               curve.

           F() An extension to linkshare criterion,  used  to  limit  at  which  speed  linkshare
               criterion  is  allowed  to dequeue. Fit-time 'ft' corresponds to packets' heads as
               well (ft = F^(-1)(w)). Based on UL service curve.

       Be sure to make clean distinction between session's RT, LS and UL service curves  and  the
       above "utility" functions.

REALTIME CRITERION

       RT   criterion  ignores  class  hierarchy  and  guarantees  precise  bandwidth  and  delay
       allocation. We say that a packet is eligible for sending, when the current  real  time  is
       later than the eligible time of the packet. From all eligible packets, the one most suited
       for sending is the one with the shortest deadline time. This sounds simple,  but  consider
       the following example:

       Interface 10Mbit, two classes, both with two-piece linear service curves:

           ·   1st class - 2Mbit for 100ms, then 7Mbit (convex - 1st slope < 2nd slope)

           ·   2nd class - 7Mbit for 100ms, then 2Mbit (concave - 1st slope > 2nd slope)

       Assume  for a moment, that we only use D() for both finding eligible packets, and choosing
       the most fitting one, thus eligible time would be computed as D^(-1)(w) and deadline  time
       would  be  computed as D^(-1)(w+l). If the 2nd class starts sending packets 1 second after
       the 1st class, it's of course impossible to guarantee 14Mbit, as the interface  capability
       is  only  10Mbit.   The only workaround in this scenario is to allow the 1st class to send
       the packets earlier that would normally be allowed. That's where  separate  E()  comes  to
       help.  Putting all the math aside (see HFSC paper for details), E() for RT concave service
       curve is just like D(), but for the RT convex service curve - it's constructed using  only
       RT service curve's 2nd slope (in our example
        7Mbit).

       The  effect  of  such E() - packets will be sent earlier, and at the same time D() will be
       updated - so the current deadline time calculated from it will be later.  Thus,  when  the
       2nd  class  starts sending packets later, both the 1st and the 2nd class will be eligible,
       but the 2nd session's deadline time will be smaller and its packets will  be  sent  first.
       When  the  1st  class  becomes  idle  at  some  later point, the 2nd class will be able to
       "buffer" up again for later active period of the 1st class.

       A short remark - in a situation, where the total amount  of  bandwidth  available  on  the
       interface  is larger than the allocated total realtime parts (imagine a 10 Mbit interface,
       but 1Mbit/2Mbit and 2Mbit/1Mbit classes), the sole speed of the interface could suffice to
       guarantee the times.

       Important  part of RT criterion is that apart from updating its D() and E(), also V() used
       by LS criterion is updated. Generally the RT criterion is secondary to LS  one,  and  used
       only   if   there's  a  risk  of  violating  precise  realtime  requirements.  Still,  the
       "participation" in bandwidth distributed by LS criterion  is  there,  so  V()  has  to  be
       updated  along  the  way.  LS  criterion  can  than properly compensate for non-ideal fair
       sharing situation, caused by RT scheduling. If you use UL service curve its  F()  will  be
       updated  as  well  (UL  service curve is an extension to LS one - see UPPERLIMIT CRITERION
       section).

       Anyway - careless specification of LS and  RT  service  curves  can  lead  to  potentially
       undesired  situations  (see CORNER CASES for examples). This wasn't the case in HFSC paper
       where LS and RT service curves couldn't be specified separately.

LINKSHARING CRITERION

       LS criterion's task is to distribute bandwidth according  to  specified  class  hierarchy.
       Contrary  to  RT  criterion, there're no comparisons between current real time and virtual
       time - the decision is based solely on direct comparison of virtual times  of  all  active
       subclasses  -  the  one  with  the  smallest  vt  wins  and  gets scheduled. One immediate
       conclusion from this fact is that absolute values don't matter - only ratios between  them
       (so for example, two children classes with simple linear 1Mbit service curves will get the
       same treatment from LS  criterion's  perspective,  as  if  they  were  5Mbit).  The  other
       conclusion is, that in perfectly fluid system with linear curves, all virtual times across
       whole class hierarchy would be equal.

       Why is VC defined in term of virtual time (and what is it)?

       Imagine an example: class A with two children - A1 and A2, both with let's say 10Mbit SCs.
       If  A2  is  idle,  A1 receives all the bandwidth of A (and update its V() in the process).
       When A2 becomes active, A1's virtual time is already far later than A2's one.  Considering
       the  type  of  decision made by LS criterion, A1 would become idle for a long time. We can
       workaround this situation by adjusting virtual time of the class becoming active -  we  do
       that  by  getting such time "up to date". HFSC uses a mean of the smallest and the biggest
       virtual time of currently active children fit for sending. As it's not real  time  anymore
       (excluding trivial case of situation where all classes become active at the same time, and
       never become idle), it's called virtual time.

       Such approach has its price though. The problem is analogous  to  what  was  presented  in
       previous section and is caused by non-linearity of service curves:

       1)  either it's impossible to guarantee service curves and satisfy fairness during certain
           time periods:

           Recall the example from RT section, slightly modified (with 3Mbit  slopes  instead  of
           2Mbit ones):

           ·   1st class - 3Mbit for 100ms, then 7Mbit (convex - 1st slope < 2nd slope)

           ·   2nd class - 7Mbit for 100ms, then 3Mbit (concave - 1st slope > 2nd slope)

           They  sum up nicely to 10Mbit - the interface's capacity. But if we wanted to only use
           LS for guarantees and fairness - it simply won't work. In LS context, only V() is used
           for  making decision which class to schedule. If the 2nd class becomes active when the
           1st one is in its second slope, the fairness will be preserved -  ratio  will  be  1:1
           (7Mbit:7Mbit),  but  LS  itself  is  of course unable to guarantee the absolute values
           themselves - as it would have to go beyond of what the interface is capable of.

       2)  and/or it's impossible to guarantee service curves of all classes  at  the  same  time
           [fairly or not]:

           This  is  similar  to  the  above  case,  but  a bit more subtle. We will consider two
           subtrees, arbitrated by their common (root here) parent:

           R (root) - 10Mbit

           A  - 7Mbit, then 3Mbit
           A1 - 5Mbit, then 2Mbit
           A2 - 2Mbit, then 1Mbit

           B  - 3Mbit, then 7Mbit

           R arbitrates between left subtree (A)  and  right  (B).  Assume  that  A2  and  B  are
           constantly  backlogged,  and at some later point A1 becomes backlogged (when all other
           classes are in their 2nd linear part).

           What happens now? B (choice made by R) will always get 7 Mbit as R is only (obviously)
           concerned  with  the ratio between its direct children. Thus A subtree gets 3Mbit, but
           its children would want (at the point when A1 became backlogged) 5Mbit + 1Mbit. That's
           of course impossible, as they can only get 3Mbit due to interface limitation.

           In  the left subtree - we have the same situation as previously (fair split between A1
           and A2, but violated guarantees), but in the whole tree - there's no fairness  (B  got
           7Mbit,  but A1 and A2 have to fit together in 3Mbit) and there's no guarantees for all
           classes (only B got what it wanted). Even if we violated fairness in the A subtree and
           set A2's service curve to 0, A1 would still not get the required bandwidth.

UPPERLIMIT CRITERION

       UL criterion is an extensions to LS one, that permits sending packets only if current real
       time is later than fit-time ('ft'). So the  modified  LS  criterion  becomes:  choose  the
       smallest  virtual  time  from  all active children, such that fit-time < current real time
       also holds. Fit-time is calculated from F(), which is based on UL service  curve.  As  you
       can  see, its role is kinda similar to E() used in RT criterion. Also, for obvious reasons
       - you can't specify UL service curve without LS one.

       The main purpose of the UL service curve is to limit HFSC to bandwidth  available  on  the
       upstream  router (think adsl home modem/router, and linux server as NAT/firewall/etc. with
       100Mbit+ connection to mentioned modem/router).  Typically, it's used to create  a  single
       class  directly under root, setting a linear UL service curve to available bandwidth - and
       then creating your class structure from that class downwards. Of course,  you're  free  to
       add a UL service curve (linear or not) to any class with LS criterion.

       An  important  part  about  the  UL service curve is that whenever at some point in time a
       class doesn't qualify for linksharing due to its fit-time, the next time it  does  qualify
       it  will  update  its virtual time to the smallest virtual time of all active children fit
       for linksharing. This way, one of the main things the LS  criterion  tries  to  achieve  -
       equality  of  all  virtual times across whole hierarchy - is preserved (in perfectly fluid
       system with only linear curves, all virtual times would be equal).

       Without that, 'vt' would lag  behind  other  virtual  times,  and  could  cause  problems.
       Consider  an  interface with a capacity of 10Mbit, and the following leaf classes (just in
       case you're skipping this text quickly - this example shows behavior that doesn't happen):

       A - ls 5.0Mbit
       B - ls 2.5Mbit
       C - ls 2.5Mbit, ul 2.5Mbit

       If B was idle, while A and C were constantly backlogged, A and C would normally (as far as
       LS  criterion  is concerned) divide bandwidth in 2:1 ratio. But due to UL service curve in
       place, C would get at most 2.5Mbit, and A would get the remaining 7.5Mbit. The longer  the
       backlogged  period,  the  more the virtual times of A and C would drift apart. If B became
       backlogged  at  some  later  point  in  time,  its  virtual   time   would   be   set   to
       (A's  vt  +  C's  vt)/2,  thus  blocking A from sending any traffic until B's virtual time
       catches up with A.

SEPARATE LS / RT SCs

       Another difference from the original HFSC paper is that RT and LS  SCs  can  be  specified
       separately.  Moreover,  leaf  classes  are allowed to have only either RT SC or LS SC. For
       interior classes, only LS SCs make sense: any RT SC will be ignored.

CORNER CASES

       Separate service curves for LS and RT criteria can lead to certain traps  that  come  from
       "fighting"  between  ideal  linksharing and enforced realtime guarantees. Those situations
       didn't exist in original HFSC paper, where specifying separate LS / RT service curves  was
       not discussed.

       Consider an interface with a 10Mbit capacity, with the following leaf classes:

       A - ls 5.0Mbit, rt 8Mbit
       B - ls 2.5Mbit
       C - ls 2.5Mbit

       Imagine A and C are constantly backlogged. As B is idle, A and C would divide bandwidth in
       2:1 ratio, considering LS service curve (so in theory - 6.66 and 3.33). Alas RT  criterion
       takes  priority,  so A will get 8Mbit and LS will be able to compensate class C for only 2
       Mbit - this will cause discrepancy between virtual times of A and C.

       Assume this situation lasts for a long time with no idle periods, and suddenly  B  becomes
       active.  B's  virtual  time will be updated to (A's vt + C's vt)/2, effectively landing in
       the middle between A's and C's virtual time. The effect - B, having no RT guarantees, will
       be punished and will not be allowed to transfer until C's virtual time catches up.

       If  the  interface  had  a higher capacity, for example 100Mbit, this example would behave
       perfectly fine though.

       Let's look a bit closer at the above example - it "cleverly" invalidates one of the  basic
       things  LS  criterion  tries  to  achieve  -  equality  of  all virtual times across class
       hierarchy. Leaf classes without RT service curves are literally left  to  their  own  fate
       (governed by messed up virtual times).

       Also,  it doesn't make much sense. Class A will always be guaranteed up to 8Mbit, and this
       is more than any absolute bandwidth that could happen from  its  LS  criterion  (excluding
       trivial case of only A being active). If the bandwidth taken by A is smaller than absolute
       value from LS criterion, the unused part will be automatically assigned  to  other  active
       classes (as A has idling periods in such case). The only "advantage" is, that even in case
       of low bandwidth on average, bursts would be handled at the speed defined by RT criterion.
       Still, if extra speed is needed (e.g. due to latency), non linear service curves should be
       used in such case.

       In the other words: the LS criterion is meaningless in the above example.

       You can quickly "workaround" it by making sure  each  leaf  class  has  RT  service  curve
       assigned  (thus  guaranteeing all of them will get some bandwidth), but it doesn't make it
       any more valid.

       Keep in mind - if you use nonlinear curves and irregularities explained above happen  only
       in the first segment, then there's little wrong with "overusing" RT curve a bit:

       A - ls 5.0Mbit, rt 9Mbit/30ms, then 1Mbit
       B - ls 2.5Mbit
       C - ls 2.5Mbit

       Here,  the vt of A will "spike" in the initial period, but then A will never get more than
       1Mbit until B & C catch up. Then everything will be back to normal.

LINUX AND TIMER RESOLUTION

       In certain situations, the scheduler can throttle itself and setup so called  watchdog  to
       wakeup dequeue function at some time later. In case of HFSC it happens when for example no
       packet is eligible for scheduling, and UL service curve is used  to  limit  the  speed  at
       which  LS criterion is allowed to dequeue packets. It's called throttling, and accuracy of
       it is dependent on how the kernel is compiled.

       There're 3 important options in  modern  kernels,  as  far  as  timers'  resolution  goes:
       'tickless system', 'high resolution timer support' and 'timer frequency'.

       If  you have 'tickless system' enabled, then the timer interrupt will trigger as slowly as
       possible, but each time a scheduler throttles itself (or any  other  part  of  the  kernel
       needs  better  accuracy),  the rate will be increased as needed / possible. The ceiling is
       either 'timer frequency' if 'high resolution  timer  support'  is  not  available  or  not
       compiled  in,  or  it's  hardware  dependent  and  can  go  far  beyond the highest 'timer
       frequency' setting available.

       If 'tickless system' is not enabled, the timer will trigger at a fixed rate  specified  by
       'timer frequency' - regardless if high resolution timers are or aren't available.

       This  is important to keep those settings in mind, as in scenario like: no tickless, no HR
       timers, frequency set to 100hz  -  throttling  accuracy  would  be  at  10ms.  It  doesn't
       automatically mean you would be limited to ~0.8Mbit/s (assuming packets at ~1KB) - as long
       as your queues are prepared to cover for timer inaccuracy. Of  course,  in  case  of  e.g.
       locally  generated  UDP traffic - appropriate socket size is needed as well. Short example
       to make it more understandable (assume hardcore anti-schedule settings  -  HZ=100,  no  HR
       timers, no tickless):

       tc qdisc add dev eth0 root handle 1:0 hfsc default 1
       tc class add dev eth0 parent 1:0 classid 1:1 hfsc rt m2 10Mbit

       Assuming  packet  of  ~1KB size and HZ=100, that averages to ~0.8Mbit - anything beyond it
       (e.g. the above example with specified rate over  10x  larger)  will  require  appropriate
       queuing  and  cause bursts every ~10 ms. As you can imagine, any HFSC's RT guarantees will
       be seriously invalidated by that.  Aforementioned example is mainly important if you  deal
       with  old hardware - as is particularly popular for home server chores. Even then, you can
       easily set HZ=1000 and have very accurate scheduling for typical adsl speeds.

       Anything modern (apic or even hpet msi based timers  +  'tickless  system')  will  provide
       enough accuracy for superb 1Gbit scheduling. For example, on one of my cheap dual-core AMD
       boards I have the following settings:

       tc qdisc add dev eth0 parent root handle 1:0 hfsc default 1
       tc class add dev eth0 parent 1:0 classid 1:1 hfsc rt m2 300mbit

       And a simple:

       nc -u dst.host.com 54321 </dev/zero
       nc -l -p 54321 >/dev/null

       ...will  yield  the  following  effects  over  a  period  of  ~10  seconds   (taken   from
       /proc/interrupts):

       319: 42124229   0  HPET_MSI-edge  hpet2 (before)
       319: 42436214   0  HPET_MSI-edge  hpet2 (after 10s.)

       That's roughly 31000/s. Now compare it with HZ=1000 setting. The obvious drawback of it is
       that cpu load can be rather high with servicing that many timer  interrupts.  The  example
       with  300Mbit RT service curve on 1Gbit link is particularly ugly, as it requires a lot of
       throttling with minuscule delays.

       Also note that it's just an example showing the capabilities  of  current  hardware.   The
       above  example  (essentially a 300Mbit TBF emulator) is pointless on an internal interface
       to begin with: you will pretty much always want a regular LS service curve there,  and  in
       such a scenario HFSC simply doesn't throttle at all.

       300Mbit RT service curve (selected columns from mpstat -P ALL 1):

       10:56:43 PM  CPU  %sys     %irq   %soft   %idle
       10:56:44 PM  all  20.10    6.53   34.67   37.19
       10:56:44 PM    0  35.00    0.00   63.00    0.00
       10:56:44 PM    1   4.95   12.87    6.93   73.27

       So,  in  the  rare  case  you need those speeds with only a RT service curve, or with a UL
       service curve: remember the drawbacks.

CAVEAT: RANDOM ONLINE EXAMPLES

       For reasons unknown (though well guessed), many examples you can google love to overuse UL
       criterion  and stuff it in every node possible. This makes no sense and works against what
       HFSC tries to do (and does pretty damn  well).  Use  UL  where  it  makes  sense:  on  the
       uppermost  node  to  match upstream router's uplink capacity. Or in special cases, such as
       testing (limit certain subtree to some speed), or customers that must never get more  than
       certain  speed.  In  the  last  case  you  can usually achieve the same by just using a RT
       criterion without LS+UL on leaf nodes.

       As for the router case - remember it's good to differentiate between "traffic  to  router"
       (remote console, web config, etc.) and "outgoing traffic", so for example:

       tc qdisc add dev eth0 root handle 1:0 hfsc default 0x8002
       tc class add dev eth0 parent 1:0 classid 1:999 hfsc rt m2 50Mbit
       tc class add dev eth0 parent 1:0 classid 1:1 hfsc ls m2 2Mbit ul m2 2Mbit

       ... so "internet" tree under 1:1 and "router itself" as 1:999

LAYER2 ADAPTATION

       Please refer to tc-stab(8)

SEE ALSO

       tc(8), tc-hfsc(8), tc-stab(8)

       Please direct bugreports and patches to: <net...@vger.kernel.org>

AUTHOR

       Manpage created by Michal Soltys (sol...@ziu.info)