IEN 189

                     ISSUES IN INTERNETTING

                        PART 4:  ROUTING

                          Eric C. Rosen

                  Bolt Beranek and Newman Inc.

                            June 1981

IEN 189                              Bolt Beranek and Newman Inc.
                                                    Eric C. Rosen

                     ISSUES IN INTERNETTING

                        PART 4:  ROUTING

4.  Routing

     This is the fourth in a series of papers  that  discuss  the

issues  involved  in designing an internet.  Familiarity with the

previous papers (IENs 184, 187, and 188) is presupposed.

     The topic of the present paper is routing.  We will  discuss

the  issues  involved  in  choosing  a  routing algorithm for the

internet, and  we  will  propose  a  particular  algorithm.   The

algorithm  we  propose  will  be  based  on the routing algorithm

currently operating in the ARPANET, called "SPF  routing."   This

algorithm  is  described in [1] and [2], which interested readers

will certainly want to look at.  Although we  will  try  to  make

this paper relatively self-contained, we will of course focus our

discussion  on those aspects of the algorithm which might have to

be modified to work in the internet.

     Any discussion of the proper routing algorithm to use  in  a

particular  Network  Structure must begin with a consideration of

just what characteristics we want the routing algorithm to  have.

That  is, we must decide in advance just what we want the routing

algorithm to do.  Everyone will agree that the routing  algorithm

ought  to be able to deliver data from an arbitrary source Switch

to an arbitrary  destination  Switch,  as  long  as  there  is  a

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physical  path  between them.  Or at least, the routing algorithm

should make the probability of being able to do this  arbitrarily

high.  However, this is a very minimal criterion (as indicated by

the  fact that everyone would agree to it).  There are many other

requirements we must place on the routing algorithm if we  intend

to  design  a  robust and high performance Network Structure.  We

will  present  some  requirements  and  some   possible   routing

algorithms  which fulfill the requirements to a greater or lesser

degree.  We hope that by the end of this paper, we will have made

a case that our proposed routing algorithm does a better  job  of

meeting more of the desired requirements than does any other that

we know of.

4.1  Flexibility and Topological Changes

     One extremely important, though little noticed, feature that

we  should require of a routing algorithm is that it enable us to

make arbitrary changes in the topology of the Network  Structure,

without the need to make manual changes in the internal tables of

the  Switches.   This  is a capability that has always existed in

the ARPANET.   IMPs  can  be  added,  removed,  or  moved  around

arbitrarily,  and  the  routing algorithm automatically adapts to

the new topology without any  manual  intervention.   This  seems

simple  enough, but it does place some significant constraints on

the nature of the routing algorithm.  For example, it immediately

rules out fixed routing.  By "fixed routing,"  we  refer  to  any

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scheme  where  a  set  of  routes  to  each destination Switch is

"compiled into" each Switch.  In fixed routing schemes, there  is

generally  a  "primary  route",  to  be  used  when  the  Network

Structure is not  suffering  from  any  outages,  and  a  set  of

alternate or secondary routes to be used if some component of the

primary route should fail.  We know of one network which does use

this  sort  of fixed routing, and as a result, they are forced to

adhere to a very strict rule which allows them to add  or  remove

Switches  only  once  every  six months.  Certainly, we would not

want to build such a restriction into the internet.

     Fixed routing also prohibits  certain  important  day-to-day

operational  procedures  that are often used in the ARPANET.  For

example, it is quite common, when an  IMP  is  brought  down  for

preventive  maintenance,  to "splice" that IMP out of the network

by wiring together two of its modems.  This causes two IMPs  that

ordinarily  have  a  common  neighbor  to  suddenly become direct

neighbors of  each  other.   (A  similar  function  can  also  be

performed  by  the  telephone  company,  in case the power to the

modems is shut off, or if the  site  cannot  be  reached.)   This

ability to preserve network bandwidth even when a site is down is

quite  important  to  robust network performance.  Yet it is very

difficult, if not impossible, to do this if  the  network  has  a

fixed routing algorithm.  It is not yet clear to what extent such

day-to-day  "firefighting"  techniques  will be applicable in the

internet, but it certainly  does  not  seem  wise  to  design  an

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internet  routing  algorithm  which  would  be  too inflexible to

permit the use of such techniques.

     Another very useful capability which is difficult to combine

with  fixed  routing  is  the  ability  to   create   arbitrarily

configured  test networks in the lab, and then to connect them to

the real network.  This is something that is done quite often  in

the  ARPANET,  usually  for  the  purposes  of  testing  out  new

software, and we will definitely  need  this  capability  in  the

internet in order to test out new gateway software (as well as to

test out patches and bug fixes to the old).

     It  is also worth noting that implementing a scheme of fixed

routing with a primary route and alternates to be used in case of

outages is not nearly as trivial as it may seem.   Remember  that

it  is not enough for each individual Switch, when its Pathway to

a particular neighbor fails, to pick an alternate neighbor as its

next hop to some destination.  Rather, any  outage  requires  ALL

the  Switches to pick alternates in a COORDINATED MANNER, so that

the routing produced  by  the  use  of  the  alternate  paths  is

loop-free.  This is quite a difficult problem, and if there are a

large  number  of Switches and Pathways, any combination of which

could fail, this means that a  very  large  number  of  alternate

paths  must  be maintained, requiring a consequently large amount

of table space.

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     We will not be giving much serious consideration to the  use

of  fixed routing in the internet.  We mention it largely for the

sake of completeness, and because there is  a  natural  tendency,

which  we  wish  to oppose, to suppose that fixed routing must be

simpler, cheaper and more reliable than  dynamic  routing.   This

tendency  ignores the day-to-day operational problems involved in

the use of fixed routing, as  well  as  the  difficult  technical

problems involved the the creation of fixed routing tables.

     Preserving  maximum  flexibility to make topological changes

requires the Switches to be able to determine, dynamically,  just

who  their  neighbors  are.   (Remember  that  two  Switches of a

Network Structure are neighbors if and only if they are connected

by a Pathway,  i.e.,  by  a  communications  path  containing  no

intermediate  Switch  of  the  same  Network  Structure.)  In the

ARPANET,  each  IMP  is  initialized  to  know  how  many   modem

interfaces  it  has,  and  does  not  determine that dynamically.

However, initialization only tells the IMP how many interfaces it

has; it does not tell the IMP  who  its  neighbor  is  over  each

interface.    The   IMPs   determine   who  their  neighbors  are

dynamically, via the line up/down protocol, and  a  line  between

two  IMPs  cannot come up unless and until each of the IMPs knows

the identity of the other.

     The situation in  the  present  Catenet  gateways  is  quite

different.   Each  gateway  has  a  table  of potential neighbors

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"assembled  in."   When  a  gateway  comes  up,  it  attempts  to

communicate  (via  a special gateway neighbor protocol) with each

of  the  gateways  in  its  pre-assembled  neighbor  table.   Two

gateways  are  considered neighbors only if this communication is

successful.  Gateways will also consider themselves neighbors  of

other  gateways  that  communicate  with  them  according  to the

gateway neighbor protocol, even if the other gateway  is  not  in

the  pre-assembled neighbor table.  This means that two gateways,

G1 and G2, cannot become neighbors unless either G1  is  in  G2's

pre-assembled  neighbor  table,  or  G2  is in G1's pre-assembled

neighbor table.

     Of course, in a real operational  environment,  it  is  very

important  to  ensure  that  site-dependent  information  is  not

assembled or compiled in.  Rather, it must be separately loadable

(over the network itself)  by  the  Network  Control  Center,  or

whatever  equivalent  organization  we  create  for operating the

internet.   In  fact,  site-dependent  information  ought  to  be

preserved  over  reload of site-independent information, and vice

versa.  (This discipline is followed in the ARPANET.)   Designing

the  gateways  according to this discipline is a very non-trivial

task, which must be planned for by the gateway designers  at  the

earliest  stage  of gateway design.  Otherwise, we will build for

ourselves  a  very  difficult  set  of  unnecessary   operational


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     However, it is not a very good idea to have a fixed table of

neighbors  in  each  Switch,  even  if  this  table is separately

loadable.  This just does not give us the flexibility  we  desire

for  making  arbitrary topological changes.  If there has not yet

been any difficulty with the Catenet's current  scheme,  that  is

probably  because  of  the small number of gateways and component

networks in the current internet environment.  As the  number  of

gateways  increases,  the need to have them dynamically determine

who their neighbors are becomes increasingly more important.

     However, having gateways discover  (dynamically)  who  their

neighbors  are  is  a  more  difficult  problem  than having IMPs

discover who their neighbors are.  The  interfaces  on  the  IMPs

function  as  point-to-point  lines,  so there can be at most one

other IMP on the other end of a line, and any data sent out  that

line can be expected to reach just that IMP.  Therefore it is not

very  hard  for an IMP to discover which IMP is at the other end.

An IMP simply sends its identity (a unique number which it  reads

from  its  hardware  configuration  cards)  down  the  line  in a

message, and if the line is operational, the message  must  reach

the  IMP  on  the  other  end.   For  two gateways connected by a

packet-switching  network,  the  problem  is  more   complicated,

because, unlike telephone circuits, a packet-switching network is

not   a   point-to-point   line  with  a  relatively  transparent

interface.  In order  for  one  gateway  to  identify  itself  to

another,  it  must be able to address the other, using the Access

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Protocol of the packet-switching  network  which  serves  as  the

Pathway  between  them.  This seems to mean that for a gateway to

be able to send its identity to a neighbor, it must already  know

the neighbor's name.  This seems like Catch-22 -- there is no way

to  determine  dynamically  who  your neighbor is, unless you can

address him, but there is  no  way  to  address  him  unless  you

already know who he is.

     This   problem  can  be  made  more  tractable  through  the

cooperation  of  the  packet-switching  networks  underlying  the

Pathways  which connect the gateways.  A packet-switching network

could recognize that certain of its own components  (which  might

be either Switches or Hosts within its own Network Structure) are

also  Switches  within  a  Network  Structure  which is one level

higher in a hierarchy.  For example, in the ARPANET, there  might

be   some  special  protocol  (call  it  the  "gateway  discovery

protocol"), carried out on the host-IMP level, by  which  certain

hosts  identify  themselves  as  internet  gateways.   Whenever a

gateway connected to a particular IMP comes up or goes down, this

information could be broadcast to all  other  IMPs.   Whenever  a

gateway  comes up, the IMP it is connected to could tell it which

of the other hosts are internet gateways.  In this way, the  IMPs

could  keep  the gateways informed as to which other gateways are

up  or  down  at  any  particular  time.   This  sort  of  scheme

eliminates the need for the gateways to know in advance who their

neighbors  might  be,  and  moves  the responsibility for keeping

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track  of  the  gateways  and  their  up/down   status   to   the

packet-switching  network  itself,  which  is  better equipped to

carry out this responsibility.

     Such a scheme would not be very difficult, in principle,  to

build  into  the  ARPANET.   Information  about gateways could be

subsumed into the routing information.  That is, an IMP connected

to a gateway could represent the gateway  as  a  stub  node,  and

report  on  it  as  such  in  its  ordinary routing updates.  (Of

course, this is only  feasible  if  the  number  of  gateways  is

relatively  small when compared to the number of IMPs.  Otherwise

the additional overhead this would add to the ARPANET's  internal

routing  algorithm would make the scheme infeasible.  However, it

does seem likely that the number of gateways on the ARPANET  will

always  be  much  smaller  than the number of IMPs.)  This scheme

would automatically cause the information about the  gateways  to

be  broadcast  to  all IMPs as part of the routing updates.  (See

section 4.5 for a description of the routing  update  procedure.)

Each  IMP  which is connected directly to a gateway could forward

information about other  gateways  to  its  own  gateway  as  the

information  is received.  The most difficult problem might be to

get enough "security" in the gateway-to-IMP protocol so that only

real gateways could declare themselves to be gateways.  (Some  of

the  issues  involved  in  preventing  a  host from "fooling" the

network into thinking it is a different host than  it  really  is

are  discussed  in  IEN 183.  See the discussion of LAD messages.

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However, that note does not consider the real issue  of  security

that arises here.)

     This  scheme  for having gateways dynamically discover their

neighbors through the cooperation of the networks underlying  the

internal  Pathways  of  the internet is an important step towards

the solution of the "flying gateway" problem.  The flying gateway

problem is the following.  Suppose that N is  a  packet-switching

network  which  is one of the component networks of the internet.

Now suppose that  due  to  some  sort  of  emergency  or  natural

disaster,  N  becomes partitioned into two "pieces", call them N1

and N2, and that  this  partition  is  expected  to  last  for  a

significant  amount  of time.  If H1 is a Host in N1, and H2 is a

Host in N2, then H1 and H2 will no longer be able to  communicate

through the network N.  (Of course, H1 and H2 might still be able

to  communicate  though  the  internet,  if  there is an internet

gateway on N1 and an internet gateway on N2, and a route  between

these two gateways other than the "direct" route via N.  In fact,

the  addressing  scheme  proposed  in  IEN 188 will automatically

cause traffic from H1 to H2 to be delivered over  this  alternate



TO  H2 OVER THE NETWORK N.)  However, in some cases, there may be

no such alternate route, or else  its  characteristics  might  be

unsatisfactory.   In  addition,  it  must  be remembered that the

partition of network N might actually result in the partition  of

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the internet itself, so that some pairs of Hosts which ordinarily

communicate over the internet can no longer reach each other.  In

such  cases, it might be desirable, at the level of the internet,

to treat N1 and N2 as separate component networks, and  to  place

an  internet  gateway  between  them so that internet traffic can

flow from N1 to N2.   One  possible  scenario  is  for  this  new

gateway  to  be  an airborne packet radio, hence the name "flying


     If a flying gateway can be connected to both N1 and N2,  and

if  the network N has a gateway discovery protocol of the sort we

have been advocating, then the flying gateway need merely come up

on N1 and N2, declaring itself to be an  internet  gateway.   The

gateway  discovery  protocol  run in the network pieces N1 and N2

will cause the other internet gateways in N1  and  N2  to  become

aware  that  they  have a new neighbor, the flying gateway.  Once

the gateways in N1 and N2 become aware of their new neighbor,  it

automatically begins to participate in the routing algorithm (see

section  4.5  for  details of the routing updating algorithm that

brings this about), and routing automatically begins to  use  the

flying  gateway  for store-and-forwarding internet traffic.  Thus

any partition of the internet is automatically brought to an end.

     In addition to using the flying  gateway  as  a  transit  or

intermediate  gateway  for  internet  traffic,  it  may  also  be

desirable to use it as a  destination  Switch  in  the  internet.

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That is, it may be desirable to allow the other internet Switches

(gateways) to use the flying gateway as the address to which they

route  traffic  for  Hosts  in  N1  or N2.  This is slightly more

complicated  than  simply  using  the  flying   gateway   as   an

intermediate Switch.  The logical-to-physical address translation

tables  in  the  gateways  (we are assuming the addressing scheme

proposed in IEN 188) will not, in general, map any  Host  logical

addresses into the address of the flying gateway, which after all

is  not  ordinarily  on  the  internet.   However, as long as the

flying gateway indicates that it is a special,  flying,  gateway,

and  as  long  as this information is made known to all the other

gateways, this problem is simple enough to  solve.   If  F  is  a

flying  gateway,  and  G  is an ordinary gateway, and F and G are

neighbors, then any logical address which maps to  G  but  cannot

currently  be  reached  through  any  ordinary  gateway should be

mapped to F.  (As we shall see, the routing algorithm we  propose

makes  available to each Switch all information about which pairs

of Switches are neighbors.)  Attempting to reach the  destination

Host  via the flying gateway F will either be successful, or else

should result in  the  return  of  a  DNA  message,  which  would

indicate  that the Host cannot be reached from the flying gateway

either.  The only remaining problem is  for  the  flying  gateway

itself  to  determine  which of the two pieces of the partitioned

network  contain  some  particular  Host  for  which  it  is  the

destination  Switch.   Any  data  for  destination  Host  H which

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arrives at Switch F can potentially be sent to  either  piece  of

the  partitioned network.  The situation is no different than the

problem of how an ordinary gateway, which has two Pathways  to  a

particular  Host,  one of which is non-operational, decides which

one to use.  Note that the individual Hosts do  not  need  to  be

aware  at  all  of the existence of the flying gateway, since the

logical addressing scheme automatically finds the right  physical

address.   Of  course, for this mechanism to be at all effective,

there  must  be  a  robust  and  efficient  Host-Switch   up/down

protocol,  which  works  through  the  cooperation of the network

underlying the Pathway between Host and Switch.

     Unfortunately, not every component network  of  an  internet

can  be  expected  to  cooperate this way in a "gateway discovery

protocol."  In fact, if two  Switches  of  the  internet  Network

Structure are connected by a Pathway which is itself an internet,

rather  than a single packet-switching network, then this sort of

cooperation  in  the  "gateway  discovery  protocol"   might   be

extremely  difficult  if  not  impossible.  It seems though to be

quite important to get the communications  media  which  underlie

the  Pathways  to  participate  in  such  a  protocol,  for  that

significantly increases both the reliability and the  flexibility

of  the  internetting  scheme.   It  does  not  seem possible for

Switches  which  are  connected  by  uncooperative  Pathways   to

determine dynamically who their neighbors are.  In such cases, we

may  then have to live with hand-built neighbor tables (as in the

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present Catenet), and a protocol which the  Switches  attempt  to

carry  out  with their neighbors to see which potential neighbors

are really reachable.  Networks which do not  provide  a  gateway

discovery  protocol,  however,  cannot be patched together with a

flying gateway if they should partition.

     Even  for  Switches  which  are  connected  by   cooperative

Pathways,  it  is desirable to have a protocol which the Switches

attempt to run with each one of their neighbors, to  see  whether

they  really  can send and receive data to or from each neighbor.

Suppose, for example,  that  two  Switches  are  connected  by  a

Pathway  which  is  a very congested network.  In such a network,

the messages which are used to tell  the  internet  Switches  who

their  "neighbors"  are  might  well  be flowing, even though the

congestion prevents ordinary (user) data from flowing.   This  is

not  at all unlikely, if the gateway discovery protocol makes use

of the network's routing updates, which would probably be of much

higher priority than ordinary data packets.  Since we don't  want

to  use  this  Pathway  for  internet traffic unless it can carry

data, some independent means of determining this may  be  needed.

The  situation  is  somewhat more complicated if the Pathway is a

packet-switching network with different "acceptance classes",  so

that  only  certain  classes of traffic are accepted at any given

time, depending perhaps on the internal loading conditions of the

network.  If a Pathway is only accepting a certain  sub-class  of

data  traffic,  any internet Switches which are connected to that

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Pathway must  be  able  to  determine  which  classes  are  being

accepted  (presumably  the  network  underlying  the Pathway will

inform the Switches as to  any  access  restrictions),  and  this

information  will  have  to be fed back into the internet routing

algorithm, so that traffic which cannot be placed  on  a  certain

Pathway is not routed there nonetheless.

     The   reader   will   doubtless   have  noticed  that  these

considerations, of determining who one's neighbors  are,  and  of

determining  whether the Pathway to each neighbor is operational,

are quite similar to the considerations adduced in IEN 187 in the

discussion of Pathway up/down protocols to be run between a  Host

and  a  Switch.   What  we  have  been  discussing  is  really an

inter-Switch Pathway up/down  protocol.   The  gateway  discovery

protocol  corresponds  to  what  we  called  a "low-level up/down

protocol", and the type of protocol  discussed  in  the  previous

paragraph corresponds to what we called the "higher-level up/down


4.2  Why We Cannot Require Optimality

     What else would we like the routing algorithm to do, besides

giving  us  the  maximum flexibility to make topological changes?

Generally, we tend to feel that a really good  routing  algorithm

should  optimize  something,  delay  or  throughput, for example.

However, true optimality is really not possible.  If we are given

a complete description of a network,  including  its  topological

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structure,  and  the  capacities  and speeds of all its lines and

Switches, and if we are also given the traffic requirement (as  a

Switch-Switch traffic matrix which tells us how much traffic each

Switch  will  originate which is destined for each other Switch),

and if the packet inter-arrival rates and sizes vary according to

certain specific probabilistic distributions, and if the  traffic

is in a steady-state condition, it is just a mathematical problem

to  devise  a  set  of routing tables for the Switches which will

minimize the network average delay.  Applied mathematicians  have

devoted  a great deal of effort to devising algorithms to produce

this optimal solution.  There are a large number of problems with

attempting to use this sort of  "optimal  routing  algorithm"  as

the operational routing algorithm of a network:

     1) Packet arrival rates and sizes do  not  necessarily  vary

        according  to  the  probabilistic distributions which are

        assumed by optimal routing algorithms.

     2) Optimal  routing   algorithms   are   ALWAYS   based   on

        mathematical models of the relationship between delay and

        throughput which are not supported by empirical data.

     3) Actual traffic requirements are quite variable,  and  may

        not  really  approach  a  steady-state  for a long enough

        period of time  to  enable  true  optimization.   Traffic

        requirements are also generally unknown, and difficult to

        predict or measure.

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     4) Most algorithms to compute the optimal  routes  are  real

        number   crunchers,  and  require  large  floating  point

        computers.  These algorithms would have to be  run  in  a

        central   location,  producing  routing  tables  for  all

        Switches, and then distributing them somehow (centralized

        routing), with  consequent  problems  of  robustness  and


     5) There  are  distributed  optimizing   algorithms   (e.g.,

        Gallager's  algorithm),  but  they are not implementable.

        That is, the proofs of these algorithms make  assumptions

        which  could not be made to hold in the real software and

        real hardware of a real network.   Hence  the  algorithms

        would  not  be  expected to give optimal results (or even

        anything   close   to   optimal)   in   real    networks.

        Furthermore,  such  algorithms  seem  to rely on updating

        protocols  which  are  insufficiently   robust   in   the

        operational  environment.   These algorithms also seem to

        contain  parameters  whose  precise  settings  are  quite

        important   to   proper   performance,   but  whose  most

        appropriate values are unknown  and  quite  difficult  to


     We  realize  that  these rather brusque comments may make it

seem like we are giving short  shrift  to  the  consideration  of

optimizing  algorithms.   We  have  made these comments simply in

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order to state our reasons for not giving  further  consideration

to  such  algorithms.   Arguing  in  support  of  these  reasons,

however, would require another paper.

     Another problem with optimal  routing  algorithms  which  is

more  specific  to  the  internet  environment has to do with the

requirement that the capacities  of  the  network  components  be

known.  With telephone circuits as the "links", it is possible to

assign  a  fixed  capacity and fixed propagation and transmission

delays to each  link.   With  packet-switching  networks  as  the

"links",  it  is  doubtful  that  this  even makes sense.  If two

gateways are connected by the ARPANET, there is no number we  can

assign  as  the  capacity  of the "link" connecting the gateways!

The amount of throughput that can be sent  between  two  gateways

via  the ARPANET is a highly variable quantity, with dependencies

on hundreds of other things going on within the ARPANET.   It  is

hard  enough  to  get  a  handle  on  just  what other things the

throughput of a given connection depends on; we  certainly  can't

express this dependency as a function, or assign numerical values

to  the  "capacity."   This  seems  to  mean that currently known

optimal routing algorithms are really quite  useless  within  the

context of the internet.  Of course, they are not too useful even

in  individual  networks,  when  considered  as  the  operational

routing algorithm of the network.  They are,  however,  sometimes

useful as a benchmark to which the operational routing algorithms

can  be  compared.   That is, it is a meaningful question to ask,

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"how close does SPF routing in the  ARPANET  come  to  optimal?",

where "optimal" is defined as the result produced by some optimal

algorithm,  run off-line.  Within the context of the internet, it

is difficult even to give meaning to this question.  There is  no

mathematical model of the internet to which we can appeal.

     This also raises an interesting question about the design of

the  internet topology, i.e., where to place the gateways and how

best to interconnect them.  The usual mathematical techniques for

trying to optimize network topological design  also  assume  some

fixed  assignment  of capacity to the links; it's not obvious how

such techniques can be extended to the internet.

4.3  Some Issues in SPF Routing

     Even if we give up the quest for optimal routing, there  are

still  a number of substantive things we can require of a routing

algorithm.  For example, we would  like  to  have  some  form  of

distributed  routing, rather than centralized routing, simply for

reasons of robustness.   ("Distributed  routing"  refers  to  any

routing  scheme  in  which  each  Switch computes its own routing

table.)  What this means basically is an algorithm based more  or

less  on the routing algorithm of the ARPANET, i.e., an algorithm

which runs in each Switch and computes the shortest path to  each

other  Switch,  based  upon (dynamically determined) knowledge of

the connectivity  of  the  internet  Network  Structure,  and  an

assignment   of  "length"  to  each  Pathway  that  connects  two

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Switches.  Routing algorithms of this sort can  be  characterized

by  three separable components: (a) the algorithm used to compute

the shortest path, given the assignment of lengths  to  Pathways,

(b) the algorithm used to assign a length to a given Pathway, and

(c)  the  protocol  used  by  the  Switches  for  sharing routing


     The most efficient shortest path algorithm that we  know  of

is  the  SPF  algorithm  of the ARPANET [1,3] (which is basically

just a modification of Dijkstra's shortest path  algorithm),  and

we  propose to base an internet routing algorithm on this.  There

are other algorithms for performing a shortest path  computation,

but  the  SPF  algorithm  seems  to  dominate them.  One possible

alternative to SPF would be something based  on  the  distributed

computation  of  the original ARPANET routing algorithm (which is

the basis for the current Catenet routing), but we  have  studied

that  algorithm  at  great  length  and in great detail and it is

inferior to SPF in a large variety of ways [3].  There  are  many

other shortest path algorithms (such as Floyd's algorithm, or the

algorithm advocated by Perlman in IEN 120), but the efficiency of

these  algorithms does not compare with that of SPF.  We will not

consider the issue of choosing  a  shortest  path  algorithm  any


     In  the ARPANET, the "length" assigned to a line is just the

average per-packet delay over that line during a preceding period

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of ten seconds.  The current Catenet routing algorithm assigns  a

length  of  1  to each Pathway, irrespective of the delay.  Other

possible assignments of lengths to Pathways  are  also  possible.

We  will  recommend  the use of measured delay as the best metric

for the internet routing algorithm to use, and we argue for  this

proposal  in  sections 4.3.1 and 4.3.3.  Section 4.3.2 covers the

related topic of "load splitting."  (One purpose of that  section

is  to  show  that the two topics are indeed related, and in ways

more subtle than generally realized.)  In section 4.4, we discuss

some of the issues in the design of an algorithm to  measure  the


     In  the  ARPANET,  a  routing  update  generated by an IMP A

specifies the average per-packet delay on each  of  A's  outgoing

lines.   Every  update generated by an IMP is sent to every other

IMP in the network, not just to the neighboring IMPs, as  in  the

Catenet  routing  algorithm.   This  updating  protocol,  and its

applicability to the internet, are discussed in section 4.5.

     Although a routing scheme can be divided into  a  number  of

separable  components,  it  is important to keep in mind that the

ultimate characteristics of the routing scheme will  result  from

the  combination  of  the  components.   A routing scheme must be

judged as a whole.  The reader should try to focus throughout  on

how  the  components  work together, and resist the temptation to

judge each component separately.

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4.3.1  Min-Hop Routing (Why Not to Use it)

     The simplest routing scheme which is based  on  having  each

Switch  compute  its  shortest  path  to  each  other  Switch  is

"min-hop" routing.  In min-hop routing, all Pathways are assigned

unit length, so that the shortest path between  two  Switches  is

just   that  path  which  has  fewer  Pathways  than  any  other.

(Generally, ties are broken arbitrarily.)  This sort  of  routing

is  used  in the current Catenet, where traffic is routed through

the  fewest  possible  number  of   intermediate   networks   (or

equivalently,   through   the   fewest   number  of  intermediate

gateways.)  This form of routing is quite simple,  and  does  not

require  us  to  worry about anything as complicated as detecting

changes in load or delay in  remote  components  of  the  Network

Structure.  Such changing conditions within the Network Structure

have  no  effect at all on the routing.  This form of routing can

be done with the minimal amount of overhead (in terms of the need

to send routing updates from Switch to Switch).  Updates need  to

be sent only when the Pathways go down or come up.  Any algorithm

which  attempts  to  be more responsive to changing conditions in

the Network Structure than  min-hop  routing  still  needs  these

up/down   updates,   plus   more  besides.   Min-hop  routing  is

definitely what one would use if one wanted to put  in  a  "quick

and   dirty"  routing  algorithm,  and  put  off  worrying  about

complexities until some  unspecified  later  time.   It  is  also

possible  to  argue  for  min-hop routing in the internet on more

principled grounds, as follows:
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     "In general, it is not unreasonable to expect that the  more
     component networks an internet packet goes through, the less
     likely  it  is to get to its destination, and the longer its
     delay is likely to be, if it does reach its destination.  We
     might expect that the number of component networks a message
     goes through would generally correlate fairly high with  the
     delay  of  the message, and would generally correlate fairly
     low with the obtainable throughput of a host-host transfer."

     Unfortunately, this sort of reasoning  is  only  valid  when

applied   to  a  Network  Structure   consisting  of  homogeneous

Pathways, which have  similar  characteristics  with  respect  to

delay,  throughput,  and reliability.  This is rather unlikely to

be the case in the internet, whose distinguishing  characteristic

is  the  heterogeneity  of its Pathways.  Where the Pathways of a

Network Structure have widely varying characteristics, delay  and

throughput  are not very likely to correlate well simply with the

number of hops.

     It is true that the delay-oriented routing  of  the  ARPANET

generally  gives  the min-hop paths.  (Remember, though, that the

ARPANET,  unlike  the   internet,   has   generally   homogeneous

Pathways.)   Min-hop  routing is all right for the "normal" case,

where there  are  no  areas  of  congestion  in  the  network  or

internet,  no areas where the delay is unusually high compared to

other areas.   Routing,  however,  is  no  different  from  other

computer  system  applications,  in that a scheme that works well

only in  the  normal  case  just  is  not  robust  enough  to  be

satisfactory.   (Think  of  a magnetic tape driver which works in

the normal case, where no tape errors are encountered, but  which

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crashes  the  system  in  the  presence of "unusual" events, like

errors on the tape.  Such a  driver  may  be  acceptable  if  one

accesses  one tape a month, but not if one needs to read or write

ten tapes a day.  The analogy is that min-hop routing may perform

acceptably in an experimental network with little traffic, but is

much less likely to be acceptable in a heavily loaded operational

network.)  It is extremely common for some area of the network to

be much more congested than another, so that traffic flows  which

traverse  a  particular  area experience a very much longer delay

(and lower throughput) than traffic flows which avoid that  area.

Significant  imbalances  in  load cause significant reductions in

the  correlation  between  hop-count   and   performance.    Such

imbalances  may not be present in a network initially, but if the

ARPANET experience is any indication, imbalances start  to  occur

with  increasing  frequency as network utilization grows.  If the

routing algorithm cannot account  for  such  imbalances,  network

performance  problems  will  start  to occur with ever-increasing

frequency  as  the  network  gains  more  users.   This  was  our

experience  with the original ARPANET routing algorithm.  For all

its widely publicized faults, it  provided  generally  acceptable

performance as long as the network was very lightly utilized, but

its  failures became more and more evident as the ARPANET shifted

from a research prototype to a  communications  utility.   If  we

expect  our  network or internet to be heavily used by real users

who are sending  real  data  that  they  really  need  for  their

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THESE AREAS, instead of blindly sending  traffic  into  congested

areas.   At a certain level of congestion, sending traffic into a

congested area is like sending it into a black hole; the  traffic

will  never  leave  the  area  to  progress  to  its destination.

Sending traffic into a congested area  also  induces  a  feedback

effect,   causing  the  congestion  to  spread  farther  than  it

otherwise would, and making it that much  less  likely  that  the

congestion  will  dissipate.   Any routing algorithm which cannot

take this into account will not be robust enough to survive in  a

real operational environment.

     Min-hop  routing also has another disadvantage which is more

specific to the internet environment.  Let  N1,  N2,  and  N3  be

three  networks,  and suppose we have to get some traffic from N1

to N3 by using N2 as a transit network.  Let  G12  be  a  gateway

connecting  N1 and N2, let G23 be a gateway connecting N2 and N3,

and let G2X  be  a  gateway  which  connects  N2  to  some  other

unspecified network.  If we use min-hop routing, then any traffic

which must go from G12 to G23 must go "directly", through network

N2, without stopping at G2X, because the path G12-G2X-G23 has one

more  hop  than the path G12-G23.  Perhaps this doesn't seem like

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much of a restriction; why would one want to have traffic stop at

the intermediate gateway G2X when it could go directly  from  G12

to G23?  Actually, two possible reasons come to mind immediately.

The  first  reason  has  to  do  with the possible effects of the

network's end-end protocol.   In  the  ARPANET,  for  example,  a

source  host  is  allowed  to  send  only  8  messages to a given

destination host before receiving the RFNM for the first of the 8

messages.   Hence  the  throughput  obtainable  on  a   host-host

connection is inversely related to the amount of time it takes to

get  a  RFNM  from  the  destination host to the source host.  It

follows that higher throughputs are obtainable between hosts that

are "near" each other than between hosts that are "far" from each

other.  It is also possible that G12 and G2X will be near to each

other, and that G2X and G23 will be near to each other, but  that

G12  and  G23  will  be  far  from each other.  So the throughput

obtainable in a transfer between G12 and G23  may  be  less  than

that  obtainable in a transfer between G12 and G2X, and less than

that obtainable in a transfer between G2X and  G23.   It  follows

that the throughput obtainable between G12 and G23 via G2X may be

higher  than  the  throughput  obtainable  between  G12  and  G23

directly.  Basically, by using an  additional  gateway  hop,  the

ninth  message  from  G12  can  be put into the network while the

first message is still in transit from G2X to G23, while  without

the  intermediate hop, this is not possible.  Of course, the best

solution to this sort of problem would  be  to  fix  the  end-end

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protocol  so  that  it  does not impose this sort of restriction.

Our present point, however, is that our routing algorithm  should

not rule out the possibility of this sort of strategy.  Note that

by  using an intermediate gateway hop, we might not only increase

throughput, but also decrease the delay (since  a  ninth  message

would not be blocked as long.)

     (It  is  interesting  to  think  about  whether this sort of

strategy might not be useful entirely within the ARPANET.)

     Another possible scenario in which an  intermediate  gateway

hop  might  be  useful  occurs  if  the  intermediate  gateway is

multi-homed.  It is possible that an intermediate gateway will be

homed to two IMPs which are distant from each  other  within  the

network.   If  so,  the  intermediate  gateway  may be used as an

"expressway" around a congested area of the network.

     If we replace the intermediate gateway G2X with two gateways

G24 and G42, we also have the possibility of sending traffic from

N1 through G12 into N2 to G24 through N4 to G42 into  N2  to  G23

and  thence into the destination network N3.  This is akin to the

oft-discussed expressway problem, but cannot  be  handled  within

the  framework  of  min-hop routing.  Of course, it might be very

difficult to take account of such factors, but one would not want

to have a routing scheme which makes it absolutely impossible.

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     Still another disadvantage of min-hop routing in the Catenet

is the following.  The current Catenet  routing  algorithm,  when

faced  with  three  gateways  on  the same network, considers the

three to be  equidistant.   However,  the  delay  and  throughput

obtainable from gateway A to gateway B may be very much different

than the throughput obtainable from gateway A to gateway C.  In a

large  distributed  network like the ARPANET, some pairs of hosts

are  connected   by   high-performance   paths,   and   some   by

low-performance  paths (either because they are separated by many

hops, or because the path between them  is  under-trunked,  etc.)

Allowing  the  routing  algorithm  to  be sensitive to this could

potentially have a large impact on the internet performance.

     There may not be any  network  that  actually  uses  min-hop

routing,  except  for  the Catenet.  There are, however, networks

that use a variant of  it,  which  we  might  call  "fixed  cost"

routing.  In fixed cost routing, each Pathway is still assigned a

constant  length,  but  not  all  Pathways  are assigned the same

length, and some Pathways have a length which is not equal to  1.

In a scheme like this, one attempts to assign values of length so

that  slow-speed  lines  appear  longer  than  high-speed  lines,

reliable lines appear shorter than  unreliable  ones,  and  lines

with  high  propagation  delays appear longer than lines with low

propagation delays.  This sort of routing is used in DATAPAC  and

in   DECNET.   Both  those  network  architectures  have  routing

algorithms based on the original ARPANET routing algorithm.   The

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designers of those architectures apparently realized that min-hop

routing  is  not  very  satisfactory  if  the  links  are  not of

relatively homogeneous quality, but were  probably  wary  of  the

problems that the ARPANET's original algorithm had in adapting to

changing  traffic conditions.  They avoided these problems by not

adapting at all to changing traffic conditions.  Of course,  this

is  the  weakness  in  fixed cost routing.  It may be better than

min-hop routing  in  a  lightly  loaded  Network  Structure  with

heterogeneous Pathways, but in a heavily loaded Network Structure

with unbalanced load it really is no better than min-hop routing,

and will still send traffic right into congested areas.

     We  have been emphasizing the claim that routing ought to be

able to detect congestion and route traffic around it.  Some  may

wonder  whether  we are confusing the proper functions of routing

with the proper functions of congestion control.  That is not the

case.  Congestion control schemes  generally  try  to  limit  the

amount  of  traffic  entering  a  network  so as to prevent or to

reduce the overloading of some resource or of the whole  network.

When  congestion  actually  exists in the network, however, it is

the job of routing to try to send traffic  around  the  congested

areas;  otherwise  the  routing actually causes the congestion to

increase.  Of course, one might attempt  to  design  the  routing

algorithm  under  the  assumption that there will be a congestion

control scheme that will make  congestion  impossible.   However,

such  a  design  could not be very robust.  If we want to build a

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robust Network Structure which will continue to operate  under  a

variety  of  unforeseen  conditions,  then  we want each software

module or protocol to be designed with the  assumption  that  the

other  modules  or  protocols  will  be  less  than optimal.  The

resulting system will be much less prone to  system-wide  failure

than one which is designed so that no part of it will work at all

unless every part of it works perfectly.  Although we will not be

discussing explicitly, in this paper, any schemes for controlling

the  amount  of  traffic  which  is  input  to the internet, that

doesn't mean that we can ignore the  way  in  which  the  routing

algorithm affects and is affected by the existence of congestion.

Particular  problems  related  to  overload  of network resources

should be discussed in whatever context they  arise  in,  without

worrying about whether the problem is properly called "congestion

control"  or "routing."  There is in general no way of telling in

advance whether the best solution to a particular  problem  is  a

routing  solution  or  a congestion control solution, and putting

labels on the problems just restricts our thinking.

4.3.2  Load Splitting

     Routing  in  the  ARPANET  has  always   been   "single-path

routing."   We  mean  by  this  that  at  any  given  moment, the

ARPANET's routing algorithm provides only a single  path  between

each  pair of IMPs.  All traffic which enters the network at some

particular time, originating at IMP A and  destined  for  IMP  B,

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will  travel  over  the  same  path.  Actually, this statement is

somewhat oversimplified, since there might be a change of  routes

while some traffic is already in transit.  The point, however, is

that at any given time, each source or intermediate IMP will send

all  traffic  for  a  particular  destination  IMP  to  a  unique

neighbor; it cannot split the traffic among several neighbors.

     Routing in the  Catenet  is  currently  somewhat  different.

Suppose  gateway  A  has  two  neighbors,  B  and C, and has some

traffic to send to gateway E.  The routing  algorithm  run  in  A

assigns  a distance value to the path to gateway E via neighbor B

and a distance value to the path to E via  neighbor  C.   If  the

distance  from A to E via B is the same as the distance from A to

E via C, then gateway A will alternate between use  of  B  and  C

when  sending traffic to E.  That is, A makes simultaneous use of

two distinct paths to E.  Such a scheme would  be  somewhat  more

difficult  to  put  into  SPF routing, because in SPF routing, no

assignment of distance values from A to E via  each  of  the  two

neighbors  is  generated.  Rather, only one path is computed, via

one of the neighbors, and only the distance on that one  path  is

known.   Distance  on  other  paths  is  not  computed by the SPF

algorithm.  (On the other hand, the SPF algorithm  generates  the

entire  path,  so that each Switch knows which other Switches its

traffic will be routed through on the  way  to  the  destination.

The  original  ARPANET algorithm does not do this, but only tells

each Switch which of its neighbors to use when sending traffic to

the destination.)
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     What is the significance of this?  It seems to  be  commonly

regarded  as  obvious that multi-path routing, or load splitting,

is an important advantage, so that routing algorithms that permit

it are better than routing algorithms that do not.  However, when

one asks advocates of multi-path routing why it  is  better  than

single-path   routing,   a   very  common  answer  seems  to  be,

"Multi-path  routing  is  better  because  it  provides  multiple

paths."    This   sort   of   answer   is   rather   superficial.

Multiple-path routing is NOT a goal  in  and  of  itself;  IT  IS


If a multi-path routing algorithm results in  smaller  delays  or

larger  throughput than some other algorithm, then that is a good

reason for favoring it over the  other  algorithm.   Now,  it  is

certainly true that any routing algorithm which OPTIMIZES network

delay  or  throughput  will  be  a  multiple-path algorithm.  THE

CONVERSE, HOWEVER,  IS  NOT  TRUE.   A  routing  algorithm  which

provides  multiple  paths  does not necessarily optimize delay or

throughput.  In fact, merely because a routing algorithm provides

multiple paths, it  does  not  follow  that  it  provides  better

performance  in  any  respect  than  some other routing algorithm

which provides only a single path between a pair of Switches.  An

algorithm which provides a single good path may be  far  superior

to an algorithm which provides several poor ones.

     To see this, let's look at some possible effects of the load

splitting  in the Catenet routing algorithm.  Let A, B, C, D, and

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E be five gateways, and suppose that there are two possible paths

from A to E, namely ABDE and ACDE.  The Catenet routing algorithm

would regard these two paths as equidistant, since that algorithm

regards two paths as equidistant if they contain the same  number

of intermediate gateways.  Therefore gateway A would perform load

splitting  on  its  traffic  to E, sending half of the traffic to

neighbor B and half  to  neighbor  C.   Does  this  provide  more

throughput  than  the  use  of  a single one of these paths?  Not

necessarily.  If the bottleneck on the paths from A to E  is  the

Pathway  DE,  then  the  use  of these two paths provides no more

throughput than the use of either one alone.  In fact, if  DE  is

the  bottleneck, the use of the two paths will probably result in

lower throughput than the use of  a  single  path.   The  use  of

several paths increases the likelihood of the packets from A to E

arriving  out  of  order  at  the  destination host.  Yet as more

packets arrive out of order, more TCP  resources  are  needed  to

handle them, and the TCP just has that much more work to do.  TCP

buffers  that  are  occupied  by  out-of-order  packets cannot be

"allocated" for receiving more packets, so  acknowledgments  must

be  delayed, and windows must be kept smaller.  The result of all

this will be higher  delays  and  lower  throughputs.   This  was

probably  not  the  intention  of load splitting, but is a likely

consequence of it.

     Suppose there really are two independent paths from A  to  E

which  are  "equidistant", say ABDE and ACFE.  Even here, sending

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half the packets on each path may only degrade  performance.   To

see this, suppose each of the Pathways AB, BD, DE, AC, and CF has

a  capacity  of  50  kbps,  but that link FE has a capacity of 10

kbps.  Suppose also that we want to send 50 kbps of traffic  from

A  to  E.   If  we  alternate packets between these two paths, by

trying to send 25 kbps of traffic each way, we will  be  able  to

get at most 35 kbps of traffic through to the destination, and we

will  cause  severe  congestion  on  link FE (which will probably

result in its being able to carry even less  than  the  rated  10

kbps, further lowering the network throughput.)  Had we used only

the  single  path  ABCD,  we  would  have  been able to pass more

traffic.  Again, we  see  a  situation  where  the  use  of  load

splitting can reduce throughput and increase delay.

     This  sort  of  problem  might  at  first  appear  to be too

unlikely to be worth worrying about.   However,  it  has  already

occurred  in  the  Catenet, and has caused a significant problem.

In fact, in the Catenet's actual problem, half of the traffic was

sent on a path whose capacity was sufficient to  handle  all  the

traffic,  and  the  other  half of the traffic was sent on a path

whose capacity was essentially zero (because a network  partition

made  the  destination  host  unreachable on that path).  In this

case, load splitting resulted in  the  throughput  being  cut  in

half,  as  half  the  traffic  was routed down a black hole!  The

problem was "solved" by  eliminating  one  of  the  two  possible

paths,  thereby  eliminating  the  possibility of load splitting.

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However, this does not seem like a proper way to deal  with  this

problem in the general case.

     The  Catenet's  load  splitting  has been defended from this

latter objection as follows:  "If there were no  load  splitting,

maybe  all  the traffic would have been sent into the black hole,

not just half."  This is less a defense than a sad commentary  on

the  state of the Catenet routing; to accept this sort of defense

is just to give up entirely on the problem of  internet  routing.

     Someone  may  reply to our first criticism of load splitting

by saying "maybe the bottlenecks will  be  Pathways  AB  and  AC,

rather  than DE, in which case the use of two paths does increase

the throughput."  This reply is correct, but not very  important.

The  sort  of  load  splitting done in the Catenet might, by pure

chance, increase throughput in some particular case.   The  point

though  is  that it is no more likely to increase throughput than

to decrease it.  Certainly there is no reason to suppose that the

cases in which it might help are any more likely  to  occur  than

the cases in which it hurts.  In our experience with the ARPANET,

schemes  that  seem  a priori as likely to hurt as to help always

end up hurting more than helping.  (In networking,  Murphy's  law

is  more  than just a joke.)  Choosing equidistant paths for load

splitting will generally result in paths  which  are  only  small

variants  of each other (if it results in any paths at all, since

there are not necessarily several  equidistant  paths  between  a

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pair  of  Switches),  and  there is no reason to suppose that the

bottleneck will not be common to each path.  Even if  we  do  get

two  paths  which  do  not  share  a  bottleneck,  unless  we try

explicitly to apportion the flows to the relative  capacities  of

the  two  paths (rather than just dividing the traffic 50-50), we

will not, in general, gain any increase in throughput.

     In chapter 4 of [6], we  actually  devised  a  multiple-path

routing  scheme,  based  on  SPF,  whose  purpose was to maximize

throughput.  In this  scheme,  we  make  sure  that  any  set  of

simultaneously    used    paths    between   two   Switches   are

"bottleneck-disjoint", (i.e., they don't share a bottleneck),  so

that  we  know  that we can get more throughput by use of several

paths.   We  also  devised  a  flow  apportionment  scheme  which

attempts  to  match  flows  (or  parts of flows) to the available

capacity of each path.   Anyone  interested  in  seeing  what  it

really  takes  to  do  multi-path  routing  should  look  at that

chapter.  The scheme proposed there is  quite  complex,  however,

and  it  is  not obvious that it will work.  Some simulation work

will eventually be done on it.  Until that sort of  algorithm  is

much  better  understood,  it  would  not be very wise to use the

internet to experiment with it.  It will be difficult  enough  to

adapt  a  well-understood  and  much-used routing algorithm (like

that currently in the ARPANET) to the internet environment.   The

internet  is certainly not a place for experimenting with new and

untried routing algorithms.

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     Although it  is  quite  difficult  to  design  a  multi-path

routing  procedure  that  results  in significant improvements of

delay or throughput over single-path  routing,  there  are  other

reasons  for  requiring multiple paths between a pair of Switches

that are more easily dealt with.  For example, we may be required

to have different paths  between  a  pair  of  internet  gateways

because of ACCESS CONTROL RESTRICTIONS.  That is, certain classes

of  packets  may  not  be  allowed to traverse certain classes of

networks, so that different routes  would  be  required  for  the

different  classes of traffic.  We may also decide that different

types of service that may be requested by the user should  travel

over different paths, even if the source and destination gateways

are  the  same  for the different traffic classes (e.g., maybe we

don't want to use multi-hop satellite  networks  for  interactive

traffic.)   This  is  easily  handled within the framework of SPF

routing.  Remember that the SPF algorithm produces  the  shortest

path  to  a destination, based on an assignment of lengths to the

Pathways.  Rather than simply assigning a unique length  to  each

Pathway,  we  can  assign  a  set  of lengths, indexed by traffic

classes.  We can then produce a set of routing tables, indexed by

traffic type, such that the routing table  for  a  given  traffic

type   contains   the   "shortest"  path,  based  on  the  length

assignments for that traffic type.  For example, if traffic class

C is not permitted to  traverse  Pathway  P,  the  length  of  P,

indexed  by  C,  can  be set to infinity.  This ensures that that

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Pathway will not be part of any path found in the routing  tables

indexed  by  C.   We  even  have the flexibility to assign to P a

length which, while not infinite, is much larger than the  length

of  any  other  Pathway.  In this case, that Pathway will be used

for traffic of class C only if  EVERY  path  to  the  destination

includes  it  (i.e.,  only if it can't be avoided).  This sort of

load splitting might be quite important in the internet,  and  is

also quite simple to handle.

4.3.3  Delay vs. Throughput

     In  the  ARPANET,  each  IMP  measures the average delay per

packet on each of its outgoing  lines.   This  average  delay  is

assigned  as  the  "length"  of  the line, and shortest paths are

computed on that basis.  We have studied the performance of  this

algorithm a great deal [5].  It tends to use min-hop routes under

conditions of light or of uniform load.  However, it does seem to

take  account  quite well of the varying delays that are produced

by  lines  of  different  transmission   or   propagation   delay

characteristics.   Since congestion causes large increases in the

delays,  congestion  is  generally  detected   by   the   routing

algorithm,  and  traffic  really is routed around congested areas

when that is possible.  While we cannot claim  that  our  routing

algorithm  gives  the  optimal delay, the characteristics that it

does have seem to be the characteristics  that  we  would  really

like  to see in any robust, operational network, and particularly

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in the internet.  The routing tends to  be  stable  on  what  are

intuitively  the  best  paths, except when exceptional conditions

arise which make it clear that  some  other  path  is  likely  to

provide  better performance.  It is this sort of routing which we

propose for the internet.

     Before discussing further the use of delay-oriented  routing

in  the  internet, we would like to briefly consider the issue of

throughput-oriented routing.  In the previous section, we  argued

against  the  use  of multi-path routing as a means of optimizing

throughput, largely  on  the  grounds  that  doing  it  right  is

extremely difficult (much more so than one might at first think),

that  the ways of doing it right are quite poorly understood, and

that the internet is not  a  good  testing  ground  for  new  and

untried algorithms.  However, one often hears that there are high

throughput  applications  (bulk  traffic) for which delay doesn't

matter, and one may wonder whether there  is  not  some  kind  of

single-path   routing   which   is   more  appropriate  for  such

applications than is delay-oriented routing.  One scheme that  is

very commonly suggested is that of routing traffic on the path of

maximum  excess  capacity, instead of on the path of least delay.

     Given an algorithm for  determining  the  amount  of  excess

capacity  on  each  Pathway  (which  could  be quite difficult to

design for the internet environment -- how do we  know  what  the

excess  capacity  of  a  packet-switching  network is?), it is no

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difficult matter to modify the SPF algorithm to produce the paths

of maximum excess capacity.  However, it would not be a good idea

to use the resultant routes for bulk traffic.  For one thing,  we

must  understand that such a routing algorithm would not maximize

total  network  throughput.   (By   "maximizing   total   network

throughput",  we  mean  maximizing the amount of traffic that the

network can handle.)  Suppose, for example, we wanted to send  40

kbps  of traffic, and had the choice of using a one-hop path with

excess capacity of 50 kbps, or a 10-hop path, each of whose links

had an excess capacity of 100 kbps (so that the  total  composite

path  has  an excess capacity of 100 kbps).  By using the shorter

path, we use up a total of 40 kbps of network capacity,  capacity

which  is now unavailable for other traffic.  By using the longer

path (which is the path of maximum excess capacity), we use up  a

total  of  10x40 kbps (40 kbps per hop), thereby using up a total

of 400 kbps which is no longer available for other  traffic.   In

terms of maximizing the total network throughput, we do better by

using  the  one-hop  path, rather than the path of maximum excess


     Maybe we are less interested  in  maximizing  total  network

throughput  than  in  finding  a path for some particular traffic

flow which has enough capacity to handle the required  throughput

of that flow.  We still would not want to use the path of maximum

excess  capacity,  for that path might have a delay which is much

too long.  Although we often hear that certain classes of traffic

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(e.g., file transfer) care only about throughput, not delay, this

is really a gross oversimplification.  In file transfer, we don't

care how long  it  takes  for  the  first  packet  to  reach  its



PACKETS.  Of course, if there are long delays between the packets

of a file transfer, the throughput will be very low.  Hence it is

not  quite  true  to  say  that  file  transfers and the like are

unconcerned with delay.  If higher level protocols like  TCP  are

being used, then routing over a path of long delay will certainly

result  in  lower  throughput.   The reason is as follows.  A TCP

sender will only send a certain amount of data,  until  he  fills

the window specified by the TCP receiver.  The size of the window

is  very  likely  to depend on such network-independent things as

the amount of resources (e.g., buffers) in the destination  host.

If  the  path  between  source and destination host is very long,

then the sending TCP will fill the window, and then have to wait,

idly, for some  period  of  time  while  his  data  gets  to  the

destination,  and  while the message indicating the re-opening of

the window is transmitted from the  receiving  TCP.   Since  this

network-imposed  long  delay causes the sending TCP to have to be

idle for some period of time, it holds down the  throughput.   So

it   seems   that   all   things   considered,   simply   routing

high-throughput application traffic on the path of maximum excess

capacity is unlikely to actually result in high throughput.

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     If we really wanted to  do  single-path  throughput-oriented

routing,  we  would  need something like the following.  We would

want to route traffic on the shortest path  (i.e.,  the  path  of

least   delay)  which  does  not  contain  any  components  whose

available capacity is too small to handle the needed  throughput.

This  would prevent us from choosing a path with arbitrarily long

delays, or a path with too little capacity.  Unfortunately, it is

almost impossible to find out either what throughput is needed by

an application,  or  to  find  out  just  what  the  capacity  of

particular  components  of  the  internet  is.   We might want to

consider some strategy such as not sending batch traffic on paths

which include components which are very heavily loaded.  This  is

fertile  ground for experimentation.  Our present point, however,

is that the delay-oriented SPF routing  of  the  ARPANET  already

provides  the  basic  structure  that we need to accommodate this

sort of strategy.  If we knew that  we  wanted  bulk  traffic  to

avoid   certain   Pathways   (e.g.,   Pathways  with  too  little

bandwidth), we could have SPF routing compute the shortest routes

that did not  include  those  Pathways,  by  using  the  "indexed

length"  scheme  described in section 4.3.2.  There is no need to

consider different sorts of routing schemes.

4.3.4  Knowing the "Whole Picture"

     The use of the SPF algorithm requires that every Switch know

the complete topology of the Network Structure.  That  is,  every

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Switch  must  know  of  all  the  other Switches, must know which

Switches are "directly connected" to which  other  Switches,  and

must  know the "length" of each Pathway.  This is not to say that

this information is "compiled in", or even  loaded  in  manually.

Rather,  it  is  determined  dynamically,  in  real-time, through

interpretation of the routing updates (see section 4.5).   It  is

this  uniform  global  knowledge  of the topology and the Pathway

lengths  that  enables  each  Switch  to  run  a  shortest   path

algorithm,  while  producing routes which are consistent with the

routes produced by other Switches, so that routing loops  do  not

form.   The  SPF algorithm does not merely tell a Switch to which

of its neighbors  it  should  send  packets  for  destination  D.

Rather,  it  computes  the entire path to the destination Switch.

However, when a packet is routed, it does not carry with  it  the

identity  of  the entire route, as computed by its source Switch.

Each Switch just forwards the packet to the next "hop" along  its

route.   The  fact  that  all  Switches have the same information

about the topology is what ensures that this routing will be free

of loops.

     Since each Switch performs its routing based on  a  complete

picture  of  the  topology  of the Network Structure, we can call

this sort of routing scheme a "whole picture"  scheme.   In  this

section,  we will compare "whole picture" schemes with some other

schemes which do not require the Switches to have uniform  global

knowledge of the topology.  We argue that "whole picture" schemes

are always superior.
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     The original ARPANET routing scheme, and the current Catenet

routing  scheme,  are  not  "whole  picture"  schemes.   In these

routing schemes,  no  Switch  need  have  any  knowledge  of  the

topology, other than who its own immediate neighbors are, and the

lengths  of  the  Pathways  to  its  immediate  neighbors.  These

algorithms function as follows.  When a Switch first comes up, it

forms a hypothesis as to the best neighbor to which to send  data

for each possible destination Switch.  This initial hypothesis is

based  only on its own local information about the lengths of the

Pathways  to  its  neighbors.   It  then  informs  its  immediate

neighbors of its hypotheses, and is informed of their hypotheses.

It   then  forms  a  new  hypothesis,  based  on  its  own  local

information  AND  the  hypotheses  communicated  to  it  by   its

neighbors.   It  then  exchanges  hypotheses  with  its neighbors

again, and again, and again, until  its  own  hypotheses  are  in

complete  agreement  with  those of its neighbors, at which point

stability is reached.

     To see the difference between this sort  of  routing  scheme

and the "whole picture" scheme, consider the following situation.

Suppose we have 100 people in a room, sitting in chairs which are

properly lined up so that we can talk of each person's having two

immediate  neighbors.   We  also have a picture of an object, and

our goal is to have ALL the people agree on the identity  of  the

depicted   object.   Now  we  have  a  choice  of  two  different

procedures for bringing this about:

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Procedure 1: Cut the diagram into 100 pieces, and give one  piece
             to  each person.  Each person is now allowed to look
             at his one piece, and then form a hypothesis  as  to
             what  is  depicted  in  the full picture.  Then each
             person  can  exchange  hypotheses  only   with   his
             immediate  neighbors.   Then  each person can form a
             new hypothesis and exchange that with his  immediate
             neighbors.   The  procedure  terminates when all 100
             people agree on what is depicted.

Procedure 2: Make 100 Xerox copies of the diagram, and distribute
             the copies to each person.

If we really think it is important for each person to  know  what

is  depicted  in  the  picture,  then  we  will  certainly follow

procedure 2,  which  will  make  the  whole  picture  immediately

available  to all participants.  Procedure 1 would only be useful

as a party game.  It would  be  quite  amusing  to  see  all  the

ridiculous  hypotheses  that  are  formed before all participants

converge to the correct one, IF they ever do manage to  converge.

Even  if  they  do converge, it might take quite a long time.  We

must remember that different people form hypotheses at  different

rates,  and can communicate them at different rates.  Some people

may simply refuse to talk to certain neighbors at all.  If  one's

left-hand  neighbor  has  formed  a  good  hypothesis,  but one's

right-hand neighbor has not, one's own hypothesis is likely to be

thrown off the track, which in turn is likely  to  mislead  one's

left-hand  neighbor  into  a  poorer  hypothesis  during the next

"iteration."  This is not a very optimal procedure  for  bringing

about convergence of opinion.

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     However,   this   situation   is   really   too  simple  and

straightforward to be truly analogous to routing.  To improve the

analogy, we must suppose that the picture is constantly changing,

even as the people are still forming hypotheses.  In procedure 2,

this change  is  accounted  for  by  simultaneously  giving  each

participant  a  new copy of the picture.  In procedure 1, changes

in the picture are accounted for as follows:  if the part of  the

picture  originally  given to person P has changed, then give him

the corresponding piece of the same picture; he can now use  this

piece  when  forming  his hypotheses, and should forget about the

previous piece.  When the procedures are thus  modified  to  take

account  of  changes  in  the picture, the situation described is

more analogous to routing, and the advantages of procedure 2 over

procedure 1 are even more pronounced.

     The  ARPANET's  current  routing  algorithm  is  similar  to

procedure  2,  since  the whole picture is made available to each

Switch.   The  ARPANET's  original  routing  algorithm,  and  the

Catenet's  current  one, are more similar to procedure 1; perhaps

they should be called "jigsaw puzzle"  algorithms.   All  of  the

problems  of  procedure  1  have their analogies in those routing

algorithms.   It   should   be   obvious   that   in   terms   of

responsiveness,   accuracy,   and   consistency,   whole  picture

algorithms are superior to jigsaw puzzle algorithms.  Many of the

problems of the  original  ARPANET  routing  algorithm,  such  as

looping  and  very  slow  response  to topological change, can be

attributed to its "jigsaw puzzle" nature.
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     Even if one agrees that we ought to  avoid  "jigsaw  puzzle"

algorithms,  one might still claim that we need not have a "whole

picture" algorithm.  One might wish to argue that a given  Switch

needs  to know only the topology of a "region" which contains it.

This region would be larger than a  single  Switch,  but  smaller

than   the   set  of  all  Switches.   A  region  would  also  be

geographically contiguous, so that if two  Switches  are  in  the

same  region, then there is a path between them which is entirely

within the region.  Then traffic which does not need to  leave  a

region  to  get  from  its source to its destination is in effect

routed by a "whole picture" scheme.  Traffic which must leave the

region, however,  does  not  have  its  whole  route  preplanned.

Switches  within one region will know only how to get traffic out

of the region.  Other Switches in the next region will  know  how

to get the traffic through that region, etc.  It seems, one might

argue,  that this sort of regionalized routing scheme ought to be

possible.  After all, consider the  analogy  with  ordinary  road

travel.   If  one wants to travel from Boston to Los Angeles, one

need not preplan the entire route.  One  can  just  head  in  the

general  direction  of Los Angeles, with no need to know anything

about the roads which are close to Los Angeles until one actually

gets close.  A similar scheme ought to work with data.

     One problem, however, with the suggested analogy, is that it

does not even hold in the case of ordinary automobile travel.  If

one were planning an automobile trip to LA,  one  would  want  to

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know  about  any  record-setting  blizzards  in  the midwest long

before one actually approached the midwest.  One  would  want  to

know  about the status of Mt. St. Helen's volcano long before one

approaches Oregon.  One might  try  not  to  be  passing  through

Chicago  at  rush hour.  Avoiding any of these potential disaster

areas could require quite a bit of advance planning.  Of  course,

the  amount of advance planning that one performs when travelling

is a matter of personality;  some  people  are  more  adventurous

than others, and might actually enjoy a disaster or two along the

way.   Users  of a data communications utility, however, whatever

personality traits they may have, generally  do  not  want  their

data to be sent on an adventure.  Rather, they want their data to

be   treated  with  a  conservatism  and  caution  which  require

considerable preplanning.

     In any case, the analogy between the road system and a  data

communications  network  is  very  misleading because of the very

rich interconnectivity of the road system.  No  matter  how  many

problems  an  automobile  driver  encounters as he approaches Los

Angeles, he still has a large number of choice points, in that he

can take any number of relatively short  detours  around  problem

areas.   In data networks, however, the connectivity is much less

rich, and the closer the data gets to its destination, the  fewer

choice   points   there   are.    With   a   sufficiently  sparse

connectivity, the entire path could even  be  determined  by  the

very first routing choice that is made, so that no detours around

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problem areas are possible once the "trip" begins.  The situation

is as if someone drove from Boston to Nevada, then found that all

roads from Nevada to California were closed, and that he then had

to  drive  all  the way back to Boston to start on a new route to

California.  This sort  of  sub-optimality  is  inherent  to  any

regionalized routing scheme for data communications networks.

     In  fact, the situation could be even worse.  If Switches in

Boston know nothing about what is happening  between  Nevada  and

California,  then data for California which arrives at Nevada and

then is sent back from Nevada to  Boston  for  alternate  routing

will  just  loop  back  to  Nevada.  The data will be stuck in an

infinite loop, never reaching its destination.  In IEN 179, Danny

Cohen proposes a regional routing scheme  like  this,  apparently

not  realizing  that  it  suffers  from loops.  His proposal also

includes a form of hierarchical addressing which is closely bound

up with routing, so that a Switch in Boston  might  not  even  be

able  to  distinguish  data  for Nevada from data for California.

That is,  in  Cohen's  scheme,  data  for  Nevada  and  data  for

California  would  be  indistinguishable  at the Boston Switches;

all such data would appear to be addressed to Nevada.   Only  the

Switches  at Nevada would look further down the address hierarchy

to  determine  whether  the  data  needs  further  forwarding  to

California.   Any such scheme is hopelessly loop-prone, except in

a Network Structure whose connectivity is  extraordinarily  rich,

much more so than the Catenet's will ever be.

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     It might seem like these objections would also have to apply

to the internet, since a gateway does not know all about IMPs and

packet  radios  and  SIMPs,  etc.,  in  the  component  networks.

However, the looping problem is avoided in the internet since  it

is  organized  in  a  strict  hierarchy  of  Network  Structures.

Switches in one Network Structure need not  know  anything  about

Switches  in  any  other  Network  Structure,  but they must have

complete information (Whole Picture) about Switches in  the  same

Network  Structure.   All  (source or intermediate) Switches in a

particular Network Structure always route data  to  a  Switch  in

that same Network Structure.  This imposition of strict hierarchy

prevents  looping,  as  long as the lower levels of hierarchy are

controlled by the higher levels.  In  the  internet,  this  means

that,  e.g.,  if  a  gateway hands a packet to an ARPANET IMP for

delivery to an ARPANET Host or to another internet  gateway,  the

ARPANET  is  required  to  deliver the packet as specified by the

gateway, or to say why not.  It must not simply pass  the  packet

back  to the gateway, or a loop will form.  (This sort of looping

has been frequently noticed between  IMPs  and  port  expanders.)

This  does  not imply that an ARPANET IMP cannot pass a packet to

an  internet  gateway  for  delivery  (through   an   "expressway

network")  to another ARPANET IMP, but only that once an internet

gateway decides to send a packet into the  ARPANET,  the  ARPANET

must  get that packet to the intended destination, or else inform

the gateway that it cannot do so.

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     It is also important to note that the hierarchical levels in

the internet tend to be  "horizontal",  rather  than  "vertical".

That  is,  in  an internet spanning North America, there would be

internet gateways located all across the continent,  as  well  as

IMPs   and   packet  radios  and  PSATs  located  throughout  the

continent.  This is  quite  different  from  regionalization,  in

which  Switches  which  are  close geographically are in a common

region.  This distinction is very important if we  are  to  avoid

such problems as looping.

     Although  building  the  internet  as  a strict hierarchy of

Network Structures avoids  the  problems  of  looping,  there  is

always  some  degree  of  sub-optimality  introduced whenever the

topological knowledge of the Switches is restricted in  any  way,

even  if  the  restriction  is  just  to Switches within the same

Network Structure.  This is a point to which we return in section

4.6,  where  we  discuss  some  of  the  basic   limitations   of


4.4  Measuring Pathway Delay

     One  of  the  most basic problems in devising a scheme to do

delay-oriented routing is to figure out a way  to  determine  the

delay.   In the ARPANET, the delay measurement algorithm is quite

straightforward.  When a packet arrives at an IMP, it is  stamped

with  its  arrival time.  When it is transmitted to the next IMP,

it is stamped with the time of transmission.  ARPANET packets are

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buffered in an IMP until acknowledged  by  the  next  IMP;  if  a

packet  has  to  be retransmitted, its transmission time stamp is

overwritten with the  time  of  latest  transmission.   When  the

packet  is acknowledged by the receiving IMP, the arrival time is

subtracted from the transmission time, yielding  the  total  time

the  packet  spent  in the IMP.  The propagation delay (i.e., the

speed of light delay along the phone line from  one  IMP  to  the

next)  is  then  added  in to compute the total amount of time it

took to get the packet from one IMP to the next.  There are three

important aspects of this delay measurement algorithm:

     1) It is necessary to measure the amount  time  each  packet

        spends  within  the  Switch.   This  should be as easy to

        apply to a gateway as to an IMP.

     2) It is necessary to determine how long it takes  a  packet

        to  travel  from  one  Switch to another over the Pathway

        connecting them.  If the Pathway is a telephone line,  as

        in  the  ARPANET, this is just the propagation delay, and

        is a constant which can be separately measured  and  then

        stored  in a table.  On the other hand, if the Pathway is

        a packet-switching network, or even an internet, this  is

        much  more  difficult  to  determine, and is certainly no


     3) There must be some way to account for packets that  don't

        get through, or don't get through immediately, due either

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        to  errors or to congestion.  In the ARPANET, if a packet

        doesn't get through on its  first  IMP-IMP  transmission,

        and  has to be retransmitted 200 ms.  later, this 200 ms.

        gets added into the  packet's  delay.   This  is  a  very

        important feature, since it enables the delay measurement

        to  reflect  the effect of congestion or of a very flakey

        line.   But  unless   the   gateways   run   a   reliable

        transmission   protocol  among  themselves,  it  will  be

        difficult to make sure that our delay measurement  really

        reflects  these  factors.   If we are trying to send data

        through a network which is dropping most of the  data  we

        send  it, we want to make sure that our delay measurement

        routines produce a high value of delay, so  that  traffic

        will  tend  to  be  routed  around  this  very flakey and

        unreliable Pathway.  (Remember that if too  much  traffic

        is  dropped, some (higher) level of protocol will have to

        do a lot  of  retransmissions,  resulting  in  very  high

        delays and low throughputs.)

     The problem of how to measure delay is more tractable in the

case  of  AREA  ROUTING  than  in the more general internet case.

Recall that by "area routing," we mean a sort of internet all  of

whose  component  networks are basically identical (see IEN 184).

For example, we might at some future time decide  to  divide  the

ARPANET  into  areas,  connected by gateways, so that the ARPANET

itself turns into a hierarchical network.  If we  decide  to  use

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the  same  routing  algorithm  at the high level (i.e., among the

intra-ARPANET gateways) as we use at the lower level (i.e., among

the individual IMPs in a  particular  area),  then  the  gateways

could  obtain the delay measurement information directly from the

routing updates sent by the individual IMPs.  That is, the  lower

level routing algorithm could provide information to the gateways

enabling  them  to  deduce their delay to other gateways.  If the

gateways  are   also   ordinary   IMPs,   this   information   is

automatically  available.   If  the gateways are hosts on the low

level ARPANET, a special protocol would have to be  developed  to

enable  the  IMPs to transmit the routing updates to the gateways

they are connected to (though this  wouldn't  be  much  different

from  the  protocol that IMPs now use to transmit routing updates

to their neighboring IMPs).  Of course, if we were to implement a

scheme like this, we would still want to make the ARPANET  appear

as  a single Pathway (with no intermediate Switches) at the level

of the Network Structure of the Catenet.  That  is,  the  Catenet

would  be  a  third  hierarchical layer over the two hierarchical

levels of the ARPANET, which would be transparent to it.

     In the more general internet case, we  cannot  rely  on  the

component   networks  to  provide  us  with  the  sort  of  delay

information we  would  like  to  use  for  the  internet  routing

algorithm;  the  internet  Switches will have to have some way of

gathering this information themselves.  In general, it  will  not

be possible for a Switch to measure the one-way delay from itself

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to  its neighbors.  (We wouldn't want to rely on the radio clocks

that are now beginning to be deployed  at  the  gateways;   while

these  might  be  useful for doing measurements, we wouldn't want

the reliability of the  entire  operational  internet  system  to

depend  on  a radio broadcast over which we have no control.)  It

is possible, however, to measure round-trip  delay  between  each

pair  of  neighboring  gateways.   In  the  ARPANET, for example,

round-trip time is easily measured by keeping  track  of  when  a

message  is  sent  to  a neighboring gateway, and then noting the

time  when  the  RFNM  is  received.   One-way  delay  would   be

approximated by dividing the round-trip delay in half.

     It  is  certainly  true that the round-trip delay is not, in

general, exactly twice the one-way delay.  However, it seems like

a good enough  approximation  to  use  in  the  internet  routing

algorithm.    All  we  really  require  is  that  it  be  roughly

proportional to the one-way  delay,  in  that  both  one-way  and

round-trip  delays  tend  to  rise  and  fall  together, and that

congestion in the Pathway (component network) tends to make  both

increase.    Of   course,  before  designing  the  precise  delay

measurement scheme that we would want to use in the internet,  we

would  have to run a series of tests and experiments to see which

of several possible delay measurement  algorithms  gives  us  the

results  we want.  This would be similar to the extensive testing

of the ARPANET's delay measurement algorithm that  is  documented

in [4].

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     Unfortunately,  there  are many networks which do not return

anything like  RFNMs  that  could  be  used  to  gauge  even  the

round-trip   delay.    (Many   networks,  e.g.,  SATNET  and  the

forthcoming wideband network, do not even tell  you  whether  you

are  sending traffic to a host which is down.)  So we will need a

gateway-gateway protocol in which gateways  receiving  data  from

other  (neighboring) gateways send back replies which can be used

for timing.

     This does not mean that every packet sent from  one  gateway

to  another  must  be  acknowledged  by  the  receiving  gateway.

Rather, we would propose something like the  following.   Suppose

we  have,  as  part of the gateway-gateway protocol, a bit that a

sending gateway can set which requires the receiving  gateway  to

acknowledge  the  packet.   The sending gateway can have a random

number generator, which lets it select packets at random in which

to set this bit.  These packets will have their round-trip  delay

measured,   and   will  constitute  a  random  (and  hopefully  a

representative) sample.  The packets need not be buffered in  the

sending  gateway  pending  acknowledgment,  but they will need to

have unique identifiers so  they  can  be  kept  track  of.   The

round-trip  delay  of  each packet is then easily determined when

the acknowledge is received.  (This probably implies though  that

gateways  will  have  to run a protocol with their neighbors when

they first come up in order to synchronize  sequence  numbers  to

use  for  identifying packets uniquely.)  There will also have to

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be a time-out, so that a packet which is not acknowledged  within

a certain amount of time (perhaps dependent on the expected delay

of the packet, based on previous measurements) will be considered

to  have  been  lost  on  the Pathway between gateways (or in the

receiving gateway).  Packets  which  have  been  lost  should  be

assigned a very high delay, so that the routing algorithm assigns

a  very high delay to Pathways which lose a lot of packets.  This

will tend to cause  internet  traffic  to  avoid  such  Pathways.

There  doesn't  seem to be any problem in principle with a scheme

like this, but we will  probably  need  to  do  some  statistical

analysis   in   order  to  determine  the  best  random  sampling

technique, and to figure out how many packets we  might  need  to

keep  track  of during some period of time (i.e., how big a table

do  we  need  to  keep  track  of  packets  which  are   awaiting


     This  sort  of random sampling can also be used as part of a

Pathway up/down protocol.  If a certain percentage of the sampled

packets do not get through, it might be good to assume  that  the

Pathway  is  not  of  sufficient  quality  to be operational, and

should appear to be down as far as the internet routing algorithm

is concerned.  In the absence of real data traffic, we could  run

the  up/down  protocol  with  randomly  generated  test  packets.

Randomly generated test traffic or randomly sampled data  traffic

will  give  us  a better result than periodic test traffic, since

measurements based on random  sampling  are  less  likely  to  be

correlated with other network phenomena.)
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     After  we compute the delay for individual packets, we still

face the following two questions:

     1) The delay  which  the  routing  algorithm  assigns  to  a

        particular  Pathway  will  be  a function of the measured

        delays of the individual packets sent  on  that  Pathway.

        But what function should it be?

     2) Once a  Switch  determines  the  delay  on  the  Pathways

        emanating  from itself, it must inform all other Switches

        of these values  (in  routing  updates).   What  protocol

        should it use for disseminating these updates?

The  second  question  will  be  discussed  in  section 4.5.  The

remainder of this section will deal with the first question.

     After  measuring  the  delays  of  individual  packets,  the

individual  delays  must  be  put  through some sort of smoothing

function before  they  can  be  used  as  input  to  the  routing

algorithm.   For  example,  in  the ARPANET, we take the average,

every 10 seconds, of the delays experienced by  all  the  packets

traversing  a  particular  line in the previous 10 seconds.  This

average is used as input to the routing algorithm  (i.e.,  it  is

assigned  as  the  "length"  of  the  line when the shortest-path

computation is run.)  We didn't choose this smoothing function at

random; we chose it because it  meets  certain  desiderata.   Our

real purpose in measuring delay on a particular line is to enable

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us  to  predict  the delay that will be seen by packets which are

routed over that line in the future.  Knowing the  average  delay

during  some  period in the past is of no value except insofar as

it enables us to make predictions about the future.  We found  in

the  ARPANET  that  for  a  given  level  of  traffic, the delays

experienced by the individual packets would vary quite a bit, but

the  delay  when  averaged  over  10  seconds  stayed  relatively

constant.  (It is interesting that everyone who does measurements

of  individual packet delay always discovers this large variance,

and always expresses great surprise.  This "surprising" result is

so often re-discovered that it should cease to  be  a  surprise.)

When designing the delay measurement routines for the ARPANET, we

investigated  some  other  smoothing functions (everyone seems to

have his own favorite), but none  gave  more  reasonable  results

than  the  simple  average  we  adopted  (which  is not a running

average, but rather starts  over  again  from  scratch  every  10

seconds).   We  also  tried  averaging  periods  of  less than 10

seconds, but found what we regarded as too much  variation,  even

when the traffic load was stable.

     Note  that if we take an average every 10 seconds, we cannot

react to a change of conditions in less than 10 seconds,  and  we

are often criticized by people who claim that it is important for

routing to be able to react more quickly.  Our reply, however, is

simply  that  it  takes  10  seconds  to  be  able  to  detect  a

significant change in delay.  Averages taken over smaller periods

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show too much variation under  constant  load  to  be  useful  in

predicting the future delay, and hence are not useful in routing.

In other words, averages taken over smaller periods give spurious

results,  "detecting"  changes  when  in fact there are none.  We

want to change routing in response to  real  changes  in  network

conditions, but not in response to the normal range of stochastic

variations  in delay.  Any change in routing made on the basis of

a shorter-term average is at least as likely to be harmful as  to

be helpful.  That is, if we attempt to make routing changes based

on  delay  data which is not sufficiently smoothed, we are really

making changes at random, since we  have  left  too  much  random

variation  in  the  delay data.  And it seems that a good routing

algorithm should not make changes at random.  Of course, it would

be nice if we could make routing changes instantaneously based on

instantaneously detected changes in real network conditions,  but

this is not possible simply because there is no instantaneous way

of detecting important changes in network conditions.

     It  is  important  to realize, however, that the measurement

periods in the various IMPs are  not  synchronized.   Although  a

given  IMP generates updates no more often than every 10 seconds,

some IMP or  other  is  generating  an  update  about  every  500

milliseconds.   Mathematical analysis indicates that synchronized

measurement and updating periods should be  avoided,  since  they

give worst case performance [4].

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     There  are  other  important  reasons for not making routing

changes too often.  During the lifetime of a single packet in the

network, we want the routing to be relatively constant,  so  that

the  packet can get to its destination without having to take too

many detours.  If we changed the routing every  millisecond,  for

example,  a  single  packet  in  transit though the network would

experience many routing changes while  in  transit,  which  would

probably  cause  it  to  have a longer delay than necessary.  The

rate at which we change routing should be  low  relative  to  the

average  transit  time  of a packet through the network.  Another

reason for not making routing changes too frequently  has  to  do

with  the  time  it  takes  routing  updates to travel around the

network.  We want to make sure that the information carried in  a

routing  update is not totally obsolete by the time the update is

received.  This implies that the  smoothing  interval  for  delay

measurements has to be long relative to the time it takes updates

to traverse the network.

     In the ARPANET, 10 seconds is much longer than the amount of

time  it  takes  to  get  updates around, or the amount of time a

packet spends in transit in the network.  We chose 10 seconds  as

the  averaging  interval  because  it  seemed  to be the shortest

period that was long enough to give us  a  reasonable  amount  of

smoothing.   If  we  think that in the internet, however, average

transit times might be measured in the tens of  seconds,  we  may

have  to  make our smoothing interval considerably longer than 10

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seconds, perhaps as long as a minute.  This could seriously limit

the responsiveness of the routing algorithm to  changing  network

conditions.  However, there is nothing we can do about this.  THE



simple  reason  that  it will just take longer to disseminate the

information needed for routing around the network.   The  transit

time  of a network places an upper limit on the responsiveness of

that network's routing algorithm.  Any  attempt  to  exceed  this

upper limit (with kludges or heuristics) will just be futile, and

will  result only in unstable and mysterious behavior on the part

of the  routing  algorithm,  reducing,  rather  than  increasing,


     This  is  not  to say that each Switch must generate routing

updates as often as every 10 seconds.  If there is no  change  in

delay  from  one  10-second  period  to another, then there is no

reason to generate an update.  Or if there is a change, but it is

not "significant", then there is no reason to generate an update.

In the ARPANET, a delay change is considered to be significant if

it exceeds a certain (parameterized)  threshold.   We  devised  a

scheme  wherein the threshold decreases with time, so that a very

large change is always  "significant",  but  a  small  change  is

significant  only  if  it  persists  for a long time.  Of course,

routing updates  must  be  generated  not  only  in  response  to

measured  changes in delay, but also if a line goes down or comes

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     We would expect that the details of  the  delay  measurement

and  smoothing  algorithms  will  have  to  be  different  in the

internet than in the ARPANET, but the principles  outlined  above

would  seem  to  apply in the internet environment also.  WE WILL



AS PATHWAYS  IN  THE  INTERNET,  and  it  may  be  that  somewhat

different  smoothing  algorithms  will  have  to  be used for the

different kinds of Pathways.  However, there doesn't seem  to  be

any   problem   in  principle  with  doing  this  sort  of  delay


     An interesting issue arises if a given pair of  gateways  is

connected  by  two  or  more distinct Pathways.  For example, two

gateways might both be connected to ARPANET and SATNET,  so  that

each  can  be  reached  from  the  other  by  either of those two

networks.  Or, a gateway might be multi-homed on the ARPANET,  so

that it has two distinct access lines over which it can reach all

the  other  ARPANET  gateways.   In  such  cases,  do  we want to

separately report the delay on each of the distinct Pathways,  or

do  we want (at the level of routing) to represent the connection

between each pair of gateways as a  single,  unique  line,  whose

delay  is  some  function  of  the delay of the distinct Pathways

which really exist?  This issue is a generalization of  an  issue

we  have  been looking at in the context of the ARPANET, which we

call "parallel trunking."  In parallel trunking, a single pair of

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IMPs is connected by two or more trunks, and the  same  issue  of

how  to  represent them in routing (as individual trunks, or as a

single, composite, trunk) arises.  When the trunks are  telephone

lines,  the problem is relatively easy to deal with.  Routing can

treat them as a single trunk, with a delay which is  the  average

delay  of  all packets sent over the composite trunk.  The actual

decision as to  which  particular  component  trunk  to  use  for

transmitting  a particular packet can be made locally, by the IMP

to which the parallel trunks are connected; there is no need  for

routing to play a role in this decision.

     In  the  case  where  the  parallel trunks are of comparable

lengths (so that there is not much difference in the  propagation

delays),  the  trunks  can  serve a common queue according to the

standard FIFO single-queue multiple-server  discipline.   If  the

trunks  are more heterogeneous, say one is a terrestrial line and

one  is  a  satellite  line,  a  somewhat  more  complex  queuing

discipline  is  required.   We  would  like  to  avoid  using the

satellite  line  until  the  load  is  such  that  if  only   the

terrestrial  line  were  used,  packets  would experience a delay

comparable to that they experience over the satellite line.  With

this sort of queuing discipline, packets sent  to  the  composite

line  experience  a  delay which is independent of the particular

component (land-line or satellite line) that they use.  That  is,

no  packet is forced to suffer the quarter-second satellite delay

unless the terrestrial line is so backed up that  the  delay  for

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packets  sent  over it is comparable to the delay of packets sent

over the satellite line.  This sort of scheme seems to ensure the

best delay performance for the composite trunk.   (Actually,  the

mathematics  of  queuing  theory  suggests that a smaller average

delay for the composite trunk might be achieved  by  starting  to

use  the  satellite  line  sooner.   That  is, a somewhat smaller

average delay might be achievable if a few packets  are  given  a

much  longer delay by being forced over the satellite line sooner

than they would be with  the  queuing  discipline  we  suggested.

Considerations  of fairness would seem to rule that out, however;

how would you like it if your data got a  much  higher  delay  so

that  someone  else's  could  get  a  slightly  smaller  one?  In

addition, the queuing  discipline  we  suggested  would  seem  to

produce a smaller variance in delays, thereby making the measured

average delay on the composite trunk a better predictor of future

performance,  and  the  better we can predict future performance,

the better performance our routing algorithm can provide.)

     Basically, there is no reason for routing to be aware that a

particular line consists of several  parallel  components  rather

than a single component, because, if the argument above is right,

any  decision  as  to  which  component  to  use can be best made

locally, at the IMP from which the parallel lines emanate.   That

is, the global routing algorithm cannot really make effective use

of  information about which lines consist of parallel components,

and should not be burdened with information that it  cannot  use.

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This  is  good,  since  the  SPF  algorithm  cannot really handle

parallel lines between a pair of Switches except by  representing

them  as  a single line.  (A careful study of the algorithm would

show that much of the algorithm's space and time efficiency would

be sacrificed if it had to be modified to handle parallel  trunks

as separate trunks.  Since this efficiency is the main thing that

recommends the SPF algorithm over other shortest-path algorithms,

we  must  be  sure that we don't destroy the effectiveness of the

algorithm by making poorly thought-out changes to it.)

     In the internet environment, however, we have a more complex

problem with parallel trunks than in the ARPANET.  The scheme  we

outlined  for using parallel trunks in the ARPANET depends on our

being able to know when the load on the composite trunk  is  such

that  exclusive  use  of  the faster component would cause delays

that are just as high as we get when we use the slower component.

This is not difficult to know if the components are  phone  lines

of one sort or another, since the relation between load and delay

is  pretty  well-defined  if  we know the length of the lines and

their capacity.  If the components  of  a  parallel  "trunk"  are

really  packet-switching  networks, however, it is much harder to

figure out which components are slow and which are fast,  and  it

is hard to figure out when the load on the fast component is such

that we have to start using the slow one.

     It  seems  that  by separately measuring the delays obtained

over the "parallel trunks" in the internet case, we ought  to  be

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able to devise some algorithm for splitting the traffic among the

parallel   components   in   a   way   which   gives   reasonable

delay/throughput performance.   However,  we  don't  yet  have  a

solution  to  this  problem,  which  we  will  put  aside for the

present.  Whatever  scheme  we  eventually  decide  on,  however,

should  be  compatible with treating the parallel components as a

single line at the level of routing.  Of course, if we decide  to

have different routes for different traffic types (say, excluding

satellite  networks  for  interactive traffic, but using them for

batch traffic), then the  problem  is  eased  somewhat  since  we

partially  solve  the  problem a priori.  There would still be no

need to represent the parallel lines as separate lines.   Rather,

we  would  represent  them as a single line, with different delay

characteristics for different traffic types.

4.5  Routing Updates

4.5.1  Overhead

     Everyone seems to be in agreement that the overhead  due  to

routing  updates  should  be kept low.  At least, no one seems to

advocate that the overhead should be made  high.   Unfortunately,

"apple pie" pronouncements like this aren't much help in actually

designing  a  routing  scheme.  In evaluating a routing algorithm

from the perspective of overhead, one must understand the way  in

which overhead is traded off against functionality.

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     One  advantage  of  the  SPF  routing  algorithm  is that it

provides a lot of handles that can be used to  control  overhead.

In SPF routing, a routing update generated by a particular Switch

identifies each neighbor of that Switch, and gives the delay over

the Pathway to that Switch.  Thus the size of an update generated

by a particular Switch is proportional to the number of neighbors

that  the  Switch  has,  generally a fairly small number (no more

than 5 in the ARPANET, and probably of a similar magnitude in the


     In the current Catenet routing algorithm, the  size  of  the

routing updates is a function of the total number of gateways (or

equivalently,  of  the  total  number  of  component networks), a

number which can increase by a great deal over the years.  In the

SPF algorithm, the size of the  updates  is  a  function  of  the

connectivity  of  the internet, which could not increase anywhere

near as much or as rapidly as the number of  gateways.   (In  the

two years that SPF has been running in the ARPANET, the number of

IMPs  has  increased  by  a  third, with another similar increase

expected in the next several months, while the connectivity,  and

hence the average update size, has remained relatively constant.)

This is important, since we wouldn't want to get ourselves into a

situation where the update size eventually becomes so big (due to

network  growth)  that we can no longer fit a whole update into a

single packet (a situation that was imminent during the last days

of the original ARPANET routing algorithm.)  In the internet, the

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maximum size of an update packet is constrained by the  component

network  which  has  the  smallest maximum packet size.  It seems

likely that any component network whose packets are large  enough

to  carry  the enormous TCP and IP headers should have no trouble

carrying the routing updates.

     The amount of overhead due to routing updates is not only  a

function  of  the  update  size,  but  also  of the rate at which

updates are generated.  In the ARPANET, since each  IMP  averages

the  delay  on  its  outgoing  lines over a period of 10 seconds,

changes in delay on the lines emanating  from  a  particular  IMP

cannot  occur,  by  definition,  more  often  than  once every 10

seconds.  In  addition  to  generating  updates  when  the  delay

changes,  updates  must  also  be generated when lines go down or

come up.  In the ARPANET, a line which goes down cannot  come  up

for at least 60 seconds.  So in an IMP with 5 neighbors, the most

updates  that  can be generated in a minute is 11 (due to each of

the lines either going down or coming up during the  minute,  for

5,  and a delay change every 10 seconds, for 6).  It is important

to note that this is the maximum rate at  which  updates  can  be

generated,  not  the  average rate.  Since IMPs need not generate

routing updates unless they have a "significant change" in  delay

to  report,  the average rate can be much lower.  In the ARPANET,

the average rate for generating updates is actually about one per

IMP per 40 seconds.  This is a very limited amount  of  overhead.

Of  course,  the  overhead  will  increase  as the number of IMPs

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increases, because there are just more IMPs to generate  updates.

However,  the  amount  of  overhead  is always under our control,

since  we  can  always  alter  the  averaging  interval,  or  the

threshold  of significant change in delay, to force updates to be

generated less frequently and thereby to reduce overhead.   These

same principles apply to the internet also, so it doesn't seem as

if we will be generating enormous amounts of routing overhead.

     There  are  some things we might want to do which would tend

to make the routing updates longer than so  far  indicated.   For

example,  if  we  defined  several  priorities  of traffic at the

internet  level,  and  mapped  these  priorities   to   different

priorities of some particular component network, we might want to

separately  measure  the  delay  across  that  network  for  each

priority.  We might also want to compute a separate set of routes

across the internet for each priority.  If we adopted  some  such

scheme,  we would need to report in each update several different

delays for each Pathway,  indexed  by  priority.   These  indexed

delays  could  then be used for computing a set of routing tables

indexed by priority, allowing traffic of different priorities  to

use  different  routes.   Of  course,  this  would  lengthen  the

updates, adding more  overhead.   Part  of  the  decision  as  to

whether to adopt such a scheme would involve an evaluation of the

trade-offs  between  the  cost of this increased overhead and the

benefit of the expected improvement in performance.  The  issues,

however,  are clear, and there are enough handles controlling the

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amount of overhead so that we can put into effect any decision we


     It is important to understand that  the  number  of  routing

updates  generated by a single internet event (such as the outage

of a gateway access line) is much less with SPF routing than with

the current Catenet routing algorithm.  In SPF routing,  a  given

event  causes  the  generation  of ONE routing update, which must

then be sent to every gateway (thereby  giving  each  gateway  an

up-to-date  copy  of the "whole picture").  On the other hand, in

the current Catenet routing algorithm, a  single  internet  event

causes  a  flurry  of  updates,  as all gateways send and receive

updates repeatedly to and from each neighbor, until  the  routing

tables  stabilize  and  the  process settles down.  This can take

quite a long time and quite a few updates,  particularly  if  the

number of gateways is large.

     In addition, in an internet with a large number of gateways,

the  updates  for  the current Catenet routing algorithm are very

much larger than the SPF updates would be.  It is clear that  the

routing overhead due to a single network event would be much less

with  SPF  than  it  currently  is.   However, if we plan to send

routing updates when delay changes, as opposed  to  just  when  a

gateway  access  line comes up or goes down (as at present), then

we will be generating updates in response to more network events.

This tends to drive the overhead up.  Again, the  trade-offs  are

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relatively  clear here;  the amount of overhead simply trades off

against the responsiveness of the routing algorithm  to  changing

network  conditions.   The  decision  as  to  how  to  draw  this

trade-off can be made as a policy decision, and can be changed if

performance considerations warrant it.  The  situation  with  the

current  Catenet  routing algorithm is quite different, since the

amount of overhead that it  generates  is  almost  impossible  to

compute.   In  that  algorithm,  the  number  of  routing updates

generated in response to a particular event depends on the  order

in  which  the  updates are processed by the individual gateways,

something that is essentially random and hence hard  to  predict.

The SPF algorithm has no such dependency.

     The  need for hysteresis in the Pathway up/down protocol run

between  neighboring   gateways   is   worth   emphasizing.    If

connections  between  neighboring gateways are allowed to come up

and go down with great frequency, causing a  constant  flurry  of

routing  changes,  packets  in  transit will bounce around a lot.

Putting a limit on the frequency  with  which  a  gateway-gateway

connection  can  change  state  is  needed  not only to limit the

amount of overhead generated, but also to give some stability  to

the  routing.   It  is  worth  noting  that the ARPANET, although

providing hysteresis in its own line up/down protocol,  does  not

provide any hysteresis in host up/downs.  Hosts are allowed to go

down  and  come  up repeatedly many times a minute, and this does

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result  in  problems,   causing   congestion   and   instability.

Hysteresis in the gateway's Pathway up/down protocol will have to

be ensured explicitly; we cannot rely on the ordinary host access

protocol  of  the component networks to do the right thing.  That

is, if a network interface goes down, we must keep it down for  a

period  of  time, even if the network itself allows the interface

to come back up immediately.

4.5.2  Protocol

     We turn now to the problem of how to disseminate the routing

updates around the Network Structure.  Remember that the  updates

generated  by  a particular Switch will contain information about

the delays to the  neighbors  of  that  Switch.   When  a  Switch

generates  an  update, it must broadcast that update to ALL other

Switches.  As a result, every single Switch will know the  values

of  delay  between every single pair of neighboring Switches.  It

is then straightforward to have each Switch run  a  shortest-path

algorithm  which determines the shortest path from itself to each

other Switch.  The basic idea is for  each  Switch  to  know  the

entire  topology  of  the Network Structure, so that the shortest

paths can be determined by a localized shortest  path  algorithm,

with  no need for a distributed computation.  In the ARPANET, the

IMPs do not start out with any knowledge of the  topology.   They

determine  who  their own neighbors are, and they reconstruct the

rest of the topology from the routing updates they receive.

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     It is possible to prove that, as long as all  Switches  have

the same information about the topology and the delays, then they

will  produce  routes  which are consistent and loop-free.  (That

is, the situation in which Switch A thinks its best path to B  is

through  C,  and  C  thinks  its best path to B is through A, can

never arise.)  However, if some routing updates somehow get  lost

before  being  received  by every single Switch, then there is no

guarantee of consistent loop-free routing.  In fact,  if  routing

updates  get  lost,  so  that  different  Switches have different

information about the topology or the  delays,  we  would  expect

long-term  routing  loops  to  arise, possibly making the Network

Structure useless for some period of time.  So the protocol  used

to  broadcast  the  routing  updates  needs  the highest possible

reliability.  Of course, it will always take some amount of  time

for  an  update to be broadcast around the Network Structure, and

during that time, some Switches will have received  it  and  some

not.   This  means  there  will always be a transient period when

routing loops  might  arise.   So  another  aim  of  the  routing

updating  protocol must be to keep this transient period as short

as possible.  In the ARPANET, we have an updating protocol  which

seems   to   provide  these  characteristics  of  extremely  high

reliability and low delay.  Some of its aspects adapt readily  to

the  internet,  but  others are more difficult to adapt.  In what

follows,  we  first  describe  the  ARPANET's  routing   updating

protocol, and then discuss its applicability to the internet.

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     Suppose  IMP  A  has  to  generate  a routing update, either

because of some "significant" change in the  measured  delay,  or

because of a line up/down state change.  Each update generated by

A  has  a sequence number, which is incremented by 1 for each new

update.  (In the ARPANET, we use 6-bit  sequence  numbers,  which

wrap around after 63.)  After creating the update, IMP A sends it

to  each of its neighbors.  The update is transmitted as a packet

of extremely high priority;  only the packets used  in  the  line

up/down  protocol  are  of  higher priority.  We use the notation

"A(n)" to refer to the update generated by IMP  A  with  sequence

number  n.   Now let's look at what happens when a copy of update

A(n) is received by an IMP B.   (IMP  B  is  intended  to  be  an

arbitrary  IMP  somewhere in the network, possibly identical to A

or to one of A's neighbors, but not necessarily so.)   If  B  has

never  received  an  update  from A before, it "accepts" A(n), by

which we mean that it (a) remembers in its tables that  the  most

recent  update  it  has  seen  from A is A(n) (i.e., the sequence

number n, the list of neighbors of A, and the delays  from  A  to

each  neighbor are stored in B's tables), (b) it forwards A(n) to

each of its neighbors,  including  the  one  from  which  it  was

received,  and  (c) the SPF algorithm is run to produce a new set

of paths, given the new delay and topology information  contained

in  A(n).   If  B  has  received  an  update  from  A  before, it

determines whether A(n) is more recent than  the  update  it  has

already  seen,  and  "accepts"  it  (as  just  defined) if it is;

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otherwise it simply  discards  A(n).   The  determination  as  to

whether  A(n) is more recent than some previously received update

A(m) is made by a sequence number comparison (which,  of  course,

must account for the fact that sequence numbers can wrap around);

A(n) is not considered to be more recent than itself.

     If  one  thinks a bit about this inductive definition of the

protocol, one sees that each IMP  in  the  network  will  receive

every  update  which is generated by any IMP, and further that it

will generally receive a copy of  each  update  on  each  of  its

lines.   This means of broadcasting an update from one IMP to all

other IMPs is called "flooding."  It is  highly  reliable,  since

updates  cannot  be  lost  in  the  network due to IMP crashes or

partitions.  If there is  any  path  at  all  between  two  IMPs,

flooding  will get the update from one to the other.  (Of course,

if there is no path at all from A to B, then updates  cannot  get

from one IMP to the other.  However, this is not a problem, since

if  traffic  from  A  cannot even reach B, then it cannot use B's

outgoing lines, so there is no need for A to know the  delays  of

B's  outgoing  lines  in  this  case.   In  saying  that flooding

prevents updates from getting lost due to network partitions,  we

are  thinking of the case where an update is in transit from A to

B when a partition forms, such that A  and  B  are  in  the  same

partition  segment,  but  the update is in a segment which is now

isolated from either A or B.  Flooding ensures delivery  in  this


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     Flooding  also  ensures  that  an  update  travels  over the

shortest (in terms of delay)  possible  path.   Basically,  every

possible  path  is  attempted,  so  the  update  necessarily gets

through first on the shortest path, by definition.  In  addition,

this means of transmitting routing updates does not depend in any

way  on  the routing algorithm itself.  Since routing updates are

sent out all lines, there is no  need  to  look  in  the  routing

tables   to  decide  where  to  send  the  routing  update.   The

transmission of routing updates is independent of routing,  which

eliminates   the  possibility  of  certain  sorts  of  disastrous

negative feedback.

     One might think that a protocol which sends a copy of  every

update on every line creates a tremendous amount of overhead.  In

the ARPANET, however, the average update packet size is 176 bits,

and  the  average  number  of  updates sent on each line (in each

direction) is less than 2 per second, for an average overhead  of

less  than 1% of a 50 kbps line.  And this is with almost 75 IMPs

generating updates.

     Of course, a protocol like flooding is only as  reliable  as

are  the  individual  point-to-point  transmissions  from  IMP to

neighboring IMP.  We ensure reliability  at  this  level  with  a

positive  acknowledgment  retransmission  scheme.  Note, however,

that no explicit acknowledgments are required.  If  IMP  X  sends

update  A(n)  to neighboring IMP Y, and then X receives from Y an

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update A(m), where A(m)  is  at  least  as  recent  as  A(n),  we

consider that Y has acknowledged X's transmission of A(n).  Since

an  IMP  which  accepts  an  update  sends  it  to all neighbors,

including the one from which it was received, in  general,  if  X

sends  A(n)  to Y, Y will send A(n) back to X, thereby furnishing

the acknowledgment.  We say "in general", since there is a little

further twist.  As another  reliability  feature,  we  make  each

update  carry  complete  information,  and forbid the carrying of

incremental information in updates.   That  is,  each  and  every

update  generated by an IMP A contains all the latest information

about A's neighbors and its delay to them, so  that  each  update

can  be  fully  understood  in  isolation from any that have gone

before.  This  means  that  if  update  A(n+1)  is  received  and

processed   by  some  IMP  B,  then  the  prior  update  A(n)  is

superfluous and can just be discarded by B.   In  particular,  if

IMP X sends A(n) to neighboring IMP Y while at the same time Y is

sending  A(n+1)  to X, then X can interpret the receipt of A(n+1)

from Y as an acknowledgment of the receipt of A(n); that is, X no

longer has to worry about retransmitting A(n), since that  update

is  no  longer needed by Y.  If no "acknowledgment" for an update

is received from a particular neighbor within a specified  amount

of  time,  the  update  is  retransmitted.  Of course, it must be

specially marked as a retransmission, so that the neighboring IMP

will always "acknowledge" it (by echoing it back),  even  if  the

neighbor  has  seen it before.  This is needed to handle the case

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where  the  update  got  through  the   first   time,   but   the

acknowledgment  did  not.   It  should also be noted that all the

information in a routing update must  be  stored  in  each  IMP's

tables  in  order to run the SPF computation.  This means that if

it is necessary to retransmit an update to a particular neighbor,

the update packet can be re-created from the tables;  it  is  not

necessary   to   buffer   the   original  update  packet  pending


     We must remember that if congestion forms in  some  part  of

the  network,  we want routing to be able to adapt in a way which

can route traffic around the congestion.  For this  to  have  any

hope  of  working,  we  must be sure that ROUTING UPDATES WILL BE


DATA  PACKETS.  Therefore, routing updates in the ARPANET are not

sent by the ordinary IMP-IMP  protocol,  which  provides  only  8

logical   channels  between  a  pair  of  IMPs.   That  would  be

disastrous, since congestion often causes all 8 logical  channels

to fill up and remain filled for some time, blocking further data

transmission  between  the IMPs.  Transmission of routing updates

must be done in a way  that  is  not  subject  to  this  sort  of

protocol  blocking  during  periods of congestion.  (This sort of

"out-of-band" signalling was quite easy to put into the  ARPANET.

However,  it  is worth noting that such protocols as HDLC make no

explicit provision for out-of-band signalling, and it seems  that

many  networks  are being built in which the routing updates will

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not be able to flow when the network gets  congested.   Designers

of  such  networks  will  no  doubt  be quite surprised when they

discover  what  is  inevitable,  namely    that   their   routing

algorithms  break down completely in the face of congestion.)  We

also want to be sure that we have enough  buffers  available  for

holding routing updates, and that we process them at a relatively

high CPU priority.

     There  is one more twist to the updating protocol, having to

do with network partitions.  A network partition is  a  situation

in which there are two IMPs in the network between which there is

no  communications path.  Network partitions,  in this sense, may

be as simple as the case in which some IMP is down (an IMP  which

is  down  has  no  communications  path  to any other IMP), or as

complex as  the  case  in  which  four  line  outages  result  in

partitioning  the  network  into  two  groups of 40 IMPs.  When a

partition ends, we have  to  be  sure  that  the  two  (or  more)

segments do not get logically rejoined until routing updates from

all  IMPs  in  each  segment  get  to  all  the IMPs in the other

segments.  That is, data packets must  not  be  routed  from  one

segment  to  the  other  until  all  IMPs  in  each  segment have

exchanged routing updates with all IMPs in  the  other  segments.

Otherwise, routing loops are sure to form.  We must also remember

that the sequence numbers of IMPs in one segment may have wrapped

around  several  times  during  the  duration  of  the partition.

Therefore we must ensure that IMPs in one segment  do  not  apply

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the  usual sequence number comparison to updates from IMPs in the

other segment.

     We have dealt with these problems by  adding  the  following

three time-outs to the updating protocol:


        to  generate at least one update every minute, whether or

        not there has been any change in delay or line state.

     2) MAXIMUM UPDATE LIFETIME: If an IMP B has not received any

        updates generated by IMP A for a  whole  minute,  then  B

        will  "accept" the next update it sees that was generated

        by A, regardless of the sequence number.

     3) WAITING PERIOD: When a line is ready to come  up,  it  is

        held in a special "waiting" state for a minute.  While in

        the  waiting  state,  no  data  can  be sent on the line.

        However, routing updates are passed over the line in  the

        normal way, as if the line were up.

     Since  the  ending  of a partition is always coincident with

some line's  coming  up,  these  three  features  ensure  that  a

partition cannot end until a full exchange of routing information

takes  place.   They also ensure (given the facts that there is a

6-bit sequence number space and that IMPs can generate at most 11

updates per minute) that sequence numbers  of  updates  generated

after  the  end  of  the partition are not compared with sequence

numbers of updates generated before the partition occurred.
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     The general idea of flooding the updates seems as  important

in the internet as in the ARPANET.  In general, we can expect the

internet  to  be  subject  to  many  more  mysterious outages and

disturbances than is the ARPANET, and the reliability  and  speed

of flooding will be essential if an internet routing algorithm is

to  have  any  hope  of  working.   The  issue of overhead may be

somewhat worrisome, though.  If an IMP has  to  send  each  of  4

neighbors a copy of each update, it is just a matter of sending a

copy of a small packet on each of 4 wideband lines.  On the other

hand,  if  a  gateway  has  to send a copy of each update to each

neighbor, this may mean that it has  to  send  4  copies  into  a

single  network,  over  a  single network interface.  This may be

somewhat more disruptive.  Of course, this problem only exists on

networks which do not have group addressing.  If a network allows

the gateways to be addressed as a group, then each gateway  needs

only  to  place one copy of each update into the network, and the

network will take responsibility for delivering it to each  other

gateway.  (This might result in each gateway's receiving back its

own  copy  of  the update, since the sending gateway will also be

part of the group, but that  is  no  problem.   As  long  as  the

gateway can identify itself as the transmitter, it can just throw

away  any  updates which it transmitted to itself.)  This idea of

sending the updates to all neighbors on a particular  network  by

using  group  addressing  fits  in well with an idea expounded in

section 4.1, namely the idea that a network  should  be  able  to

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tell which of its hosts are gateways, and should inform the other

gateways  when  a new gateway come up.  This same mechanism could

be used by the network to augment its group addressing mechanism,

to  allow  the  group  definition  to  change   dynamically   and

automatically  as  the  set  of gateways connected to it changes.

Unfortunately, few networks seem to have group addressing.   Even

SATNET has only a primitive group addressing feature, although it

seems  odd  to  have  a  broadcast  network  without  full  group

addressing capabilities.  (Group addressing is much more  complex

on  a  distributed  network  like  ARPANET  than  on  a broadcast

network.)  Perhaps as further internet development proceeds, more

of the component networks will add group addressing, in order  to

make their use of the internet more robust and efficient.

     Retransmission      of      routing     updates     on     a

gateway-to-neighboring-gateway basis, based on the scheme in  the

ARPANET,  also seems to offer no problems in principle.  However,

the retransmission time-outs might have to be  carefully  chosen,

and  tuned  to  the characteristics of the network connecting the

sending and receiving gateways.  The retransmission time  has  to

be  somewhat  longer  than  the  average round-trip delay in that

network, and this may vary considerably from network to  network.

In  principle,  however,  this  is no different from the ARPANET,

where  the  retransmission  timers  for  routing   updates   vary

according  to  the propagation delay of the phone line connecting

two IMPs.

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     There is a bit of a subtle problem that we discovered in the

ARPANET, having to do  with  the  scheme  of  using  the  updates

themselves   as   acknowledgments.   Suppose  Switch  A  has  two

neighbors, B and C.  A receives a copy of update u  from  B,  and

queues  it  for  transmission to C.  However, while u is still on

the queue to C, A receives a copy of u from C.  If A had  already

sent  u  to  C,  this  copy  from  C  would  have  served  as A's

acknowledgment that C had received the update.  But now,  with  u

on  the  queue  to  C,  if we are not careful, A will send u to C

after having received a copy of u from C.  When C gets this  copy

of  u  from  A it will not accept it (since it has already seen a

copy of u and sent that copy on to A),  which  will  cause  A  to

retransmit u to C, resulting in an unnecessary retransmission.

     In  the ARPANET, we deal with this problem by turning on the

retransmission timer as soon as an  update  is  received,  rather

than  when it is sent.  That way, an update which is still queued

for transmission when its "acknowledgment" is received will still

get transmitted unnecessarily, but the retransmission timer  gets

shut   off,  causing  only  one,  rather  than  two,  unnecessary

transmissions.  A more logical  scheme  would  be  to  check  the

transmission  queue  to  a  Switch whenever an update is received

from that Switch.  If a copy of the same  update  that  was  just

received  is  queued  for transmission, it should just be removed

from   the   queue.    This   would   prevent   any   unnecessary

transmissions.   In  the ARPANET, a few unnecessary transmissions

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don't really matter, but in the internet, if we  really  want  to

keep  the  overhead  low, it is probably worthwhile trying to get

this just right.  We must remember that network access  protocols

may  limit  the  number  of  packets  we can get into the network

during some period, which makes it  all  the  more  important  to

avoid sending unnecessary packets.

     Suppose  we find that for some reason or other, it is taking

a very long time to get updates from some gateway to one  of  its

neighbors.   This  would  show  up  as  an  excessive  number  of

retransmissions of updates.  In such a case,  we  would  probably

have  to  consider  that particular gateway-gateway Pathway to be

down, irrespective of what our ordinary Pathway up/down  protocol

tells  us.   Remember  that  in  order  to  ensure consistent and

loop-free routing, we must get the updates around the internet as

rapidly as  possible.   If  updates  cannot  travel  sufficiently

rapidly  on some Pathway, then we just cannot use that Pathway at

all for transit within the internet.   Attempting  to  keep  that

Pathway up for transit can result in relatively long-term routing

loops,  which  could  in  turn  cause  a  degradation  in network

performance which swamps the degradation caused by not using that

Pathway at all.  Especially disastrous would be  a  situation  in

which  ordinary data packets could pass, but routing updates, for

some reason, could not.  It is hard to know what might cause such

a situation (perhaps a bug in the component network that  we  are

using  as  a  Pathway),  but it is certainly something we need to

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protect against.  (Note, however, that even if  we  declare  some

gateway-gateway Pathway down, it does not follow that the network

underlying  that Pathway cannot be used as a terminus network, to

which data for Hosts can be sent and from which data  from  Hosts

can  be  received.   Even  if  some  network  is  not  usable for

providing a Pathway between two gateways on it, it may  still  be

useful  for providing a Pathway between the gateways and some set

of Hosts.)

     We have emphasized the need to transmit routing  updates  as

"out-of-band"  signals,  which bypass the ordinary communications

protocols (such as the IMP-IMP protocol in the ARPANET), so  that

when  congestion forms which causes those protocols to block, the

routing updates can still flow.  That is, we would like to have a

protocol which is both non-blocking and non-refusing.   This  may

be  quite difficult to achieve in the internet environment, where

sending an update from gateway to  gateway  requires  us  to  use

whatever  network  access  protocol  is  provided  by the Pathway

network.  Here our most  difficult  problem  might  be  with  the

ARPANET's  1822 protocol, which can cause blocking of the network

interface for tens of seconds.  We really can't delay  sending  a

routing update for 15 seconds or so while the IMP is blocking, so

whenever  this happens we would have to declare the pathway down.

     In the ARPANET, we have two ways  of  trying  to  deal  with

this.   One  way would be to send all packets into the ARPANET as

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datagrams, which cannot cause blocking.  Another way would be  to

use  the standard virtual circuit interface, but to obey the flow

control restrictions of the ARPANET (i.e., to control the  number

of  outstanding  messages  between a pair of hosts), and to avoid

the use of multi-packet messages (which can cause blocking if the

destination IMP is short of buffers, as ARPANET IMPs  chronically

are).   There  are  other situations in which blocking can occur,

but they all involve a shortage of resources at the  source  IMP,

and  in  such  cases declaring the Pathway to be down is probably

the right thing to  do.   We  do  not  want  to  be  forced  into

declaring  Pathways  down  simply  because  we  have ignored some

protocol restriction, but it seems much more sensible to  declare

a Pathway down if, say, the IMP to which a gateway is attached is

too  congested  to provide reliable service for internet packets.

     It is important to note that whatever restrictions we  apply

to  our  use  of  the  network  access protocol apply not only to

routing updates, but also to all messages sent into  the  ARPANET

from  the  gateway.  It would do no good, for example, to send in

routing updates as datagrams, while using non-datagrams for other

packets, since this would allow the other packets  to  block  the

routing  updates.  At this point, it is not quite clear just what

the best scheme would be.  The use of datagrams enables us to get

around  the  sometimes  time-consuming  but   often   unnecessary

resequencing which the ARPANET performs before delivering packets

to  the  destination  host (it is neither necessary nor desirable

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for the ARPANET to resequence routing updates  before  delivering

them  to  a  gateway),  but  it  also  reduces the reliability of

transmission through the ARPANET, and it is not obvious how  this

trades  off.   For  each  network  which  we  intend  to use as a

component of the internet, we will have to  carefully  study  the

details  of  its  network  access  protocol, and possibly do some

experiments to see how the  various  details  of  network  access

affect  the  performance,  in  terms  of  delay,  throughput, and

reliability of the network.  Only by  careful  attention  to  the

details  of  network  access  on  each particular network, and by

continuing measurements and instrumentation in  the  gateways  to

see if we are getting the expected performance from the component

networks, can we hope to make the routing updating protocol quick

and  reliable  enough  to ensure consistent and loop-free routing

throughput the internet.  There are a few general  principles  we

might  appeal  to,  such as making routing updates be the highest

priority traffic  that  we  send  into  the  component  networks.

However,  it is difficult to be sure a priori what effect even so

straightforward a principle might have.  It's not hard to imagine

a poorly designed network in which low priority  packets  receive

better   performance  than  high  priority  ones,  under  certain

circumstances.  To make the internet robust, we need to  be  able

to  detect  such  situations  (and to gather enough evidence, via

measurements, to enable us to point the finger convincingly), and

we cannot simply assume that a component network will perform  as

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     If  we  might  digress  a  little, the considerations of the

preceding paragraphs raise an interesting issue with  respect  to

the  use  of  fragmentation  in  the  gateways.   We  raised  the

possibility of not using multi-packet ARPANET messages, and  such

a  strategy  would  doubtless  require more fragmentation than is

presently done.  Fragmentation in  the  gateways  has  long  been

thought  of  as a necessary evil, necessary because some networks

have a smaller maximum packet size than  others.   If  a  gateway

receives  a  packet from network A which is too large to fit into

network B, then the gateway must either fragment it or drop it on

the floor.  However, perhaps fragmentation is sometimes useful as

an optimization procedure.  That is,  some  network  may  have  a

suitably  large  maximum  packet  size  so that fragmentation is,

strictly speaking, unnecessary.  Nevertheless, the network  might

actually  perform  better  if  given  smaller  packets,  so  that

fragmentation provides better performance.  We see this  in  some

current  Catenet problems.  It seems that the BBN-gateway between

ARPANET and SATNET often receives packets from SATNET  which  are

2000  bits  long,  or  twice  the size of an ARPANET packet.  The

gateway then presents these messages to the ARPANET as two-packet

messages.  As it happens, two-packet messages generally give  the

lowest  possible  throughput on the ARPANET (a consequence of the

limited buffer space at the destination IMPs and  the  fact  that

the ARPANET assumes that all multi-packet messages will contain 8

packets);  the  gateway  could probably obtain better performance

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from the ARPANET by fragmenting the two-packet message  into  two

single-packet  messages.   Of course, the situation is a bit more

complicated in general than this may make it seem.   If  messages

are being sent from a source host through SATNET and then through

ARPANET  to  a  destination  host, best performance might well be

achieved by sending the messages  as  2000-bit  messages  through

SATNET,  then  fragmenting  them  and  sending  them  as 1000-bit

messages through ARPANET.  However, what if the messages must  go

beyond  ARPANET,  through another network, which handles 2000-bit

messages more efficiently than 1000-bit messages?  This  sort  of

strategy,  if useful at all, is best done in combination with the

hop-by-hop fragmentation/reassembly scheme suggested in IEN  187.

     The  part  of the routing updating protocol which deals with

recovery  from  partitions  (including  the  degenerate  case  of

initialization when a Switch comes up) is somewhat more tricky to

apply  to  the  internet  environment.  In the ARPANET, we have a

number of one-minute timers.  Each IMP must generate an update at

least once per minute; a line that  is  ready  to  come  up  must

participate  in  the  updating protocol for a minute before being

declared up; and an update that has been held for a minute in  an

IMP,  with  no  later update from that update's source IMP having

been seen, is regarded as "old", in the sense that  its  sequence

number  is  no longer considered when the IMP is deciding whether

the next update it sees (from the same source) is acceptable.  In

attempting to adapt these procedures to  the  internet,  we  must

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take  notice  of the way in which these timers interact with each

other and with other features of  the  internet.   Consider,  for

example,  the  length  of  the  maximum  update  lifetime,  which

determines how long an update's sequence number remains valid for

the purposes of judging the acceptability  of  the  next  update.

There are two restrictions on the length of this timer:

     1) A Switch A should not time out  an  update  whose  source

        Switch  is  B  unless  there  really is a partition which

        destroys  the  communication  path  between   A   and   B

        (remember,   this  includes  the  degenerate  case  of  a

        partition where B simply goes down).  This means that the

        time-out period must be  greater  than  the  sum  of  the

        maximum  interval between updates PLUS the maximum amount

        of time that an update from B could take to get to A.

     2) The sequence numbering scheme used for the  updates  must

        be such that the sequence numbers cannot wrap around in a

        period  of  time  which  is  less than the maximum update


     In the ARPANET, the sequence numbers  cannot  wrap  in  less

than  a  few  minutes, each IMP generates an update at least once

per minute, and the time to get that update to all other IMPs  is

negligible  when  compared  to  a  minute,  so  a  maximum update

lifetime of one minute is fine.  In  the  internet,  however,  we

could  not  expect  to  measure  transit times in the hundreds of

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milliseconds; tens of seconds would be more like it.  So even  if

we  forced  each  gateway  to  generate  at  least one update per

minute, we would still need a maximum update lifetime of  several

minutes.   And the longer our maximum update lifetime, the larger

our sequence number space must be (to prevent wrap-around), which

means additional overhead (memory and bandwidth) to represent the

sequence numbers.

     A similar constraint applies to the "waiting  period".   The

purpose   of  the  waiting  period  is  to  ensure  that  when  a

gateway-gateway Pathway is ready to come up, it is not  permitted

to  carry  data until an update from each other gateway traverses

it.  Clearly, for this to have the  proper  effect,  the  waiting

period  must  be  longer than the sum of the maximum transit time

plus the maximum interval between the generation of updates  from

a  single  gateway.   We  would probably also have to set this to

several  minutes.   This  does   have   a   serious   operational

consequence,  namely  that  no  outage will persist for less than

several minutes.  This can be an inconvenience,  lengthening  the

time  it  takes  to  put  out  a  new software release to all the

gateways,  for  example,  and   possibly   affecting   the   MTTR

statistics, but it is something we just have to live with.  Note,

by  the  way,  that  as long as the waiting period is at least as

long as the maximum update  lifetime,  a  gateway  that  restarts

after  a  failure (or a reload) can start generating updates with

sequence number 0, irrespective of what sequence numbers  it  was

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using before, since all its prior updates will have timed out (if

the timers are set right).

4.6  Limitations of Internetting

     This  discussion  of routing in the internet points out some

of  the  inherent  limits  of  internetting.   Good   performance

requires the use of a routing updating procedure which broadcasts

the  updates  in a very reliable and quick manner.  Anything that

delays the routing updates, or makes their transmission less than

reliable, will lengthen the amount of time during which different

Switches have a different "picture"  of  the  Network  Structure,

which  in  turn  will  degrade  performance.  We believe that the

updating protocol we  developed  for  the  ARPANET  solves  these

problems in the context of the ARPANET.  It seems clear, however,

that  broadcasting  routing updates in the internet is just going

to be slower and  less  reliable  than  it  is  in  the  ARPANET.

Although  the  same  principles  seem to apply in both cases, the

characteristics of the internet  Pathways  are  not  sufficiently

stable  to  ensure the speed and reliability that we really would

like to have.  It is going to be very hard to ensure that we  can

get our routing updates through the various component networks of

the  internet in a timely and reliable manner, and it may be hard

to get the component networks  to  handle  the  internet  routing

updates  with  enough priority to prevent them from being blocked

due to congestion.  This is going to place a  limit  on  internet

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performance  which we cannot avoid no matter what architecture we


     The only way to eliminate this sort of problem would  be  to

have  the component networks themselves give special treatment to

internet control packets, such as  routing  updates.   Currently,

the  component  networks  of  the internet treat internet control

packets as mere data.  We have suggested that in some  cases,  it

is  impossible  to meet certain of our goals without special help

from the underlying networks.  For example, in our discussion  of

the  "gateway  discovery protocol", we argued that preserving the

maximum  flexibility  for  making  topological  changes  in   the

internet requires cooperation from the underlying networks.  This

point  can  be  generalized, though.  The more cooperation we can

get from the underlying networks, the  better  we  can  make  our

internet  routing  algorithm  perform, and the better we can make

the internet perform.  We would recommend therefore that  serious

consideration be given to modifying the component networks of the

Catenet to maximize their cooperation with the internet.

     Even  if the component networks of the internet cooperate to

the fullest,  there  is  another  problem  which  may  limit  the

responsiveness  of  the internet routing algorithm.  If there are

very long transit times across the internet, much longer than  we

ever  see  in  individual  networks  like  the  ARPANET, then the

responsiveness of routing is necessarily held down.  This  factor

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will  place  a natural restriction on the growth of the internet.

At a certain point, it will become just too big to be treated  as

a  single  Network  Structure,  so that further growth would make

routing too non-responsive to provide  good  service.   That  is,

eventually  we reach a point of diminishing returns, where adding

more Switches, or even adding more levels of hierarchy, begins to

significantly degrade service throughout the internet  by  making

the  routing  algorithm  too  non-responsive.  It is important to

understand that the notion of "big" here has nothing to  do  with

the  number  of Switches, but rather with the transit time across

the internet.

     If there are two Hosts which cannot, for reasons like  this,

be placed on the same internet, it may still be possible for them

to communicate, though at a somewhat reduced level of efficiency.

Each  of  the  Hosts  would  have to be on some internet, but not

necessarily on the same one.  Suppose, for  example,  that  there

are  two  different  internets,  internet A and internet B, which

cannot be combined into one larger internet because the resultant

internet would be too large to  permit  a  reasonably  responsive

routing  algorithm.   However,  it  is  still  possible  for each

internet to model the other one as an  Access  Pathway.   Suppose

that  Host  H1 on internet A needs to communicate with Host H2 on

internet B.  Then if a Switch SA of internet A can  be  connected

to  a  Switch SB of internet B, the internet A can represent Host

H2 as being homed to its Switch  SA,  via  a  Pathway  (of  whose

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internal  structure  it is unaware) which is actually internet B.

A corresponding mapping can  be  made  in  the  other  direction,

permitting  full-duplex communication.  However, neither internet

could use the other as an internal (i.e., Switch-Switch) Pathway,

or  the   resulting   configuration   would   be   insufficiently

responsive.   (This  may seem akin to the regionalization against

which  we  argued  in  section  4.3.4.   However,  since  neither

internet  uses  the  other  as  an internal Pathway, there are no

problems of looping.)  Naturally,  just  as  Hosts  on  a  common

network  can expect to get more efficient communications than can

Hosts which must communicate over an internet, Hosts on a  common

internet  will  get more efficient communications than will hosts

on different internets.

     There are other reasons besides non-responsiveness which may

make it imperative to have separate internets  which  cannot  use

each  other  as  internal  Pathways.   For example, two internets

might cover the same "territory,"  geographically  speaking,  but

may  be  under the control of two different organizations, or may

use essentially different  algorithms  or  protocols.   In  fact,

several  different  internets  might  even  cover the same set of

Hosts, and consist of the same set of component  packet-switching

networks.   (It  is  important  to remember that it is the set of

gateways which constitute the internet, not the set of  component

networks.  Imagine if every ARPA-controlled network had a Brand X

gateway  and a Brand Y gateway.  Then there would be two separate

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internets, Brand X and Brand Y, which are logically  rather  than

physically  separate.)   Our  procedure  of  having each internet

regard the other as an Access Pathway to a set of Hosts, but  not

as  an  Internal Pathway, allows communication among Hosts on the

different internets, without introducing problems of looping, and

while preserving the maintainability of the individual internets.

Of course if the two internets have different  access  protocols,

then  the Switches of one or the other internet (or both) must be

prepared to translate from one protocol to the other, but that is

a simpler problem than the ones we have been dealing with.

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1.  J.M. McQuillan, I.  Richer,  E.C.  Rosen,  "The  New  Routing
    Algorithm    for   the   ARPANET,"   IEEE   TRANSACTIONS   ON

2.  E.C. Rosen, "The Updating Protocol of ARPANET's  New  Routing
    Algorithm," COMPUTER NETWORKS, February 1980.

3.  J.M  McQuillan,  I.  Richer,  E.C.  Rosen,  ARPANET   ROUTING
    BBN Report No. 3803, April 1978.

4.  J.M.  McQuillan,  I.  Richer,  E.C.  Rosen,  D.P.  Bertsekas,
    TECHNICAL REPORT, BBN Report No. 3940, October 1978.

5.  E.C. Rosen, J.G. Herman, I. Richer, J.M.  McQuillan,  ARPANET
    REPORT, BBN Report No. 4088, March 1979.

6.  E.C. Rosen, J. Mayersohn,  P.J.  Sevcik,  G.J.  Williams,  R.
    Report No. 4473, August 1980.

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